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7ea96eefb0
BFQ may merge a new bfq_queue, stably, with the last bfq_queue
created. In particular, BFQ first waits a little bit for some I/O to
flow inside the new queue, say Q2, if this is needed to understand
whether it is better or worse to merge Q2 with the last queue created,
say Q1. This delayed stable merge is performed by assigning
bic->stable_merge_bfqq = Q1, for the bic associated with Q1.
Yet, while waiting for some I/O to flow in Q2, a non-stable queue
merge of Q2 with Q1 may happen, causing the bic previously associated
with Q2 to be associated with exactly Q1 (bic->bfqq = Q1). After that,
Q2 and Q1 may happen to be split, and, in the split, Q1 may happen to
be recycled as a non-shared bfq_queue. In that case, Q1 may then
happen to undergo a stable merge with the bfq_queue pointed by
bic->stable_merge_bfqq. Yet bic->stable_merge_bfqq still points to
Q1. So Q1 would be merged with itself.
This commit fixes this error by intercepting this situation, and
canceling the schedule of the stable merge.
Fixes: 430a67f9d6 ("block, bfq: merge bursts of newly-created queues")
Signed-off-by: Pietro Pedroni <pedroni.pietro.96@gmail.com>
Signed-off-by: Paolo Valente <paolo.valente@linaro.org>
Link: https://lore.kernel.org/r/20210512094352.85545-2-paolo.valente@linaro.org
Signed-off-by: Jens Axboe <axboe@kernel.dk>
7361 lines
254 KiB
C
7361 lines
254 KiB
C
// SPDX-License-Identifier: GPL-2.0-or-later
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/*
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* Budget Fair Queueing (BFQ) I/O scheduler.
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*
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* Based on ideas and code from CFQ:
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* Copyright (C) 2003 Jens Axboe <axboe@kernel.dk>
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*
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* Copyright (C) 2008 Fabio Checconi <fabio@gandalf.sssup.it>
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* Paolo Valente <paolo.valente@unimore.it>
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*
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* Copyright (C) 2010 Paolo Valente <paolo.valente@unimore.it>
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* Arianna Avanzini <avanzini@google.com>
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*
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* Copyright (C) 2017 Paolo Valente <paolo.valente@linaro.org>
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*
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* BFQ is a proportional-share I/O scheduler, with some extra
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* low-latency capabilities. BFQ also supports full hierarchical
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* scheduling through cgroups. Next paragraphs provide an introduction
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* on BFQ inner workings. Details on BFQ benefits, usage and
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* limitations can be found in Documentation/block/bfq-iosched.rst.
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*
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* BFQ is a proportional-share storage-I/O scheduling algorithm based
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* on the slice-by-slice service scheme of CFQ. But BFQ assigns
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* budgets, measured in number of sectors, to processes instead of
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* time slices. The device is not granted to the in-service process
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* for a given time slice, but until it has exhausted its assigned
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* budget. This change from the time to the service domain enables BFQ
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* to distribute the device throughput among processes as desired,
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* without any distortion due to throughput fluctuations, or to device
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* internal queueing. BFQ uses an ad hoc internal scheduler, called
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* B-WF2Q+, to schedule processes according to their budgets. More
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* precisely, BFQ schedules queues associated with processes. Each
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* process/queue is assigned a user-configurable weight, and B-WF2Q+
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* guarantees that each queue receives a fraction of the throughput
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* proportional to its weight. Thanks to the accurate policy of
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* B-WF2Q+, BFQ can afford to assign high budgets to I/O-bound
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* processes issuing sequential requests (to boost the throughput),
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* and yet guarantee a low latency to interactive and soft real-time
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* applications.
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*
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* In particular, to provide these low-latency guarantees, BFQ
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* explicitly privileges the I/O of two classes of time-sensitive
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* applications: interactive and soft real-time. In more detail, BFQ
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* behaves this way if the low_latency parameter is set (default
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* configuration). This feature enables BFQ to provide applications in
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* these classes with a very low latency.
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*
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* To implement this feature, BFQ constantly tries to detect whether
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* the I/O requests in a bfq_queue come from an interactive or a soft
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* real-time application. For brevity, in these cases, the queue is
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* said to be interactive or soft real-time. In both cases, BFQ
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* privileges the service of the queue, over that of non-interactive
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* and non-soft-real-time queues. This privileging is performed,
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* mainly, by raising the weight of the queue. So, for brevity, we
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* call just weight-raising periods the time periods during which a
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* queue is privileged, because deemed interactive or soft real-time.
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*
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* The detection of soft real-time queues/applications is described in
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* detail in the comments on the function
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* bfq_bfqq_softrt_next_start. On the other hand, the detection of an
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* interactive queue works as follows: a queue is deemed interactive
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* if it is constantly non empty only for a limited time interval,
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* after which it does become empty. The queue may be deemed
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* interactive again (for a limited time), if it restarts being
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* constantly non empty, provided that this happens only after the
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* queue has remained empty for a given minimum idle time.
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*
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* By default, BFQ computes automatically the above maximum time
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* interval, i.e., the time interval after which a constantly
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* non-empty queue stops being deemed interactive. Since a queue is
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* weight-raised while it is deemed interactive, this maximum time
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* interval happens to coincide with the (maximum) duration of the
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* weight-raising for interactive queues.
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*
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* Finally, BFQ also features additional heuristics for
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* preserving both a low latency and a high throughput on NCQ-capable,
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* rotational or flash-based devices, and to get the job done quickly
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* for applications consisting in many I/O-bound processes.
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*
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* NOTE: if the main or only goal, with a given device, is to achieve
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* the maximum-possible throughput at all times, then do switch off
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* all low-latency heuristics for that device, by setting low_latency
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* to 0.
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*
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* BFQ is described in [1], where also a reference to the initial,
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* more theoretical paper on BFQ can be found. The interested reader
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* can find in the latter paper full details on the main algorithm, as
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* well as formulas of the guarantees and formal proofs of all the
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* properties. With respect to the version of BFQ presented in these
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* papers, this implementation adds a few more heuristics, such as the
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* ones that guarantee a low latency to interactive and soft real-time
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* applications, and a hierarchical extension based on H-WF2Q+.
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*
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* B-WF2Q+ is based on WF2Q+, which is described in [2], together with
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* H-WF2Q+, while the augmented tree used here to implement B-WF2Q+
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* with O(log N) complexity derives from the one introduced with EEVDF
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* in [3].
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*
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* [1] P. Valente, A. Avanzini, "Evolution of the BFQ Storage I/O
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* Scheduler", Proceedings of the First Workshop on Mobile System
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* Technologies (MST-2015), May 2015.
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* http://algogroup.unimore.it/people/paolo/disk_sched/mst-2015.pdf
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*
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* [2] Jon C.R. Bennett and H. Zhang, "Hierarchical Packet Fair Queueing
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* Algorithms", IEEE/ACM Transactions on Networking, 5(5):675-689,
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* Oct 1997.
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*
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* http://www.cs.cmu.edu/~hzhang/papers/TON-97-Oct.ps.gz
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*
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* [3] I. Stoica and H. Abdel-Wahab, "Earliest Eligible Virtual Deadline
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* First: A Flexible and Accurate Mechanism for Proportional Share
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* Resource Allocation", technical report.
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*
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* http://www.cs.berkeley.edu/~istoica/papers/eevdf-tr-95.pdf
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*/
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#include <linux/module.h>
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#include <linux/slab.h>
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#include <linux/blkdev.h>
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#include <linux/cgroup.h>
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#include <linux/elevator.h>
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#include <linux/ktime.h>
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#include <linux/rbtree.h>
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#include <linux/ioprio.h>
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#include <linux/sbitmap.h>
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#include <linux/delay.h>
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#include <linux/backing-dev.h>
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#include <trace/events/block.h>
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#include "blk.h"
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#include "blk-mq.h"
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#include "blk-mq-tag.h"
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#include "blk-mq-sched.h"
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#include "bfq-iosched.h"
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#include "blk-wbt.h"
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#define BFQ_BFQQ_FNS(name) \
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void bfq_mark_bfqq_##name(struct bfq_queue *bfqq) \
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{ \
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__set_bit(BFQQF_##name, &(bfqq)->flags); \
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} \
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void bfq_clear_bfqq_##name(struct bfq_queue *bfqq) \
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{ \
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__clear_bit(BFQQF_##name, &(bfqq)->flags); \
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} \
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int bfq_bfqq_##name(const struct bfq_queue *bfqq) \
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{ \
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return test_bit(BFQQF_##name, &(bfqq)->flags); \
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}
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BFQ_BFQQ_FNS(just_created);
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BFQ_BFQQ_FNS(busy);
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BFQ_BFQQ_FNS(wait_request);
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BFQ_BFQQ_FNS(non_blocking_wait_rq);
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BFQ_BFQQ_FNS(fifo_expire);
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BFQ_BFQQ_FNS(has_short_ttime);
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BFQ_BFQQ_FNS(sync);
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BFQ_BFQQ_FNS(IO_bound);
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BFQ_BFQQ_FNS(in_large_burst);
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BFQ_BFQQ_FNS(coop);
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BFQ_BFQQ_FNS(split_coop);
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BFQ_BFQQ_FNS(softrt_update);
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#undef BFQ_BFQQ_FNS \
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/* Expiration time of async (0) and sync (1) requests, in ns. */
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static const u64 bfq_fifo_expire[2] = { NSEC_PER_SEC / 4, NSEC_PER_SEC / 8 };
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/* Maximum backwards seek (magic number lifted from CFQ), in KiB. */
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static const int bfq_back_max = 16 * 1024;
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/* Penalty of a backwards seek, in number of sectors. */
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static const int bfq_back_penalty = 2;
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/* Idling period duration, in ns. */
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static u64 bfq_slice_idle = NSEC_PER_SEC / 125;
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/* Minimum number of assigned budgets for which stats are safe to compute. */
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static const int bfq_stats_min_budgets = 194;
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/* Default maximum budget values, in sectors and number of requests. */
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static const int bfq_default_max_budget = 16 * 1024;
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/*
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* When a sync request is dispatched, the queue that contains that
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* request, and all the ancestor entities of that queue, are charged
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* with the number of sectors of the request. In contrast, if the
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* request is async, then the queue and its ancestor entities are
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* charged with the number of sectors of the request, multiplied by
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* the factor below. This throttles the bandwidth for async I/O,
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* w.r.t. to sync I/O, and it is done to counter the tendency of async
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* writes to steal I/O throughput to reads.
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*
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* The current value of this parameter is the result of a tuning with
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* several hardware and software configurations. We tried to find the
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* lowest value for which writes do not cause noticeable problems to
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* reads. In fact, the lower this parameter, the stabler I/O control,
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* in the following respect. The lower this parameter is, the less
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* the bandwidth enjoyed by a group decreases
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* - when the group does writes, w.r.t. to when it does reads;
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* - when other groups do reads, w.r.t. to when they do writes.
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*/
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static const int bfq_async_charge_factor = 3;
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/* Default timeout values, in jiffies, approximating CFQ defaults. */
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const int bfq_timeout = HZ / 8;
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/*
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* Time limit for merging (see comments in bfq_setup_cooperator). Set
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* to the slowest value that, in our tests, proved to be effective in
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* removing false positives, while not causing true positives to miss
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* queue merging.
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*
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* As can be deduced from the low time limit below, queue merging, if
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* successful, happens at the very beginning of the I/O of the involved
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* cooperating processes, as a consequence of the arrival of the very
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* first requests from each cooperator. After that, there is very
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* little chance to find cooperators.
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*/
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static const unsigned long bfq_merge_time_limit = HZ/10;
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static struct kmem_cache *bfq_pool;
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/* Below this threshold (in ns), we consider thinktime immediate. */
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#define BFQ_MIN_TT (2 * NSEC_PER_MSEC)
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/* hw_tag detection: parallel requests threshold and min samples needed. */
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#define BFQ_HW_QUEUE_THRESHOLD 3
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#define BFQ_HW_QUEUE_SAMPLES 32
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#define BFQQ_SEEK_THR (sector_t)(8 * 100)
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#define BFQQ_SECT_THR_NONROT (sector_t)(2 * 32)
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#define BFQ_RQ_SEEKY(bfqd, last_pos, rq) \
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(get_sdist(last_pos, rq) > \
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BFQQ_SEEK_THR && \
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(!blk_queue_nonrot(bfqd->queue) || \
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blk_rq_sectors(rq) < BFQQ_SECT_THR_NONROT))
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#define BFQQ_CLOSE_THR (sector_t)(8 * 1024)
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#define BFQQ_SEEKY(bfqq) (hweight32(bfqq->seek_history) > 19)
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/*
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* Sync random I/O is likely to be confused with soft real-time I/O,
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* because it is characterized by limited throughput and apparently
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* isochronous arrival pattern. To avoid false positives, queues
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* containing only random (seeky) I/O are prevented from being tagged
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* as soft real-time.
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*/
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#define BFQQ_TOTALLY_SEEKY(bfqq) (bfqq->seek_history == -1)
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/* Min number of samples required to perform peak-rate update */
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#define BFQ_RATE_MIN_SAMPLES 32
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/* Min observation time interval required to perform a peak-rate update (ns) */
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#define BFQ_RATE_MIN_INTERVAL (300*NSEC_PER_MSEC)
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/* Target observation time interval for a peak-rate update (ns) */
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#define BFQ_RATE_REF_INTERVAL NSEC_PER_SEC
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/*
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* Shift used for peak-rate fixed precision calculations.
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* With
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* - the current shift: 16 positions
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* - the current type used to store rate: u32
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* - the current unit of measure for rate: [sectors/usec], or, more precisely,
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* [(sectors/usec) / 2^BFQ_RATE_SHIFT] to take into account the shift,
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* the range of rates that can be stored is
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* [1 / 2^BFQ_RATE_SHIFT, 2^(32 - BFQ_RATE_SHIFT)] sectors/usec =
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* [1 / 2^16, 2^16] sectors/usec = [15e-6, 65536] sectors/usec =
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* [15, 65G] sectors/sec
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* Which, assuming a sector size of 512B, corresponds to a range of
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* [7.5K, 33T] B/sec
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*/
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#define BFQ_RATE_SHIFT 16
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/*
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* When configured for computing the duration of the weight-raising
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* for interactive queues automatically (see the comments at the
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* beginning of this file), BFQ does it using the following formula:
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* duration = (ref_rate / r) * ref_wr_duration,
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* where r is the peak rate of the device, and ref_rate and
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* ref_wr_duration are two reference parameters. In particular,
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* ref_rate is the peak rate of the reference storage device (see
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* below), and ref_wr_duration is about the maximum time needed, with
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* BFQ and while reading two files in parallel, to load typical large
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* applications on the reference device (see the comments on
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* max_service_from_wr below, for more details on how ref_wr_duration
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* is obtained). In practice, the slower/faster the device at hand
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* is, the more/less it takes to load applications with respect to the
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* reference device. Accordingly, the longer/shorter BFQ grants
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* weight raising to interactive applications.
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*
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* BFQ uses two different reference pairs (ref_rate, ref_wr_duration),
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* depending on whether the device is rotational or non-rotational.
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*
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* In the following definitions, ref_rate[0] and ref_wr_duration[0]
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* are the reference values for a rotational device, whereas
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* ref_rate[1] and ref_wr_duration[1] are the reference values for a
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* non-rotational device. The reference rates are not the actual peak
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* rates of the devices used as a reference, but slightly lower
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* values. The reason for using slightly lower values is that the
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* peak-rate estimator tends to yield slightly lower values than the
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* actual peak rate (it can yield the actual peak rate only if there
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* is only one process doing I/O, and the process does sequential
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* I/O).
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*
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* The reference peak rates are measured in sectors/usec, left-shifted
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* by BFQ_RATE_SHIFT.
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*/
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static int ref_rate[2] = {14000, 33000};
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/*
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* To improve readability, a conversion function is used to initialize
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* the following array, which entails that the array can be
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* initialized only in a function.
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*/
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static int ref_wr_duration[2];
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/*
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* BFQ uses the above-detailed, time-based weight-raising mechanism to
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* privilege interactive tasks. This mechanism is vulnerable to the
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* following false positives: I/O-bound applications that will go on
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* doing I/O for much longer than the duration of weight
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* raising. These applications have basically no benefit from being
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* weight-raised at the beginning of their I/O. On the opposite end,
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* while being weight-raised, these applications
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* a) unjustly steal throughput to applications that may actually need
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* low latency;
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* b) make BFQ uselessly perform device idling; device idling results
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* in loss of device throughput with most flash-based storage, and may
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* increase latencies when used purposelessly.
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*
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* BFQ tries to reduce these problems, by adopting the following
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* countermeasure. To introduce this countermeasure, we need first to
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* finish explaining how the duration of weight-raising for
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* interactive tasks is computed.
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*
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* For a bfq_queue deemed as interactive, the duration of weight
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* raising is dynamically adjusted, as a function of the estimated
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* peak rate of the device, so as to be equal to the time needed to
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* execute the 'largest' interactive task we benchmarked so far. By
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* largest task, we mean the task for which each involved process has
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* to do more I/O than for any of the other tasks we benchmarked. This
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* reference interactive task is the start-up of LibreOffice Writer,
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* and in this task each process/bfq_queue needs to have at most ~110K
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* sectors transferred.
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*
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* This last piece of information enables BFQ to reduce the actual
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* duration of weight-raising for at least one class of I/O-bound
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* applications: those doing sequential or quasi-sequential I/O. An
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* example is file copy. In fact, once started, the main I/O-bound
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* processes of these applications usually consume the above 110K
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* sectors in much less time than the processes of an application that
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* is starting, because these I/O-bound processes will greedily devote
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* almost all their CPU cycles only to their target,
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* throughput-friendly I/O operations. This is even more true if BFQ
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* happens to be underestimating the device peak rate, and thus
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* overestimating the duration of weight raising. But, according to
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* our measurements, once transferred 110K sectors, these processes
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* have no right to be weight-raised any longer.
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*
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* Basing on the last consideration, BFQ ends weight-raising for a
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* bfq_queue if the latter happens to have received an amount of
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* service at least equal to the following constant. The constant is
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* set to slightly more than 110K, to have a minimum safety margin.
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*
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* This early ending of weight-raising reduces the amount of time
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* during which interactive false positives cause the two problems
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* described at the beginning of these comments.
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*/
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static const unsigned long max_service_from_wr = 120000;
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#define RQ_BIC(rq) icq_to_bic((rq)->elv.priv[0])
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#define RQ_BFQQ(rq) ((rq)->elv.priv[1])
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struct bfq_queue *bic_to_bfqq(struct bfq_io_cq *bic, bool is_sync)
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{
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return bic->bfqq[is_sync];
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}
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static void bfq_put_stable_ref(struct bfq_queue *bfqq);
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void bic_set_bfqq(struct bfq_io_cq *bic, struct bfq_queue *bfqq, bool is_sync)
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{
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/*
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* If bfqq != NULL, then a non-stable queue merge between
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* bic->bfqq and bfqq is happening here. This causes troubles
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* in the following case: bic->bfqq has also been scheduled
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* for a possible stable merge with bic->stable_merge_bfqq,
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* and bic->stable_merge_bfqq == bfqq happens to
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* hold. Troubles occur because bfqq may then undergo a split,
|
|
* thereby becoming eligible for a stable merge. Yet, if
|
|
* bic->stable_merge_bfqq points exactly to bfqq, then bfqq
|
|
* would be stably merged with itself. To avoid this anomaly,
|
|
* we cancel the stable merge if
|
|
* bic->stable_merge_bfqq == bfqq.
|
|
*/
|
|
bic->bfqq[is_sync] = bfqq;
|
|
|
|
if (bfqq && bic->stable_merge_bfqq == bfqq) {
|
|
/*
|
|
* Actually, these same instructions are executed also
|
|
* in bfq_setup_cooperator, in case of abort or actual
|
|
* execution of a stable merge. We could avoid
|
|
* repeating these instructions there too, but if we
|
|
* did so, we would nest even more complexity in this
|
|
* function.
|
|
*/
|
|
bfq_put_stable_ref(bic->stable_merge_bfqq);
|
|
|
|
bic->stable_merge_bfqq = NULL;
|
|
}
|
|
}
|
|
|
|
struct bfq_data *bic_to_bfqd(struct bfq_io_cq *bic)
|
|
{
|
|
return bic->icq.q->elevator->elevator_data;
|
|
}
|
|
|
|
/**
|
|
* icq_to_bic - convert iocontext queue structure to bfq_io_cq.
|
|
* @icq: the iocontext queue.
|
|
*/
|
|
static struct bfq_io_cq *icq_to_bic(struct io_cq *icq)
|
|
{
|
|
/* bic->icq is the first member, %NULL will convert to %NULL */
|
|
return container_of(icq, struct bfq_io_cq, icq);
|
|
}
|
|
|
|
/**
|
|
* bfq_bic_lookup - search into @ioc a bic associated to @bfqd.
|
|
* @bfqd: the lookup key.
|
|
* @ioc: the io_context of the process doing I/O.
|
|
* @q: the request queue.
|
|
*/
|
|
static struct bfq_io_cq *bfq_bic_lookup(struct bfq_data *bfqd,
|
|
struct io_context *ioc,
|
|
struct request_queue *q)
|
|
{
|
|
if (ioc) {
|
|
unsigned long flags;
|
|
struct bfq_io_cq *icq;
|
|
|
|
spin_lock_irqsave(&q->queue_lock, flags);
|
|
icq = icq_to_bic(ioc_lookup_icq(ioc, q));
|
|
spin_unlock_irqrestore(&q->queue_lock, flags);
|
|
|
|
return icq;
|
|
}
|
|
|
|
return NULL;
|
|
}
|
|
|
|
/*
|
|
* Scheduler run of queue, if there are requests pending and no one in the
|
|
* driver that will restart queueing.
|
|
*/
|
|
void bfq_schedule_dispatch(struct bfq_data *bfqd)
|
|
{
|
|
if (bfqd->queued != 0) {
|
|
bfq_log(bfqd, "schedule dispatch");
|
|
blk_mq_run_hw_queues(bfqd->queue, true);
|
|
}
|
|
}
|
|
|
|
#define bfq_class_idle(bfqq) ((bfqq)->ioprio_class == IOPRIO_CLASS_IDLE)
|
|
|
|
#define bfq_sample_valid(samples) ((samples) > 80)
|
|
|
|
/*
|
|
* Lifted from AS - choose which of rq1 and rq2 that is best served now.
|
|
* We choose the request that is closer to the head right now. Distance
|
|
* behind the head is penalized and only allowed to a certain extent.
|
|
*/
|
|
static struct request *bfq_choose_req(struct bfq_data *bfqd,
|
|
struct request *rq1,
|
|
struct request *rq2,
|
|
sector_t last)
|
|
{
|
|
sector_t s1, s2, d1 = 0, d2 = 0;
|
|
unsigned long back_max;
|
|
#define BFQ_RQ1_WRAP 0x01 /* request 1 wraps */
|
|
#define BFQ_RQ2_WRAP 0x02 /* request 2 wraps */
|
|
unsigned int wrap = 0; /* bit mask: requests behind the disk head? */
|
|
|
|
if (!rq1 || rq1 == rq2)
|
|
return rq2;
|
|
if (!rq2)
|
|
return rq1;
|
|
|
|
if (rq_is_sync(rq1) && !rq_is_sync(rq2))
|
|
return rq1;
|
|
else if (rq_is_sync(rq2) && !rq_is_sync(rq1))
|
|
return rq2;
|
|
if ((rq1->cmd_flags & REQ_META) && !(rq2->cmd_flags & REQ_META))
|
|
return rq1;
|
|
else if ((rq2->cmd_flags & REQ_META) && !(rq1->cmd_flags & REQ_META))
|
|
return rq2;
|
|
|
|
s1 = blk_rq_pos(rq1);
|
|
s2 = blk_rq_pos(rq2);
|
|
|
|
/*
|
|
* By definition, 1KiB is 2 sectors.
|
|
*/
|
|
back_max = bfqd->bfq_back_max * 2;
|
|
|
|
/*
|
|
* Strict one way elevator _except_ in the case where we allow
|
|
* short backward seeks which are biased as twice the cost of a
|
|
* similar forward seek.
|
|
*/
|
|
if (s1 >= last)
|
|
d1 = s1 - last;
|
|
else if (s1 + back_max >= last)
|
|
d1 = (last - s1) * bfqd->bfq_back_penalty;
|
|
else
|
|
wrap |= BFQ_RQ1_WRAP;
|
|
|
|
if (s2 >= last)
|
|
d2 = s2 - last;
|
|
else if (s2 + back_max >= last)
|
|
d2 = (last - s2) * bfqd->bfq_back_penalty;
|
|
else
|
|
wrap |= BFQ_RQ2_WRAP;
|
|
|
|
/* Found required data */
|
|
|
|
/*
|
|
* By doing switch() on the bit mask "wrap" we avoid having to
|
|
* check two variables for all permutations: --> faster!
|
|
*/
|
|
switch (wrap) {
|
|
case 0: /* common case for CFQ: rq1 and rq2 not wrapped */
|
|
if (d1 < d2)
|
|
return rq1;
|
|
else if (d2 < d1)
|
|
return rq2;
|
|
|
|
if (s1 >= s2)
|
|
return rq1;
|
|
else
|
|
return rq2;
|
|
|
|
case BFQ_RQ2_WRAP:
|
|
return rq1;
|
|
case BFQ_RQ1_WRAP:
|
|
return rq2;
|
|
case BFQ_RQ1_WRAP|BFQ_RQ2_WRAP: /* both rqs wrapped */
|
|
default:
|
|
/*
|
|
* Since both rqs are wrapped,
|
|
* start with the one that's further behind head
|
|
* (--> only *one* back seek required),
|
|
* since back seek takes more time than forward.
|
|
*/
|
|
if (s1 <= s2)
|
|
return rq1;
|
|
else
|
|
return rq2;
|
|
}
|
|
}
|
|
|
|
/*
|
|
* Async I/O can easily starve sync I/O (both sync reads and sync
|
|
* writes), by consuming all tags. Similarly, storms of sync writes,
|
|
* such as those that sync(2) may trigger, can starve sync reads.
|
|
* Limit depths of async I/O and sync writes so as to counter both
|
|
* problems.
|
|
*/
|
|
static void bfq_limit_depth(unsigned int op, struct blk_mq_alloc_data *data)
|
|
{
|
|
struct bfq_data *bfqd = data->q->elevator->elevator_data;
|
|
|
|
if (op_is_sync(op) && !op_is_write(op))
|
|
return;
|
|
|
|
data->shallow_depth =
|
|
bfqd->word_depths[!!bfqd->wr_busy_queues][op_is_sync(op)];
|
|
|
|
bfq_log(bfqd, "[%s] wr_busy %d sync %d depth %u",
|
|
__func__, bfqd->wr_busy_queues, op_is_sync(op),
|
|
data->shallow_depth);
|
|
}
|
|
|
|
static struct bfq_queue *
|
|
bfq_rq_pos_tree_lookup(struct bfq_data *bfqd, struct rb_root *root,
|
|
sector_t sector, struct rb_node **ret_parent,
|
|
struct rb_node ***rb_link)
|
|
{
|
|
struct rb_node **p, *parent;
|
|
struct bfq_queue *bfqq = NULL;
|
|
|
|
parent = NULL;
|
|
p = &root->rb_node;
|
|
while (*p) {
|
|
struct rb_node **n;
|
|
|
|
parent = *p;
|
|
bfqq = rb_entry(parent, struct bfq_queue, pos_node);
|
|
|
|
/*
|
|
* Sort strictly based on sector. Smallest to the left,
|
|
* largest to the right.
|
|
*/
|
|
if (sector > blk_rq_pos(bfqq->next_rq))
|
|
n = &(*p)->rb_right;
|
|
else if (sector < blk_rq_pos(bfqq->next_rq))
|
|
n = &(*p)->rb_left;
|
|
else
|
|
break;
|
|
p = n;
|
|
bfqq = NULL;
|
|
}
|
|
|
|
*ret_parent = parent;
|
|
if (rb_link)
|
|
*rb_link = p;
|
|
|
|
bfq_log(bfqd, "rq_pos_tree_lookup %llu: returning %d",
|
|
(unsigned long long)sector,
|
|
bfqq ? bfqq->pid : 0);
|
|
|
|
return bfqq;
|
|
}
|
|
|
|
static bool bfq_too_late_for_merging(struct bfq_queue *bfqq)
|
|
{
|
|
return bfqq->service_from_backlogged > 0 &&
|
|
time_is_before_jiffies(bfqq->first_IO_time +
|
|
bfq_merge_time_limit);
|
|
}
|
|
|
|
/*
|
|
* The following function is not marked as __cold because it is
|
|
* actually cold, but for the same performance goal described in the
|
|
* comments on the likely() at the beginning of
|
|
* bfq_setup_cooperator(). Unexpectedly, to reach an even lower
|
|
* execution time for the case where this function is not invoked, we
|
|
* had to add an unlikely() in each involved if().
|
|
*/
|
|
void __cold
|
|
bfq_pos_tree_add_move(struct bfq_data *bfqd, struct bfq_queue *bfqq)
|
|
{
|
|
struct rb_node **p, *parent;
|
|
struct bfq_queue *__bfqq;
|
|
|
|
if (bfqq->pos_root) {
|
|
rb_erase(&bfqq->pos_node, bfqq->pos_root);
|
|
bfqq->pos_root = NULL;
|
|
}
|
|
|
|
/* oom_bfqq does not participate in queue merging */
|
|
if (bfqq == &bfqd->oom_bfqq)
|
|
return;
|
|
|
|
/*
|
|
* bfqq cannot be merged any longer (see comments in
|
|
* bfq_setup_cooperator): no point in adding bfqq into the
|
|
* position tree.
|
|
*/
|
|
if (bfq_too_late_for_merging(bfqq))
|
|
return;
|
|
|
|
if (bfq_class_idle(bfqq))
|
|
return;
|
|
if (!bfqq->next_rq)
|
|
return;
|
|
|
|
bfqq->pos_root = &bfq_bfqq_to_bfqg(bfqq)->rq_pos_tree;
|
|
__bfqq = bfq_rq_pos_tree_lookup(bfqd, bfqq->pos_root,
|
|
blk_rq_pos(bfqq->next_rq), &parent, &p);
|
|
if (!__bfqq) {
|
|
rb_link_node(&bfqq->pos_node, parent, p);
|
|
rb_insert_color(&bfqq->pos_node, bfqq->pos_root);
|
|
} else
|
|
bfqq->pos_root = NULL;
|
|
}
|
|
|
|
/*
|
|
* The following function returns false either if every active queue
|
|
* must receive the same share of the throughput (symmetric scenario),
|
|
* or, as a special case, if bfqq must receive a share of the
|
|
* throughput lower than or equal to the share that every other active
|
|
* queue must receive. If bfqq does sync I/O, then these are the only
|
|
* two cases where bfqq happens to be guaranteed its share of the
|
|
* throughput even if I/O dispatching is not plugged when bfqq remains
|
|
* temporarily empty (for more details, see the comments in the
|
|
* function bfq_better_to_idle()). For this reason, the return value
|
|
* of this function is used to check whether I/O-dispatch plugging can
|
|
* be avoided.
|
|
*
|
|
* The above first case (symmetric scenario) occurs when:
|
|
* 1) all active queues have the same weight,
|
|
* 2) all active queues belong to the same I/O-priority class,
|
|
* 3) all active groups at the same level in the groups tree have the same
|
|
* weight,
|
|
* 4) all active groups at the same level in the groups tree have the same
|
|
* number of children.
|
|
*
|
|
* Unfortunately, keeping the necessary state for evaluating exactly
|
|
* the last two symmetry sub-conditions above would be quite complex
|
|
* and time consuming. Therefore this function evaluates, instead,
|
|
* only the following stronger three sub-conditions, for which it is
|
|
* much easier to maintain the needed state:
|
|
* 1) all active queues have the same weight,
|
|
* 2) all active queues belong to the same I/O-priority class,
|
|
* 3) there are no active groups.
|
|
* In particular, the last condition is always true if hierarchical
|
|
* support or the cgroups interface are not enabled, thus no state
|
|
* needs to be maintained in this case.
|
|
*/
|
|
static bool bfq_asymmetric_scenario(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
bool smallest_weight = bfqq &&
|
|
bfqq->weight_counter &&
|
|
bfqq->weight_counter ==
|
|
container_of(
|
|
rb_first_cached(&bfqd->queue_weights_tree),
|
|
struct bfq_weight_counter,
|
|
weights_node);
|
|
|
|
/*
|
|
* For queue weights to differ, queue_weights_tree must contain
|
|
* at least two nodes.
|
|
*/
|
|
bool varied_queue_weights = !smallest_weight &&
|
|
!RB_EMPTY_ROOT(&bfqd->queue_weights_tree.rb_root) &&
|
|
(bfqd->queue_weights_tree.rb_root.rb_node->rb_left ||
|
|
bfqd->queue_weights_tree.rb_root.rb_node->rb_right);
|
|
|
|
bool multiple_classes_busy =
|
|
(bfqd->busy_queues[0] && bfqd->busy_queues[1]) ||
|
|
(bfqd->busy_queues[0] && bfqd->busy_queues[2]) ||
|
|
(bfqd->busy_queues[1] && bfqd->busy_queues[2]);
|
|
|
|
return varied_queue_weights || multiple_classes_busy
|
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED
|
|
|| bfqd->num_groups_with_pending_reqs > 0
|
|
#endif
|
|
;
|
|
}
|
|
|
|
/*
|
|
* If the weight-counter tree passed as input contains no counter for
|
|
* the weight of the input queue, then add that counter; otherwise just
|
|
* increment the existing counter.
|
|
*
|
|
* Note that weight-counter trees contain few nodes in mostly symmetric
|
|
* scenarios. For example, if all queues have the same weight, then the
|
|
* weight-counter tree for the queues may contain at most one node.
|
|
* This holds even if low_latency is on, because weight-raised queues
|
|
* are not inserted in the tree.
|
|
* In most scenarios, the rate at which nodes are created/destroyed
|
|
* should be low too.
|
|
*/
|
|
void bfq_weights_tree_add(struct bfq_data *bfqd, struct bfq_queue *bfqq,
|
|
struct rb_root_cached *root)
|
|
{
|
|
struct bfq_entity *entity = &bfqq->entity;
|
|
struct rb_node **new = &(root->rb_root.rb_node), *parent = NULL;
|
|
bool leftmost = true;
|
|
|
|
/*
|
|
* Do not insert if the queue is already associated with a
|
|
* counter, which happens if:
|
|
* 1) a request arrival has caused the queue to become both
|
|
* non-weight-raised, and hence change its weight, and
|
|
* backlogged; in this respect, each of the two events
|
|
* causes an invocation of this function,
|
|
* 2) this is the invocation of this function caused by the
|
|
* second event. This second invocation is actually useless,
|
|
* and we handle this fact by exiting immediately. More
|
|
* efficient or clearer solutions might possibly be adopted.
|
|
*/
|
|
if (bfqq->weight_counter)
|
|
return;
|
|
|
|
while (*new) {
|
|
struct bfq_weight_counter *__counter = container_of(*new,
|
|
struct bfq_weight_counter,
|
|
weights_node);
|
|
parent = *new;
|
|
|
|
if (entity->weight == __counter->weight) {
|
|
bfqq->weight_counter = __counter;
|
|
goto inc_counter;
|
|
}
|
|
if (entity->weight < __counter->weight)
|
|
new = &((*new)->rb_left);
|
|
else {
|
|
new = &((*new)->rb_right);
|
|
leftmost = false;
|
|
}
|
|
}
|
|
|
|
bfqq->weight_counter = kzalloc(sizeof(struct bfq_weight_counter),
|
|
GFP_ATOMIC);
|
|
|
|
/*
|
|
* In the unlucky event of an allocation failure, we just
|
|
* exit. This will cause the weight of queue to not be
|
|
* considered in bfq_asymmetric_scenario, which, in its turn,
|
|
* causes the scenario to be deemed wrongly symmetric in case
|
|
* bfqq's weight would have been the only weight making the
|
|
* scenario asymmetric. On the bright side, no unbalance will
|
|
* however occur when bfqq becomes inactive again (the
|
|
* invocation of this function is triggered by an activation
|
|
* of queue). In fact, bfq_weights_tree_remove does nothing
|
|
* if !bfqq->weight_counter.
|
|
*/
|
|
if (unlikely(!bfqq->weight_counter))
|
|
return;
|
|
|
|
bfqq->weight_counter->weight = entity->weight;
|
|
rb_link_node(&bfqq->weight_counter->weights_node, parent, new);
|
|
rb_insert_color_cached(&bfqq->weight_counter->weights_node, root,
|
|
leftmost);
|
|
|
|
inc_counter:
|
|
bfqq->weight_counter->num_active++;
|
|
bfqq->ref++;
|
|
}
|
|
|
|
/*
|
|
* Decrement the weight counter associated with the queue, and, if the
|
|
* counter reaches 0, remove the counter from the tree.
|
|
* See the comments to the function bfq_weights_tree_add() for considerations
|
|
* about overhead.
|
|
*/
|
|
void __bfq_weights_tree_remove(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq,
|
|
struct rb_root_cached *root)
|
|
{
|
|
if (!bfqq->weight_counter)
|
|
return;
|
|
|
|
bfqq->weight_counter->num_active--;
|
|
if (bfqq->weight_counter->num_active > 0)
|
|
goto reset_entity_pointer;
|
|
|
|
rb_erase_cached(&bfqq->weight_counter->weights_node, root);
|
|
kfree(bfqq->weight_counter);
|
|
|
|
reset_entity_pointer:
|
|
bfqq->weight_counter = NULL;
|
|
bfq_put_queue(bfqq);
|
|
}
|
|
|
|
/*
|
|
* Invoke __bfq_weights_tree_remove on bfqq and decrement the number
|
|
* of active groups for each queue's inactive parent entity.
|
|
*/
|
|
void bfq_weights_tree_remove(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
struct bfq_entity *entity = bfqq->entity.parent;
|
|
|
|
for_each_entity(entity) {
|
|
struct bfq_sched_data *sd = entity->my_sched_data;
|
|
|
|
if (sd->next_in_service || sd->in_service_entity) {
|
|
/*
|
|
* entity is still active, because either
|
|
* next_in_service or in_service_entity is not
|
|
* NULL (see the comments on the definition of
|
|
* next_in_service for details on why
|
|
* in_service_entity must be checked too).
|
|
*
|
|
* As a consequence, its parent entities are
|
|
* active as well, and thus this loop must
|
|
* stop here.
|
|
*/
|
|
break;
|
|
}
|
|
|
|
/*
|
|
* The decrement of num_groups_with_pending_reqs is
|
|
* not performed immediately upon the deactivation of
|
|
* entity, but it is delayed to when it also happens
|
|
* that the first leaf descendant bfqq of entity gets
|
|
* all its pending requests completed. The following
|
|
* instructions perform this delayed decrement, if
|
|
* needed. See the comments on
|
|
* num_groups_with_pending_reqs for details.
|
|
*/
|
|
if (entity->in_groups_with_pending_reqs) {
|
|
entity->in_groups_with_pending_reqs = false;
|
|
bfqd->num_groups_with_pending_reqs--;
|
|
}
|
|
}
|
|
|
|
/*
|
|
* Next function is invoked last, because it causes bfqq to be
|
|
* freed if the following holds: bfqq is not in service and
|
|
* has no dispatched request. DO NOT use bfqq after the next
|
|
* function invocation.
|
|
*/
|
|
__bfq_weights_tree_remove(bfqd, bfqq,
|
|
&bfqd->queue_weights_tree);
|
|
}
|
|
|
|
/*
|
|
* Return expired entry, or NULL to just start from scratch in rbtree.
|
|
*/
|
|
static struct request *bfq_check_fifo(struct bfq_queue *bfqq,
|
|
struct request *last)
|
|
{
|
|
struct request *rq;
|
|
|
|
if (bfq_bfqq_fifo_expire(bfqq))
|
|
return NULL;
|
|
|
|
bfq_mark_bfqq_fifo_expire(bfqq);
|
|
|
|
rq = rq_entry_fifo(bfqq->fifo.next);
|
|
|
|
if (rq == last || ktime_get_ns() < rq->fifo_time)
|
|
return NULL;
|
|
|
|
bfq_log_bfqq(bfqq->bfqd, bfqq, "check_fifo: returned %p", rq);
|
|
return rq;
|
|
}
|
|
|
|
static struct request *bfq_find_next_rq(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq,
|
|
struct request *last)
|
|
{
|
|
struct rb_node *rbnext = rb_next(&last->rb_node);
|
|
struct rb_node *rbprev = rb_prev(&last->rb_node);
|
|
struct request *next, *prev = NULL;
|
|
|
|
/* Follow expired path, else get first next available. */
|
|
next = bfq_check_fifo(bfqq, last);
|
|
if (next)
|
|
return next;
|
|
|
|
if (rbprev)
|
|
prev = rb_entry_rq(rbprev);
|
|
|
|
if (rbnext)
|
|
next = rb_entry_rq(rbnext);
|
|
else {
|
|
rbnext = rb_first(&bfqq->sort_list);
|
|
if (rbnext && rbnext != &last->rb_node)
|
|
next = rb_entry_rq(rbnext);
|
|
}
|
|
|
|
return bfq_choose_req(bfqd, next, prev, blk_rq_pos(last));
|
|
}
|
|
|
|
/* see the definition of bfq_async_charge_factor for details */
|
|
static unsigned long bfq_serv_to_charge(struct request *rq,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
if (bfq_bfqq_sync(bfqq) || bfqq->wr_coeff > 1 ||
|
|
bfq_asymmetric_scenario(bfqq->bfqd, bfqq))
|
|
return blk_rq_sectors(rq);
|
|
|
|
return blk_rq_sectors(rq) * bfq_async_charge_factor;
|
|
}
|
|
|
|
/**
|
|
* bfq_updated_next_req - update the queue after a new next_rq selection.
|
|
* @bfqd: the device data the queue belongs to.
|
|
* @bfqq: the queue to update.
|
|
*
|
|
* If the first request of a queue changes we make sure that the queue
|
|
* has enough budget to serve at least its first request (if the
|
|
* request has grown). We do this because if the queue has not enough
|
|
* budget for its first request, it has to go through two dispatch
|
|
* rounds to actually get it dispatched.
|
|
*/
|
|
static void bfq_updated_next_req(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
struct bfq_entity *entity = &bfqq->entity;
|
|
struct request *next_rq = bfqq->next_rq;
|
|
unsigned long new_budget;
|
|
|
|
if (!next_rq)
|
|
return;
|
|
|
|
if (bfqq == bfqd->in_service_queue)
|
|
/*
|
|
* In order not to break guarantees, budgets cannot be
|
|
* changed after an entity has been selected.
|
|
*/
|
|
return;
|
|
|
|
new_budget = max_t(unsigned long,
|
|
max_t(unsigned long, bfqq->max_budget,
|
|
bfq_serv_to_charge(next_rq, bfqq)),
|
|
entity->service);
|
|
if (entity->budget != new_budget) {
|
|
entity->budget = new_budget;
|
|
bfq_log_bfqq(bfqd, bfqq, "updated next rq: new budget %lu",
|
|
new_budget);
|
|
bfq_requeue_bfqq(bfqd, bfqq, false);
|
|
}
|
|
}
|
|
|
|
static unsigned int bfq_wr_duration(struct bfq_data *bfqd)
|
|
{
|
|
u64 dur;
|
|
|
|
if (bfqd->bfq_wr_max_time > 0)
|
|
return bfqd->bfq_wr_max_time;
|
|
|
|
dur = bfqd->rate_dur_prod;
|
|
do_div(dur, bfqd->peak_rate);
|
|
|
|
/*
|
|
* Limit duration between 3 and 25 seconds. The upper limit
|
|
* has been conservatively set after the following worst case:
|
|
* on a QEMU/KVM virtual machine
|
|
* - running in a slow PC
|
|
* - with a virtual disk stacked on a slow low-end 5400rpm HDD
|
|
* - serving a heavy I/O workload, such as the sequential reading
|
|
* of several files
|
|
* mplayer took 23 seconds to start, if constantly weight-raised.
|
|
*
|
|
* As for higher values than that accommodating the above bad
|
|
* scenario, tests show that higher values would often yield
|
|
* the opposite of the desired result, i.e., would worsen
|
|
* responsiveness by allowing non-interactive applications to
|
|
* preserve weight raising for too long.
|
|
*
|
|
* On the other end, lower values than 3 seconds make it
|
|
* difficult for most interactive tasks to complete their jobs
|
|
* before weight-raising finishes.
|
|
*/
|
|
return clamp_val(dur, msecs_to_jiffies(3000), msecs_to_jiffies(25000));
|
|
}
|
|
|
|
/* switch back from soft real-time to interactive weight raising */
|
|
static void switch_back_to_interactive_wr(struct bfq_queue *bfqq,
|
|
struct bfq_data *bfqd)
|
|
{
|
|
bfqq->wr_coeff = bfqd->bfq_wr_coeff;
|
|
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd);
|
|
bfqq->last_wr_start_finish = bfqq->wr_start_at_switch_to_srt;
|
|
}
|
|
|
|
static void
|
|
bfq_bfqq_resume_state(struct bfq_queue *bfqq, struct bfq_data *bfqd,
|
|
struct bfq_io_cq *bic, bool bfq_already_existing)
|
|
{
|
|
unsigned int old_wr_coeff = 1;
|
|
bool busy = bfq_already_existing && bfq_bfqq_busy(bfqq);
|
|
|
|
if (bic->saved_has_short_ttime)
|
|
bfq_mark_bfqq_has_short_ttime(bfqq);
|
|
else
|
|
bfq_clear_bfqq_has_short_ttime(bfqq);
|
|
|
|
if (bic->saved_IO_bound)
|
|
bfq_mark_bfqq_IO_bound(bfqq);
|
|
else
|
|
bfq_clear_bfqq_IO_bound(bfqq);
|
|
|
|
bfqq->last_serv_time_ns = bic->saved_last_serv_time_ns;
|
|
bfqq->inject_limit = bic->saved_inject_limit;
|
|
bfqq->decrease_time_jif = bic->saved_decrease_time_jif;
|
|
|
|
bfqq->entity.new_weight = bic->saved_weight;
|
|
bfqq->ttime = bic->saved_ttime;
|
|
bfqq->io_start_time = bic->saved_io_start_time;
|
|
bfqq->tot_idle_time = bic->saved_tot_idle_time;
|
|
/*
|
|
* Restore weight coefficient only if low_latency is on
|
|
*/
|
|
if (bfqd->low_latency) {
|
|
old_wr_coeff = bfqq->wr_coeff;
|
|
bfqq->wr_coeff = bic->saved_wr_coeff;
|
|
}
|
|
bfqq->service_from_wr = bic->saved_service_from_wr;
|
|
bfqq->wr_start_at_switch_to_srt = bic->saved_wr_start_at_switch_to_srt;
|
|
bfqq->last_wr_start_finish = bic->saved_last_wr_start_finish;
|
|
bfqq->wr_cur_max_time = bic->saved_wr_cur_max_time;
|
|
|
|
if (bfqq->wr_coeff > 1 && (bfq_bfqq_in_large_burst(bfqq) ||
|
|
time_is_before_jiffies(bfqq->last_wr_start_finish +
|
|
bfqq->wr_cur_max_time))) {
|
|
if (bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time &&
|
|
!bfq_bfqq_in_large_burst(bfqq) &&
|
|
time_is_after_eq_jiffies(bfqq->wr_start_at_switch_to_srt +
|
|
bfq_wr_duration(bfqd))) {
|
|
switch_back_to_interactive_wr(bfqq, bfqd);
|
|
} else {
|
|
bfqq->wr_coeff = 1;
|
|
bfq_log_bfqq(bfqq->bfqd, bfqq,
|
|
"resume state: switching off wr");
|
|
}
|
|
}
|
|
|
|
/* make sure weight will be updated, however we got here */
|
|
bfqq->entity.prio_changed = 1;
|
|
|
|
if (likely(!busy))
|
|
return;
|
|
|
|
if (old_wr_coeff == 1 && bfqq->wr_coeff > 1)
|
|
bfqd->wr_busy_queues++;
|
|
else if (old_wr_coeff > 1 && bfqq->wr_coeff == 1)
|
|
bfqd->wr_busy_queues--;
|
|
}
|
|
|
|
static int bfqq_process_refs(struct bfq_queue *bfqq)
|
|
{
|
|
return bfqq->ref - bfqq->allocated - bfqq->entity.on_st_or_in_serv -
|
|
(bfqq->weight_counter != NULL) - bfqq->stable_ref;
|
|
}
|
|
|
|
/* Empty burst list and add just bfqq (see comments on bfq_handle_burst) */
|
|
static void bfq_reset_burst_list(struct bfq_data *bfqd, struct bfq_queue *bfqq)
|
|
{
|
|
struct bfq_queue *item;
|
|
struct hlist_node *n;
|
|
|
|
hlist_for_each_entry_safe(item, n, &bfqd->burst_list, burst_list_node)
|
|
hlist_del_init(&item->burst_list_node);
|
|
|
|
/*
|
|
* Start the creation of a new burst list only if there is no
|
|
* active queue. See comments on the conditional invocation of
|
|
* bfq_handle_burst().
|
|
*/
|
|
if (bfq_tot_busy_queues(bfqd) == 0) {
|
|
hlist_add_head(&bfqq->burst_list_node, &bfqd->burst_list);
|
|
bfqd->burst_size = 1;
|
|
} else
|
|
bfqd->burst_size = 0;
|
|
|
|
bfqd->burst_parent_entity = bfqq->entity.parent;
|
|
}
|
|
|
|
/* Add bfqq to the list of queues in current burst (see bfq_handle_burst) */
|
|
static void bfq_add_to_burst(struct bfq_data *bfqd, struct bfq_queue *bfqq)
|
|
{
|
|
/* Increment burst size to take into account also bfqq */
|
|
bfqd->burst_size++;
|
|
|
|
if (bfqd->burst_size == bfqd->bfq_large_burst_thresh) {
|
|
struct bfq_queue *pos, *bfqq_item;
|
|
struct hlist_node *n;
|
|
|
|
/*
|
|
* Enough queues have been activated shortly after each
|
|
* other to consider this burst as large.
|
|
*/
|
|
bfqd->large_burst = true;
|
|
|
|
/*
|
|
* We can now mark all queues in the burst list as
|
|
* belonging to a large burst.
|
|
*/
|
|
hlist_for_each_entry(bfqq_item, &bfqd->burst_list,
|
|
burst_list_node)
|
|
bfq_mark_bfqq_in_large_burst(bfqq_item);
|
|
bfq_mark_bfqq_in_large_burst(bfqq);
|
|
|
|
/*
|
|
* From now on, and until the current burst finishes, any
|
|
* new queue being activated shortly after the last queue
|
|
* was inserted in the burst can be immediately marked as
|
|
* belonging to a large burst. So the burst list is not
|
|
* needed any more. Remove it.
|
|
*/
|
|
hlist_for_each_entry_safe(pos, n, &bfqd->burst_list,
|
|
burst_list_node)
|
|
hlist_del_init(&pos->burst_list_node);
|
|
} else /*
|
|
* Burst not yet large: add bfqq to the burst list. Do
|
|
* not increment the ref counter for bfqq, because bfqq
|
|
* is removed from the burst list before freeing bfqq
|
|
* in put_queue.
|
|
*/
|
|
hlist_add_head(&bfqq->burst_list_node, &bfqd->burst_list);
|
|
}
|
|
|
|
/*
|
|
* If many queues belonging to the same group happen to be created
|
|
* shortly after each other, then the processes associated with these
|
|
* queues have typically a common goal. In particular, bursts of queue
|
|
* creations are usually caused by services or applications that spawn
|
|
* many parallel threads/processes. Examples are systemd during boot,
|
|
* or git grep. To help these processes get their job done as soon as
|
|
* possible, it is usually better to not grant either weight-raising
|
|
* or device idling to their queues, unless these queues must be
|
|
* protected from the I/O flowing through other active queues.
|
|
*
|
|
* In this comment we describe, firstly, the reasons why this fact
|
|
* holds, and, secondly, the next function, which implements the main
|
|
* steps needed to properly mark these queues so that they can then be
|
|
* treated in a different way.
|
|
*
|
|
* The above services or applications benefit mostly from a high
|
|
* throughput: the quicker the requests of the activated queues are
|
|
* cumulatively served, the sooner the target job of these queues gets
|
|
* completed. As a consequence, weight-raising any of these queues,
|
|
* which also implies idling the device for it, is almost always
|
|
* counterproductive, unless there are other active queues to isolate
|
|
* these new queues from. If there no other active queues, then
|
|
* weight-raising these new queues just lowers throughput in most
|
|
* cases.
|
|
*
|
|
* On the other hand, a burst of queue creations may be caused also by
|
|
* the start of an application that does not consist of a lot of
|
|
* parallel I/O-bound threads. In fact, with a complex application,
|
|
* several short processes may need to be executed to start-up the
|
|
* application. In this respect, to start an application as quickly as
|
|
* possible, the best thing to do is in any case to privilege the I/O
|
|
* related to the application with respect to all other
|
|
* I/O. Therefore, the best strategy to start as quickly as possible
|
|
* an application that causes a burst of queue creations is to
|
|
* weight-raise all the queues created during the burst. This is the
|
|
* exact opposite of the best strategy for the other type of bursts.
|
|
*
|
|
* In the end, to take the best action for each of the two cases, the
|
|
* two types of bursts need to be distinguished. Fortunately, this
|
|
* seems relatively easy, by looking at the sizes of the bursts. In
|
|
* particular, we found a threshold such that only bursts with a
|
|
* larger size than that threshold are apparently caused by
|
|
* services or commands such as systemd or git grep. For brevity,
|
|
* hereafter we call just 'large' these bursts. BFQ *does not*
|
|
* weight-raise queues whose creation occurs in a large burst. In
|
|
* addition, for each of these queues BFQ performs or does not perform
|
|
* idling depending on which choice boosts the throughput more. The
|
|
* exact choice depends on the device and request pattern at
|
|
* hand.
|
|
*
|
|
* Unfortunately, false positives may occur while an interactive task
|
|
* is starting (e.g., an application is being started). The
|
|
* consequence is that the queues associated with the task do not
|
|
* enjoy weight raising as expected. Fortunately these false positives
|
|
* are very rare. They typically occur if some service happens to
|
|
* start doing I/O exactly when the interactive task starts.
|
|
*
|
|
* Turning back to the next function, it is invoked only if there are
|
|
* no active queues (apart from active queues that would belong to the
|
|
* same, possible burst bfqq would belong to), and it implements all
|
|
* the steps needed to detect the occurrence of a large burst and to
|
|
* properly mark all the queues belonging to it (so that they can then
|
|
* be treated in a different way). This goal is achieved by
|
|
* maintaining a "burst list" that holds, temporarily, the queues that
|
|
* belong to the burst in progress. The list is then used to mark
|
|
* these queues as belonging to a large burst if the burst does become
|
|
* large. The main steps are the following.
|
|
*
|
|
* . when the very first queue is created, the queue is inserted into the
|
|
* list (as it could be the first queue in a possible burst)
|
|
*
|
|
* . if the current burst has not yet become large, and a queue Q that does
|
|
* not yet belong to the burst is activated shortly after the last time
|
|
* at which a new queue entered the burst list, then the function appends
|
|
* Q to the burst list
|
|
*
|
|
* . if, as a consequence of the previous step, the burst size reaches
|
|
* the large-burst threshold, then
|
|
*
|
|
* . all the queues in the burst list are marked as belonging to a
|
|
* large burst
|
|
*
|
|
* . the burst list is deleted; in fact, the burst list already served
|
|
* its purpose (keeping temporarily track of the queues in a burst,
|
|
* so as to be able to mark them as belonging to a large burst in the
|
|
* previous sub-step), and now is not needed any more
|
|
*
|
|
* . the device enters a large-burst mode
|
|
*
|
|
* . if a queue Q that does not belong to the burst is created while
|
|
* the device is in large-burst mode and shortly after the last time
|
|
* at which a queue either entered the burst list or was marked as
|
|
* belonging to the current large burst, then Q is immediately marked
|
|
* as belonging to a large burst.
|
|
*
|
|
* . if a queue Q that does not belong to the burst is created a while
|
|
* later, i.e., not shortly after, than the last time at which a queue
|
|
* either entered the burst list or was marked as belonging to the
|
|
* current large burst, then the current burst is deemed as finished and:
|
|
*
|
|
* . the large-burst mode is reset if set
|
|
*
|
|
* . the burst list is emptied
|
|
*
|
|
* . Q is inserted in the burst list, as Q may be the first queue
|
|
* in a possible new burst (then the burst list contains just Q
|
|
* after this step).
|
|
*/
|
|
static void bfq_handle_burst(struct bfq_data *bfqd, struct bfq_queue *bfqq)
|
|
{
|
|
/*
|
|
* If bfqq is already in the burst list or is part of a large
|
|
* burst, or finally has just been split, then there is
|
|
* nothing else to do.
|
|
*/
|
|
if (!hlist_unhashed(&bfqq->burst_list_node) ||
|
|
bfq_bfqq_in_large_burst(bfqq) ||
|
|
time_is_after_eq_jiffies(bfqq->split_time +
|
|
msecs_to_jiffies(10)))
|
|
return;
|
|
|
|
/*
|
|
* If bfqq's creation happens late enough, or bfqq belongs to
|
|
* a different group than the burst group, then the current
|
|
* burst is finished, and related data structures must be
|
|
* reset.
|
|
*
|
|
* In this respect, consider the special case where bfqq is
|
|
* the very first queue created after BFQ is selected for this
|
|
* device. In this case, last_ins_in_burst and
|
|
* burst_parent_entity are not yet significant when we get
|
|
* here. But it is easy to verify that, whether or not the
|
|
* following condition is true, bfqq will end up being
|
|
* inserted into the burst list. In particular the list will
|
|
* happen to contain only bfqq. And this is exactly what has
|
|
* to happen, as bfqq may be the first queue of the first
|
|
* burst.
|
|
*/
|
|
if (time_is_before_jiffies(bfqd->last_ins_in_burst +
|
|
bfqd->bfq_burst_interval) ||
|
|
bfqq->entity.parent != bfqd->burst_parent_entity) {
|
|
bfqd->large_burst = false;
|
|
bfq_reset_burst_list(bfqd, bfqq);
|
|
goto end;
|
|
}
|
|
|
|
/*
|
|
* If we get here, then bfqq is being activated shortly after the
|
|
* last queue. So, if the current burst is also large, we can mark
|
|
* bfqq as belonging to this large burst immediately.
|
|
*/
|
|
if (bfqd->large_burst) {
|
|
bfq_mark_bfqq_in_large_burst(bfqq);
|
|
goto end;
|
|
}
|
|
|
|
/*
|
|
* If we get here, then a large-burst state has not yet been
|
|
* reached, but bfqq is being activated shortly after the last
|
|
* queue. Then we add bfqq to the burst.
|
|
*/
|
|
bfq_add_to_burst(bfqd, bfqq);
|
|
end:
|
|
/*
|
|
* At this point, bfqq either has been added to the current
|
|
* burst or has caused the current burst to terminate and a
|
|
* possible new burst to start. In particular, in the second
|
|
* case, bfqq has become the first queue in the possible new
|
|
* burst. In both cases last_ins_in_burst needs to be moved
|
|
* forward.
|
|
*/
|
|
bfqd->last_ins_in_burst = jiffies;
|
|
}
|
|
|
|
static int bfq_bfqq_budget_left(struct bfq_queue *bfqq)
|
|
{
|
|
struct bfq_entity *entity = &bfqq->entity;
|
|
|
|
return entity->budget - entity->service;
|
|
}
|
|
|
|
/*
|
|
* If enough samples have been computed, return the current max budget
|
|
* stored in bfqd, which is dynamically updated according to the
|
|
* estimated disk peak rate; otherwise return the default max budget
|
|
*/
|
|
static int bfq_max_budget(struct bfq_data *bfqd)
|
|
{
|
|
if (bfqd->budgets_assigned < bfq_stats_min_budgets)
|
|
return bfq_default_max_budget;
|
|
else
|
|
return bfqd->bfq_max_budget;
|
|
}
|
|
|
|
/*
|
|
* Return min budget, which is a fraction of the current or default
|
|
* max budget (trying with 1/32)
|
|
*/
|
|
static int bfq_min_budget(struct bfq_data *bfqd)
|
|
{
|
|
if (bfqd->budgets_assigned < bfq_stats_min_budgets)
|
|
return bfq_default_max_budget / 32;
|
|
else
|
|
return bfqd->bfq_max_budget / 32;
|
|
}
|
|
|
|
/*
|
|
* The next function, invoked after the input queue bfqq switches from
|
|
* idle to busy, updates the budget of bfqq. The function also tells
|
|
* whether the in-service queue should be expired, by returning
|
|
* true. The purpose of expiring the in-service queue is to give bfqq
|
|
* the chance to possibly preempt the in-service queue, and the reason
|
|
* for preempting the in-service queue is to achieve one of the two
|
|
* goals below.
|
|
*
|
|
* 1. Guarantee to bfqq its reserved bandwidth even if bfqq has
|
|
* expired because it has remained idle. In particular, bfqq may have
|
|
* expired for one of the following two reasons:
|
|
*
|
|
* - BFQQE_NO_MORE_REQUESTS bfqq did not enjoy any device idling
|
|
* and did not make it to issue a new request before its last
|
|
* request was served;
|
|
*
|
|
* - BFQQE_TOO_IDLE bfqq did enjoy device idling, but did not issue
|
|
* a new request before the expiration of the idling-time.
|
|
*
|
|
* Even if bfqq has expired for one of the above reasons, the process
|
|
* associated with the queue may be however issuing requests greedily,
|
|
* and thus be sensitive to the bandwidth it receives (bfqq may have
|
|
* remained idle for other reasons: CPU high load, bfqq not enjoying
|
|
* idling, I/O throttling somewhere in the path from the process to
|
|
* the I/O scheduler, ...). But if, after every expiration for one of
|
|
* the above two reasons, bfqq has to wait for the service of at least
|
|
* one full budget of another queue before being served again, then
|
|
* bfqq is likely to get a much lower bandwidth or resource time than
|
|
* its reserved ones. To address this issue, two countermeasures need
|
|
* to be taken.
|
|
*
|
|
* First, the budget and the timestamps of bfqq need to be updated in
|
|
* a special way on bfqq reactivation: they need to be updated as if
|
|
* bfqq did not remain idle and did not expire. In fact, if they are
|
|
* computed as if bfqq expired and remained idle until reactivation,
|
|
* then the process associated with bfqq is treated as if, instead of
|
|
* being greedy, it stopped issuing requests when bfqq remained idle,
|
|
* and restarts issuing requests only on this reactivation. In other
|
|
* words, the scheduler does not help the process recover the "service
|
|
* hole" between bfqq expiration and reactivation. As a consequence,
|
|
* the process receives a lower bandwidth than its reserved one. In
|
|
* contrast, to recover this hole, the budget must be updated as if
|
|
* bfqq was not expired at all before this reactivation, i.e., it must
|
|
* be set to the value of the remaining budget when bfqq was
|
|
* expired. Along the same line, timestamps need to be assigned the
|
|
* value they had the last time bfqq was selected for service, i.e.,
|
|
* before last expiration. Thus timestamps need to be back-shifted
|
|
* with respect to their normal computation (see [1] for more details
|
|
* on this tricky aspect).
|
|
*
|
|
* Secondly, to allow the process to recover the hole, the in-service
|
|
* queue must be expired too, to give bfqq the chance to preempt it
|
|
* immediately. In fact, if bfqq has to wait for a full budget of the
|
|
* in-service queue to be completed, then it may become impossible to
|
|
* let the process recover the hole, even if the back-shifted
|
|
* timestamps of bfqq are lower than those of the in-service queue. If
|
|
* this happens for most or all of the holes, then the process may not
|
|
* receive its reserved bandwidth. In this respect, it is worth noting
|
|
* that, being the service of outstanding requests unpreemptible, a
|
|
* little fraction of the holes may however be unrecoverable, thereby
|
|
* causing a little loss of bandwidth.
|
|
*
|
|
* The last important point is detecting whether bfqq does need this
|
|
* bandwidth recovery. In this respect, the next function deems the
|
|
* process associated with bfqq greedy, and thus allows it to recover
|
|
* the hole, if: 1) the process is waiting for the arrival of a new
|
|
* request (which implies that bfqq expired for one of the above two
|
|
* reasons), and 2) such a request has arrived soon. The first
|
|
* condition is controlled through the flag non_blocking_wait_rq,
|
|
* while the second through the flag arrived_in_time. If both
|
|
* conditions hold, then the function computes the budget in the
|
|
* above-described special way, and signals that the in-service queue
|
|
* should be expired. Timestamp back-shifting is done later in
|
|
* __bfq_activate_entity.
|
|
*
|
|
* 2. Reduce latency. Even if timestamps are not backshifted to let
|
|
* the process associated with bfqq recover a service hole, bfqq may
|
|
* however happen to have, after being (re)activated, a lower finish
|
|
* timestamp than the in-service queue. That is, the next budget of
|
|
* bfqq may have to be completed before the one of the in-service
|
|
* queue. If this is the case, then preempting the in-service queue
|
|
* allows this goal to be achieved, apart from the unpreemptible,
|
|
* outstanding requests mentioned above.
|
|
*
|
|
* Unfortunately, regardless of which of the above two goals one wants
|
|
* to achieve, service trees need first to be updated to know whether
|
|
* the in-service queue must be preempted. To have service trees
|
|
* correctly updated, the in-service queue must be expired and
|
|
* rescheduled, and bfqq must be scheduled too. This is one of the
|
|
* most costly operations (in future versions, the scheduling
|
|
* mechanism may be re-designed in such a way to make it possible to
|
|
* know whether preemption is needed without needing to update service
|
|
* trees). In addition, queue preemptions almost always cause random
|
|
* I/O, which may in turn cause loss of throughput. Finally, there may
|
|
* even be no in-service queue when the next function is invoked (so,
|
|
* no queue to compare timestamps with). Because of these facts, the
|
|
* next function adopts the following simple scheme to avoid costly
|
|
* operations, too frequent preemptions and too many dependencies on
|
|
* the state of the scheduler: it requests the expiration of the
|
|
* in-service queue (unconditionally) only for queues that need to
|
|
* recover a hole. Then it delegates to other parts of the code the
|
|
* responsibility of handling the above case 2.
|
|
*/
|
|
static bool bfq_bfqq_update_budg_for_activation(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq,
|
|
bool arrived_in_time)
|
|
{
|
|
struct bfq_entity *entity = &bfqq->entity;
|
|
|
|
/*
|
|
* In the next compound condition, we check also whether there
|
|
* is some budget left, because otherwise there is no point in
|
|
* trying to go on serving bfqq with this same budget: bfqq
|
|
* would be expired immediately after being selected for
|
|
* service. This would only cause useless overhead.
|
|
*/
|
|
if (bfq_bfqq_non_blocking_wait_rq(bfqq) && arrived_in_time &&
|
|
bfq_bfqq_budget_left(bfqq) > 0) {
|
|
/*
|
|
* We do not clear the flag non_blocking_wait_rq here, as
|
|
* the latter is used in bfq_activate_bfqq to signal
|
|
* that timestamps need to be back-shifted (and is
|
|
* cleared right after).
|
|
*/
|
|
|
|
/*
|
|
* In next assignment we rely on that either
|
|
* entity->service or entity->budget are not updated
|
|
* on expiration if bfqq is empty (see
|
|
* __bfq_bfqq_recalc_budget). Thus both quantities
|
|
* remain unchanged after such an expiration, and the
|
|
* following statement therefore assigns to
|
|
* entity->budget the remaining budget on such an
|
|
* expiration.
|
|
*/
|
|
entity->budget = min_t(unsigned long,
|
|
bfq_bfqq_budget_left(bfqq),
|
|
bfqq->max_budget);
|
|
|
|
/*
|
|
* At this point, we have used entity->service to get
|
|
* the budget left (needed for updating
|
|
* entity->budget). Thus we finally can, and have to,
|
|
* reset entity->service. The latter must be reset
|
|
* because bfqq would otherwise be charged again for
|
|
* the service it has received during its previous
|
|
* service slot(s).
|
|
*/
|
|
entity->service = 0;
|
|
|
|
return true;
|
|
}
|
|
|
|
/*
|
|
* We can finally complete expiration, by setting service to 0.
|
|
*/
|
|
entity->service = 0;
|
|
entity->budget = max_t(unsigned long, bfqq->max_budget,
|
|
bfq_serv_to_charge(bfqq->next_rq, bfqq));
|
|
bfq_clear_bfqq_non_blocking_wait_rq(bfqq);
|
|
return false;
|
|
}
|
|
|
|
/*
|
|
* Return the farthest past time instant according to jiffies
|
|
* macros.
|
|
*/
|
|
static unsigned long bfq_smallest_from_now(void)
|
|
{
|
|
return jiffies - MAX_JIFFY_OFFSET;
|
|
}
|
|
|
|
static void bfq_update_bfqq_wr_on_rq_arrival(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq,
|
|
unsigned int old_wr_coeff,
|
|
bool wr_or_deserves_wr,
|
|
bool interactive,
|
|
bool in_burst,
|
|
bool soft_rt)
|
|
{
|
|
if (old_wr_coeff == 1 && wr_or_deserves_wr) {
|
|
/* start a weight-raising period */
|
|
if (interactive) {
|
|
bfqq->service_from_wr = 0;
|
|
bfqq->wr_coeff = bfqd->bfq_wr_coeff;
|
|
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd);
|
|
} else {
|
|
/*
|
|
* No interactive weight raising in progress
|
|
* here: assign minus infinity to
|
|
* wr_start_at_switch_to_srt, to make sure
|
|
* that, at the end of the soft-real-time
|
|
* weight raising periods that is starting
|
|
* now, no interactive weight-raising period
|
|
* may be wrongly considered as still in
|
|
* progress (and thus actually started by
|
|
* mistake).
|
|
*/
|
|
bfqq->wr_start_at_switch_to_srt =
|
|
bfq_smallest_from_now();
|
|
bfqq->wr_coeff = bfqd->bfq_wr_coeff *
|
|
BFQ_SOFTRT_WEIGHT_FACTOR;
|
|
bfqq->wr_cur_max_time =
|
|
bfqd->bfq_wr_rt_max_time;
|
|
}
|
|
|
|
/*
|
|
* If needed, further reduce budget to make sure it is
|
|
* close to bfqq's backlog, so as to reduce the
|
|
* scheduling-error component due to a too large
|
|
* budget. Do not care about throughput consequences,
|
|
* but only about latency. Finally, do not assign a
|
|
* too small budget either, to avoid increasing
|
|
* latency by causing too frequent expirations.
|
|
*/
|
|
bfqq->entity.budget = min_t(unsigned long,
|
|
bfqq->entity.budget,
|
|
2 * bfq_min_budget(bfqd));
|
|
} else if (old_wr_coeff > 1) {
|
|
if (interactive) { /* update wr coeff and duration */
|
|
bfqq->wr_coeff = bfqd->bfq_wr_coeff;
|
|
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd);
|
|
} else if (in_burst)
|
|
bfqq->wr_coeff = 1;
|
|
else if (soft_rt) {
|
|
/*
|
|
* The application is now or still meeting the
|
|
* requirements for being deemed soft rt. We
|
|
* can then correctly and safely (re)charge
|
|
* the weight-raising duration for the
|
|
* application with the weight-raising
|
|
* duration for soft rt applications.
|
|
*
|
|
* In particular, doing this recharge now, i.e.,
|
|
* before the weight-raising period for the
|
|
* application finishes, reduces the probability
|
|
* of the following negative scenario:
|
|
* 1) the weight of a soft rt application is
|
|
* raised at startup (as for any newly
|
|
* created application),
|
|
* 2) since the application is not interactive,
|
|
* at a certain time weight-raising is
|
|
* stopped for the application,
|
|
* 3) at that time the application happens to
|
|
* still have pending requests, and hence
|
|
* is destined to not have a chance to be
|
|
* deemed soft rt before these requests are
|
|
* completed (see the comments to the
|
|
* function bfq_bfqq_softrt_next_start()
|
|
* for details on soft rt detection),
|
|
* 4) these pending requests experience a high
|
|
* latency because the application is not
|
|
* weight-raised while they are pending.
|
|
*/
|
|
if (bfqq->wr_cur_max_time !=
|
|
bfqd->bfq_wr_rt_max_time) {
|
|
bfqq->wr_start_at_switch_to_srt =
|
|
bfqq->last_wr_start_finish;
|
|
|
|
bfqq->wr_cur_max_time =
|
|
bfqd->bfq_wr_rt_max_time;
|
|
bfqq->wr_coeff = bfqd->bfq_wr_coeff *
|
|
BFQ_SOFTRT_WEIGHT_FACTOR;
|
|
}
|
|
bfqq->last_wr_start_finish = jiffies;
|
|
}
|
|
}
|
|
}
|
|
|
|
static bool bfq_bfqq_idle_for_long_time(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
return bfqq->dispatched == 0 &&
|
|
time_is_before_jiffies(
|
|
bfqq->budget_timeout +
|
|
bfqd->bfq_wr_min_idle_time);
|
|
}
|
|
|
|
|
|
/*
|
|
* Return true if bfqq is in a higher priority class, or has a higher
|
|
* weight than the in-service queue.
|
|
*/
|
|
static bool bfq_bfqq_higher_class_or_weight(struct bfq_queue *bfqq,
|
|
struct bfq_queue *in_serv_bfqq)
|
|
{
|
|
int bfqq_weight, in_serv_weight;
|
|
|
|
if (bfqq->ioprio_class < in_serv_bfqq->ioprio_class)
|
|
return true;
|
|
|
|
if (in_serv_bfqq->entity.parent == bfqq->entity.parent) {
|
|
bfqq_weight = bfqq->entity.weight;
|
|
in_serv_weight = in_serv_bfqq->entity.weight;
|
|
} else {
|
|
if (bfqq->entity.parent)
|
|
bfqq_weight = bfqq->entity.parent->weight;
|
|
else
|
|
bfqq_weight = bfqq->entity.weight;
|
|
if (in_serv_bfqq->entity.parent)
|
|
in_serv_weight = in_serv_bfqq->entity.parent->weight;
|
|
else
|
|
in_serv_weight = in_serv_bfqq->entity.weight;
|
|
}
|
|
|
|
return bfqq_weight > in_serv_weight;
|
|
}
|
|
|
|
static bool bfq_better_to_idle(struct bfq_queue *bfqq);
|
|
|
|
static void bfq_bfqq_handle_idle_busy_switch(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq,
|
|
int old_wr_coeff,
|
|
struct request *rq,
|
|
bool *interactive)
|
|
{
|
|
bool soft_rt, in_burst, wr_or_deserves_wr,
|
|
bfqq_wants_to_preempt,
|
|
idle_for_long_time = bfq_bfqq_idle_for_long_time(bfqd, bfqq),
|
|
/*
|
|
* See the comments on
|
|
* bfq_bfqq_update_budg_for_activation for
|
|
* details on the usage of the next variable.
|
|
*/
|
|
arrived_in_time = ktime_get_ns() <=
|
|
bfqq->ttime.last_end_request +
|
|
bfqd->bfq_slice_idle * 3;
|
|
|
|
|
|
/*
|
|
* bfqq deserves to be weight-raised if:
|
|
* - it is sync,
|
|
* - it does not belong to a large burst,
|
|
* - it has been idle for enough time or is soft real-time,
|
|
* - is linked to a bfq_io_cq (it is not shared in any sense),
|
|
* - has a default weight (otherwise we assume the user wanted
|
|
* to control its weight explicitly)
|
|
*/
|
|
in_burst = bfq_bfqq_in_large_burst(bfqq);
|
|
soft_rt = bfqd->bfq_wr_max_softrt_rate > 0 &&
|
|
!BFQQ_TOTALLY_SEEKY(bfqq) &&
|
|
!in_burst &&
|
|
time_is_before_jiffies(bfqq->soft_rt_next_start) &&
|
|
bfqq->dispatched == 0 &&
|
|
bfqq->entity.new_weight == 40;
|
|
*interactive = !in_burst && idle_for_long_time &&
|
|
bfqq->entity.new_weight == 40;
|
|
wr_or_deserves_wr = bfqd->low_latency &&
|
|
(bfqq->wr_coeff > 1 ||
|
|
(bfq_bfqq_sync(bfqq) &&
|
|
bfqq->bic && (*interactive || soft_rt)));
|
|
|
|
/*
|
|
* Using the last flag, update budget and check whether bfqq
|
|
* may want to preempt the in-service queue.
|
|
*/
|
|
bfqq_wants_to_preempt =
|
|
bfq_bfqq_update_budg_for_activation(bfqd, bfqq,
|
|
arrived_in_time);
|
|
|
|
/*
|
|
* If bfqq happened to be activated in a burst, but has been
|
|
* idle for much more than an interactive queue, then we
|
|
* assume that, in the overall I/O initiated in the burst, the
|
|
* I/O associated with bfqq is finished. So bfqq does not need
|
|
* to be treated as a queue belonging to a burst
|
|
* anymore. Accordingly, we reset bfqq's in_large_burst flag
|
|
* if set, and remove bfqq from the burst list if it's
|
|
* there. We do not decrement burst_size, because the fact
|
|
* that bfqq does not need to belong to the burst list any
|
|
* more does not invalidate the fact that bfqq was created in
|
|
* a burst.
|
|
*/
|
|
if (likely(!bfq_bfqq_just_created(bfqq)) &&
|
|
idle_for_long_time &&
|
|
time_is_before_jiffies(
|
|
bfqq->budget_timeout +
|
|
msecs_to_jiffies(10000))) {
|
|
hlist_del_init(&bfqq->burst_list_node);
|
|
bfq_clear_bfqq_in_large_burst(bfqq);
|
|
}
|
|
|
|
bfq_clear_bfqq_just_created(bfqq);
|
|
|
|
if (bfqd->low_latency) {
|
|
if (unlikely(time_is_after_jiffies(bfqq->split_time)))
|
|
/* wraparound */
|
|
bfqq->split_time =
|
|
jiffies - bfqd->bfq_wr_min_idle_time - 1;
|
|
|
|
if (time_is_before_jiffies(bfqq->split_time +
|
|
bfqd->bfq_wr_min_idle_time)) {
|
|
bfq_update_bfqq_wr_on_rq_arrival(bfqd, bfqq,
|
|
old_wr_coeff,
|
|
wr_or_deserves_wr,
|
|
*interactive,
|
|
in_burst,
|
|
soft_rt);
|
|
|
|
if (old_wr_coeff != bfqq->wr_coeff)
|
|
bfqq->entity.prio_changed = 1;
|
|
}
|
|
}
|
|
|
|
bfqq->last_idle_bklogged = jiffies;
|
|
bfqq->service_from_backlogged = 0;
|
|
bfq_clear_bfqq_softrt_update(bfqq);
|
|
|
|
bfq_add_bfqq_busy(bfqd, bfqq);
|
|
|
|
/*
|
|
* Expire in-service queue if preemption may be needed for
|
|
* guarantees or throughput. As for guarantees, we care
|
|
* explicitly about two cases. The first is that bfqq has to
|
|
* recover a service hole, as explained in the comments on
|
|
* bfq_bfqq_update_budg_for_activation(), i.e., that
|
|
* bfqq_wants_to_preempt is true. However, if bfqq does not
|
|
* carry time-critical I/O, then bfqq's bandwidth is less
|
|
* important than that of queues that carry time-critical I/O.
|
|
* So, as a further constraint, we consider this case only if
|
|
* bfqq is at least as weight-raised, i.e., at least as time
|
|
* critical, as the in-service queue.
|
|
*
|
|
* The second case is that bfqq is in a higher priority class,
|
|
* or has a higher weight than the in-service queue. If this
|
|
* condition does not hold, we don't care because, even if
|
|
* bfqq does not start to be served immediately, the resulting
|
|
* delay for bfqq's I/O is however lower or much lower than
|
|
* the ideal completion time to be guaranteed to bfqq's I/O.
|
|
*
|
|
* In both cases, preemption is needed only if, according to
|
|
* the timestamps of both bfqq and of the in-service queue,
|
|
* bfqq actually is the next queue to serve. So, to reduce
|
|
* useless preemptions, the return value of
|
|
* next_queue_may_preempt() is considered in the next compound
|
|
* condition too. Yet next_queue_may_preempt() just checks a
|
|
* simple, necessary condition for bfqq to be the next queue
|
|
* to serve. In fact, to evaluate a sufficient condition, the
|
|
* timestamps of the in-service queue would need to be
|
|
* updated, and this operation is quite costly (see the
|
|
* comments on bfq_bfqq_update_budg_for_activation()).
|
|
*
|
|
* As for throughput, we ask bfq_better_to_idle() whether we
|
|
* still need to plug I/O dispatching. If bfq_better_to_idle()
|
|
* says no, then plugging is not needed any longer, either to
|
|
* boost throughput or to perserve service guarantees. Then
|
|
* the best option is to stop plugging I/O, as not doing so
|
|
* would certainly lower throughput. We may end up in this
|
|
* case if: (1) upon a dispatch attempt, we detected that it
|
|
* was better to plug I/O dispatch, and to wait for a new
|
|
* request to arrive for the currently in-service queue, but
|
|
* (2) this switch of bfqq to busy changes the scenario.
|
|
*/
|
|
if (bfqd->in_service_queue &&
|
|
((bfqq_wants_to_preempt &&
|
|
bfqq->wr_coeff >= bfqd->in_service_queue->wr_coeff) ||
|
|
bfq_bfqq_higher_class_or_weight(bfqq, bfqd->in_service_queue) ||
|
|
!bfq_better_to_idle(bfqd->in_service_queue)) &&
|
|
next_queue_may_preempt(bfqd))
|
|
bfq_bfqq_expire(bfqd, bfqd->in_service_queue,
|
|
false, BFQQE_PREEMPTED);
|
|
}
|
|
|
|
static void bfq_reset_inject_limit(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
/* invalidate baseline total service time */
|
|
bfqq->last_serv_time_ns = 0;
|
|
|
|
/*
|
|
* Reset pointer in case we are waiting for
|
|
* some request completion.
|
|
*/
|
|
bfqd->waited_rq = NULL;
|
|
|
|
/*
|
|
* If bfqq has a short think time, then start by setting the
|
|
* inject limit to 0 prudentially, because the service time of
|
|
* an injected I/O request may be higher than the think time
|
|
* of bfqq, and therefore, if one request was injected when
|
|
* bfqq remains empty, this injected request might delay the
|
|
* service of the next I/O request for bfqq significantly. In
|
|
* case bfqq can actually tolerate some injection, then the
|
|
* adaptive update will however raise the limit soon. This
|
|
* lucky circumstance holds exactly because bfqq has a short
|
|
* think time, and thus, after remaining empty, is likely to
|
|
* get new I/O enqueued---and then completed---before being
|
|
* expired. This is the very pattern that gives the
|
|
* limit-update algorithm the chance to measure the effect of
|
|
* injection on request service times, and then to update the
|
|
* limit accordingly.
|
|
*
|
|
* However, in the following special case, the inject limit is
|
|
* left to 1 even if the think time is short: bfqq's I/O is
|
|
* synchronized with that of some other queue, i.e., bfqq may
|
|
* receive new I/O only after the I/O of the other queue is
|
|
* completed. Keeping the inject limit to 1 allows the
|
|
* blocking I/O to be served while bfqq is in service. And
|
|
* this is very convenient both for bfqq and for overall
|
|
* throughput, as explained in detail in the comments in
|
|
* bfq_update_has_short_ttime().
|
|
*
|
|
* On the opposite end, if bfqq has a long think time, then
|
|
* start directly by 1, because:
|
|
* a) on the bright side, keeping at most one request in
|
|
* service in the drive is unlikely to cause any harm to the
|
|
* latency of bfqq's requests, as the service time of a single
|
|
* request is likely to be lower than the think time of bfqq;
|
|
* b) on the downside, after becoming empty, bfqq is likely to
|
|
* expire before getting its next request. With this request
|
|
* arrival pattern, it is very hard to sample total service
|
|
* times and update the inject limit accordingly (see comments
|
|
* on bfq_update_inject_limit()). So the limit is likely to be
|
|
* never, or at least seldom, updated. As a consequence, by
|
|
* setting the limit to 1, we avoid that no injection ever
|
|
* occurs with bfqq. On the downside, this proactive step
|
|
* further reduces chances to actually compute the baseline
|
|
* total service time. Thus it reduces chances to execute the
|
|
* limit-update algorithm and possibly raise the limit to more
|
|
* than 1.
|
|
*/
|
|
if (bfq_bfqq_has_short_ttime(bfqq))
|
|
bfqq->inject_limit = 0;
|
|
else
|
|
bfqq->inject_limit = 1;
|
|
|
|
bfqq->decrease_time_jif = jiffies;
|
|
}
|
|
|
|
static void bfq_update_io_intensity(struct bfq_queue *bfqq, u64 now_ns)
|
|
{
|
|
u64 tot_io_time = now_ns - bfqq->io_start_time;
|
|
|
|
if (RB_EMPTY_ROOT(&bfqq->sort_list) && bfqq->dispatched == 0)
|
|
bfqq->tot_idle_time +=
|
|
now_ns - bfqq->ttime.last_end_request;
|
|
|
|
if (unlikely(bfq_bfqq_just_created(bfqq)))
|
|
return;
|
|
|
|
/*
|
|
* Must be busy for at least about 80% of the time to be
|
|
* considered I/O bound.
|
|
*/
|
|
if (bfqq->tot_idle_time * 5 > tot_io_time)
|
|
bfq_clear_bfqq_IO_bound(bfqq);
|
|
else
|
|
bfq_mark_bfqq_IO_bound(bfqq);
|
|
|
|
/*
|
|
* Keep an observation window of at most 200 ms in the past
|
|
* from now.
|
|
*/
|
|
if (tot_io_time > 200 * NSEC_PER_MSEC) {
|
|
bfqq->io_start_time = now_ns - (tot_io_time>>1);
|
|
bfqq->tot_idle_time >>= 1;
|
|
}
|
|
}
|
|
|
|
/*
|
|
* Detect whether bfqq's I/O seems synchronized with that of some
|
|
* other queue, i.e., whether bfqq, after remaining empty, happens to
|
|
* receive new I/O only right after some I/O request of the other
|
|
* queue has been completed. We call waker queue the other queue, and
|
|
* we assume, for simplicity, that bfqq may have at most one waker
|
|
* queue.
|
|
*
|
|
* A remarkable throughput boost can be reached by unconditionally
|
|
* injecting the I/O of the waker queue, every time a new
|
|
* bfq_dispatch_request happens to be invoked while I/O is being
|
|
* plugged for bfqq. In addition to boosting throughput, this
|
|
* unblocks bfqq's I/O, thereby improving bandwidth and latency for
|
|
* bfqq. Note that these same results may be achieved with the general
|
|
* injection mechanism, but less effectively. For details on this
|
|
* aspect, see the comments on the choice of the queue for injection
|
|
* in bfq_select_queue().
|
|
*
|
|
* Turning back to the detection of a waker queue, a queue Q is deemed
|
|
* as a waker queue for bfqq if, for three consecutive times, bfqq
|
|
* happens to become non empty right after a request of Q has been
|
|
* completed. In particular, on the first time, Q is tentatively set
|
|
* as a candidate waker queue, while on the third consecutive time
|
|
* that Q is detected, the field waker_bfqq is set to Q, to confirm
|
|
* that Q is a waker queue for bfqq. These detection steps are
|
|
* performed only if bfqq has a long think time, so as to make it more
|
|
* likely that bfqq's I/O is actually being blocked by a
|
|
* synchronization. This last filter, plus the above three-times
|
|
* requirement, make false positives less likely.
|
|
*
|
|
* NOTE
|
|
*
|
|
* The sooner a waker queue is detected, the sooner throughput can be
|
|
* boosted by injecting I/O from the waker queue. Fortunately,
|
|
* detection is likely to be actually fast, for the following
|
|
* reasons. While blocked by synchronization, bfqq has a long think
|
|
* time. This implies that bfqq's inject limit is at least equal to 1
|
|
* (see the comments in bfq_update_inject_limit()). So, thanks to
|
|
* injection, the waker queue is likely to be served during the very
|
|
* first I/O-plugging time interval for bfqq. This triggers the first
|
|
* step of the detection mechanism. Thanks again to injection, the
|
|
* candidate waker queue is then likely to be confirmed no later than
|
|
* during the next I/O-plugging interval for bfqq.
|
|
*
|
|
* ISSUE
|
|
*
|
|
* On queue merging all waker information is lost.
|
|
*/
|
|
static void bfq_check_waker(struct bfq_data *bfqd, struct bfq_queue *bfqq,
|
|
u64 now_ns)
|
|
{
|
|
if (!bfqd->last_completed_rq_bfqq ||
|
|
bfqd->last_completed_rq_bfqq == bfqq ||
|
|
bfq_bfqq_has_short_ttime(bfqq) ||
|
|
now_ns - bfqd->last_completion >= 4 * NSEC_PER_MSEC ||
|
|
bfqd->last_completed_rq_bfqq == bfqq->waker_bfqq)
|
|
return;
|
|
|
|
if (bfqd->last_completed_rq_bfqq !=
|
|
bfqq->tentative_waker_bfqq) {
|
|
/*
|
|
* First synchronization detected with a
|
|
* candidate waker queue, or with a different
|
|
* candidate waker queue from the current one.
|
|
*/
|
|
bfqq->tentative_waker_bfqq =
|
|
bfqd->last_completed_rq_bfqq;
|
|
bfqq->num_waker_detections = 1;
|
|
} else /* Same tentative waker queue detected again */
|
|
bfqq->num_waker_detections++;
|
|
|
|
if (bfqq->num_waker_detections == 3) {
|
|
bfqq->waker_bfqq = bfqd->last_completed_rq_bfqq;
|
|
bfqq->tentative_waker_bfqq = NULL;
|
|
|
|
/*
|
|
* If the waker queue disappears, then
|
|
* bfqq->waker_bfqq must be reset. To
|
|
* this goal, we maintain in each
|
|
* waker queue a list, woken_list, of
|
|
* all the queues that reference the
|
|
* waker queue through their
|
|
* waker_bfqq pointer. When the waker
|
|
* queue exits, the waker_bfqq pointer
|
|
* of all the queues in the woken_list
|
|
* is reset.
|
|
*
|
|
* In addition, if bfqq is already in
|
|
* the woken_list of a waker queue,
|
|
* then, before being inserted into
|
|
* the woken_list of a new waker
|
|
* queue, bfqq must be removed from
|
|
* the woken_list of the old waker
|
|
* queue.
|
|
*/
|
|
if (!hlist_unhashed(&bfqq->woken_list_node))
|
|
hlist_del_init(&bfqq->woken_list_node);
|
|
hlist_add_head(&bfqq->woken_list_node,
|
|
&bfqd->last_completed_rq_bfqq->woken_list);
|
|
}
|
|
}
|
|
|
|
static void bfq_add_request(struct request *rq)
|
|
{
|
|
struct bfq_queue *bfqq = RQ_BFQQ(rq);
|
|
struct bfq_data *bfqd = bfqq->bfqd;
|
|
struct request *next_rq, *prev;
|
|
unsigned int old_wr_coeff = bfqq->wr_coeff;
|
|
bool interactive = false;
|
|
u64 now_ns = ktime_get_ns();
|
|
|
|
bfq_log_bfqq(bfqd, bfqq, "add_request %d", rq_is_sync(rq));
|
|
bfqq->queued[rq_is_sync(rq)]++;
|
|
bfqd->queued++;
|
|
|
|
if (RB_EMPTY_ROOT(&bfqq->sort_list) && bfq_bfqq_sync(bfqq)) {
|
|
bfq_check_waker(bfqd, bfqq, now_ns);
|
|
|
|
/*
|
|
* Periodically reset inject limit, to make sure that
|
|
* the latter eventually drops in case workload
|
|
* changes, see step (3) in the comments on
|
|
* bfq_update_inject_limit().
|
|
*/
|
|
if (time_is_before_eq_jiffies(bfqq->decrease_time_jif +
|
|
msecs_to_jiffies(1000)))
|
|
bfq_reset_inject_limit(bfqd, bfqq);
|
|
|
|
/*
|
|
* The following conditions must hold to setup a new
|
|
* sampling of total service time, and then a new
|
|
* update of the inject limit:
|
|
* - bfqq is in service, because the total service
|
|
* time is evaluated only for the I/O requests of
|
|
* the queues in service;
|
|
* - this is the right occasion to compute or to
|
|
* lower the baseline total service time, because
|
|
* there are actually no requests in the drive,
|
|
* or
|
|
* the baseline total service time is available, and
|
|
* this is the right occasion to compute the other
|
|
* quantity needed to update the inject limit, i.e.,
|
|
* the total service time caused by the amount of
|
|
* injection allowed by the current value of the
|
|
* limit. It is the right occasion because injection
|
|
* has actually been performed during the service
|
|
* hole, and there are still in-flight requests,
|
|
* which are very likely to be exactly the injected
|
|
* requests, or part of them;
|
|
* - the minimum interval for sampling the total
|
|
* service time and updating the inject limit has
|
|
* elapsed.
|
|
*/
|
|
if (bfqq == bfqd->in_service_queue &&
|
|
(bfqd->rq_in_driver == 0 ||
|
|
(bfqq->last_serv_time_ns > 0 &&
|
|
bfqd->rqs_injected && bfqd->rq_in_driver > 0)) &&
|
|
time_is_before_eq_jiffies(bfqq->decrease_time_jif +
|
|
msecs_to_jiffies(10))) {
|
|
bfqd->last_empty_occupied_ns = ktime_get_ns();
|
|
/*
|
|
* Start the state machine for measuring the
|
|
* total service time of rq: setting
|
|
* wait_dispatch will cause bfqd->waited_rq to
|
|
* be set when rq will be dispatched.
|
|
*/
|
|
bfqd->wait_dispatch = true;
|
|
/*
|
|
* If there is no I/O in service in the drive,
|
|
* then possible injection occurred before the
|
|
* arrival of rq will not affect the total
|
|
* service time of rq. So the injection limit
|
|
* must not be updated as a function of such
|
|
* total service time, unless new injection
|
|
* occurs before rq is completed. To have the
|
|
* injection limit updated only in the latter
|
|
* case, reset rqs_injected here (rqs_injected
|
|
* will be set in case injection is performed
|
|
* on bfqq before rq is completed).
|
|
*/
|
|
if (bfqd->rq_in_driver == 0)
|
|
bfqd->rqs_injected = false;
|
|
}
|
|
}
|
|
|
|
if (bfq_bfqq_sync(bfqq))
|
|
bfq_update_io_intensity(bfqq, now_ns);
|
|
|
|
elv_rb_add(&bfqq->sort_list, rq);
|
|
|
|
/*
|
|
* Check if this request is a better next-serve candidate.
|
|
*/
|
|
prev = bfqq->next_rq;
|
|
next_rq = bfq_choose_req(bfqd, bfqq->next_rq, rq, bfqd->last_position);
|
|
bfqq->next_rq = next_rq;
|
|
|
|
/*
|
|
* Adjust priority tree position, if next_rq changes.
|
|
* See comments on bfq_pos_tree_add_move() for the unlikely().
|
|
*/
|
|
if (unlikely(!bfqd->nonrot_with_queueing && prev != bfqq->next_rq))
|
|
bfq_pos_tree_add_move(bfqd, bfqq);
|
|
|
|
if (!bfq_bfqq_busy(bfqq)) /* switching to busy ... */
|
|
bfq_bfqq_handle_idle_busy_switch(bfqd, bfqq, old_wr_coeff,
|
|
rq, &interactive);
|
|
else {
|
|
if (bfqd->low_latency && old_wr_coeff == 1 && !rq_is_sync(rq) &&
|
|
time_is_before_jiffies(
|
|
bfqq->last_wr_start_finish +
|
|
bfqd->bfq_wr_min_inter_arr_async)) {
|
|
bfqq->wr_coeff = bfqd->bfq_wr_coeff;
|
|
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd);
|
|
|
|
bfqd->wr_busy_queues++;
|
|
bfqq->entity.prio_changed = 1;
|
|
}
|
|
if (prev != bfqq->next_rq)
|
|
bfq_updated_next_req(bfqd, bfqq);
|
|
}
|
|
|
|
/*
|
|
* Assign jiffies to last_wr_start_finish in the following
|
|
* cases:
|
|
*
|
|
* . if bfqq is not going to be weight-raised, because, for
|
|
* non weight-raised queues, last_wr_start_finish stores the
|
|
* arrival time of the last request; as of now, this piece
|
|
* of information is used only for deciding whether to
|
|
* weight-raise async queues
|
|
*
|
|
* . if bfqq is not weight-raised, because, if bfqq is now
|
|
* switching to weight-raised, then last_wr_start_finish
|
|
* stores the time when weight-raising starts
|
|
*
|
|
* . if bfqq is interactive, because, regardless of whether
|
|
* bfqq is currently weight-raised, the weight-raising
|
|
* period must start or restart (this case is considered
|
|
* separately because it is not detected by the above
|
|
* conditions, if bfqq is already weight-raised)
|
|
*
|
|
* last_wr_start_finish has to be updated also if bfqq is soft
|
|
* real-time, because the weight-raising period is constantly
|
|
* restarted on idle-to-busy transitions for these queues, but
|
|
* this is already done in bfq_bfqq_handle_idle_busy_switch if
|
|
* needed.
|
|
*/
|
|
if (bfqd->low_latency &&
|
|
(old_wr_coeff == 1 || bfqq->wr_coeff == 1 || interactive))
|
|
bfqq->last_wr_start_finish = jiffies;
|
|
}
|
|
|
|
static struct request *bfq_find_rq_fmerge(struct bfq_data *bfqd,
|
|
struct bio *bio,
|
|
struct request_queue *q)
|
|
{
|
|
struct bfq_queue *bfqq = bfqd->bio_bfqq;
|
|
|
|
|
|
if (bfqq)
|
|
return elv_rb_find(&bfqq->sort_list, bio_end_sector(bio));
|
|
|
|
return NULL;
|
|
}
|
|
|
|
static sector_t get_sdist(sector_t last_pos, struct request *rq)
|
|
{
|
|
if (last_pos)
|
|
return abs(blk_rq_pos(rq) - last_pos);
|
|
|
|
return 0;
|
|
}
|
|
|
|
#if 0 /* Still not clear if we can do without next two functions */
|
|
static void bfq_activate_request(struct request_queue *q, struct request *rq)
|
|
{
|
|
struct bfq_data *bfqd = q->elevator->elevator_data;
|
|
|
|
bfqd->rq_in_driver++;
|
|
}
|
|
|
|
static void bfq_deactivate_request(struct request_queue *q, struct request *rq)
|
|
{
|
|
struct bfq_data *bfqd = q->elevator->elevator_data;
|
|
|
|
bfqd->rq_in_driver--;
|
|
}
|
|
#endif
|
|
|
|
static void bfq_remove_request(struct request_queue *q,
|
|
struct request *rq)
|
|
{
|
|
struct bfq_queue *bfqq = RQ_BFQQ(rq);
|
|
struct bfq_data *bfqd = bfqq->bfqd;
|
|
const int sync = rq_is_sync(rq);
|
|
|
|
if (bfqq->next_rq == rq) {
|
|
bfqq->next_rq = bfq_find_next_rq(bfqd, bfqq, rq);
|
|
bfq_updated_next_req(bfqd, bfqq);
|
|
}
|
|
|
|
if (rq->queuelist.prev != &rq->queuelist)
|
|
list_del_init(&rq->queuelist);
|
|
bfqq->queued[sync]--;
|
|
bfqd->queued--;
|
|
elv_rb_del(&bfqq->sort_list, rq);
|
|
|
|
elv_rqhash_del(q, rq);
|
|
if (q->last_merge == rq)
|
|
q->last_merge = NULL;
|
|
|
|
if (RB_EMPTY_ROOT(&bfqq->sort_list)) {
|
|
bfqq->next_rq = NULL;
|
|
|
|
if (bfq_bfqq_busy(bfqq) && bfqq != bfqd->in_service_queue) {
|
|
bfq_del_bfqq_busy(bfqd, bfqq, false);
|
|
/*
|
|
* bfqq emptied. In normal operation, when
|
|
* bfqq is empty, bfqq->entity.service and
|
|
* bfqq->entity.budget must contain,
|
|
* respectively, the service received and the
|
|
* budget used last time bfqq emptied. These
|
|
* facts do not hold in this case, as at least
|
|
* this last removal occurred while bfqq is
|
|
* not in service. To avoid inconsistencies,
|
|
* reset both bfqq->entity.service and
|
|
* bfqq->entity.budget, if bfqq has still a
|
|
* process that may issue I/O requests to it.
|
|
*/
|
|
bfqq->entity.budget = bfqq->entity.service = 0;
|
|
}
|
|
|
|
/*
|
|
* Remove queue from request-position tree as it is empty.
|
|
*/
|
|
if (bfqq->pos_root) {
|
|
rb_erase(&bfqq->pos_node, bfqq->pos_root);
|
|
bfqq->pos_root = NULL;
|
|
}
|
|
} else {
|
|
/* see comments on bfq_pos_tree_add_move() for the unlikely() */
|
|
if (unlikely(!bfqd->nonrot_with_queueing))
|
|
bfq_pos_tree_add_move(bfqd, bfqq);
|
|
}
|
|
|
|
if (rq->cmd_flags & REQ_META)
|
|
bfqq->meta_pending--;
|
|
|
|
}
|
|
|
|
static bool bfq_bio_merge(struct request_queue *q, struct bio *bio,
|
|
unsigned int nr_segs)
|
|
{
|
|
struct bfq_data *bfqd = q->elevator->elevator_data;
|
|
struct request *free = NULL;
|
|
/*
|
|
* bfq_bic_lookup grabs the queue_lock: invoke it now and
|
|
* store its return value for later use, to avoid nesting
|
|
* queue_lock inside the bfqd->lock. We assume that the bic
|
|
* returned by bfq_bic_lookup does not go away before
|
|
* bfqd->lock is taken.
|
|
*/
|
|
struct bfq_io_cq *bic = bfq_bic_lookup(bfqd, current->io_context, q);
|
|
bool ret;
|
|
|
|
spin_lock_irq(&bfqd->lock);
|
|
|
|
if (bic)
|
|
bfqd->bio_bfqq = bic_to_bfqq(bic, op_is_sync(bio->bi_opf));
|
|
else
|
|
bfqd->bio_bfqq = NULL;
|
|
bfqd->bio_bic = bic;
|
|
|
|
ret = blk_mq_sched_try_merge(q, bio, nr_segs, &free);
|
|
|
|
if (free)
|
|
blk_mq_free_request(free);
|
|
spin_unlock_irq(&bfqd->lock);
|
|
|
|
return ret;
|
|
}
|
|
|
|
static int bfq_request_merge(struct request_queue *q, struct request **req,
|
|
struct bio *bio)
|
|
{
|
|
struct bfq_data *bfqd = q->elevator->elevator_data;
|
|
struct request *__rq;
|
|
|
|
__rq = bfq_find_rq_fmerge(bfqd, bio, q);
|
|
if (__rq && elv_bio_merge_ok(__rq, bio)) {
|
|
*req = __rq;
|
|
return ELEVATOR_FRONT_MERGE;
|
|
}
|
|
|
|
return ELEVATOR_NO_MERGE;
|
|
}
|
|
|
|
static struct bfq_queue *bfq_init_rq(struct request *rq);
|
|
|
|
static void bfq_request_merged(struct request_queue *q, struct request *req,
|
|
enum elv_merge type)
|
|
{
|
|
if (type == ELEVATOR_FRONT_MERGE &&
|
|
rb_prev(&req->rb_node) &&
|
|
blk_rq_pos(req) <
|
|
blk_rq_pos(container_of(rb_prev(&req->rb_node),
|
|
struct request, rb_node))) {
|
|
struct bfq_queue *bfqq = bfq_init_rq(req);
|
|
struct bfq_data *bfqd;
|
|
struct request *prev, *next_rq;
|
|
|
|
if (!bfqq)
|
|
return;
|
|
|
|
bfqd = bfqq->bfqd;
|
|
|
|
/* Reposition request in its sort_list */
|
|
elv_rb_del(&bfqq->sort_list, req);
|
|
elv_rb_add(&bfqq->sort_list, req);
|
|
|
|
/* Choose next request to be served for bfqq */
|
|
prev = bfqq->next_rq;
|
|
next_rq = bfq_choose_req(bfqd, bfqq->next_rq, req,
|
|
bfqd->last_position);
|
|
bfqq->next_rq = next_rq;
|
|
/*
|
|
* If next_rq changes, update both the queue's budget to
|
|
* fit the new request and the queue's position in its
|
|
* rq_pos_tree.
|
|
*/
|
|
if (prev != bfqq->next_rq) {
|
|
bfq_updated_next_req(bfqd, bfqq);
|
|
/*
|
|
* See comments on bfq_pos_tree_add_move() for
|
|
* the unlikely().
|
|
*/
|
|
if (unlikely(!bfqd->nonrot_with_queueing))
|
|
bfq_pos_tree_add_move(bfqd, bfqq);
|
|
}
|
|
}
|
|
}
|
|
|
|
/*
|
|
* This function is called to notify the scheduler that the requests
|
|
* rq and 'next' have been merged, with 'next' going away. BFQ
|
|
* exploits this hook to address the following issue: if 'next' has a
|
|
* fifo_time lower that rq, then the fifo_time of rq must be set to
|
|
* the value of 'next', to not forget the greater age of 'next'.
|
|
*
|
|
* NOTE: in this function we assume that rq is in a bfq_queue, basing
|
|
* on that rq is picked from the hash table q->elevator->hash, which,
|
|
* in its turn, is filled only with I/O requests present in
|
|
* bfq_queues, while BFQ is in use for the request queue q. In fact,
|
|
* the function that fills this hash table (elv_rqhash_add) is called
|
|
* only by bfq_insert_request.
|
|
*/
|
|
static void bfq_requests_merged(struct request_queue *q, struct request *rq,
|
|
struct request *next)
|
|
{
|
|
struct bfq_queue *bfqq = bfq_init_rq(rq),
|
|
*next_bfqq = bfq_init_rq(next);
|
|
|
|
if (!bfqq)
|
|
return;
|
|
|
|
/*
|
|
* If next and rq belong to the same bfq_queue and next is older
|
|
* than rq, then reposition rq in the fifo (by substituting next
|
|
* with rq). Otherwise, if next and rq belong to different
|
|
* bfq_queues, never reposition rq: in fact, we would have to
|
|
* reposition it with respect to next's position in its own fifo,
|
|
* which would most certainly be too expensive with respect to
|
|
* the benefits.
|
|
*/
|
|
if (bfqq == next_bfqq &&
|
|
!list_empty(&rq->queuelist) && !list_empty(&next->queuelist) &&
|
|
next->fifo_time < rq->fifo_time) {
|
|
list_del_init(&rq->queuelist);
|
|
list_replace_init(&next->queuelist, &rq->queuelist);
|
|
rq->fifo_time = next->fifo_time;
|
|
}
|
|
|
|
if (bfqq->next_rq == next)
|
|
bfqq->next_rq = rq;
|
|
|
|
bfqg_stats_update_io_merged(bfqq_group(bfqq), next->cmd_flags);
|
|
}
|
|
|
|
/* Must be called with bfqq != NULL */
|
|
static void bfq_bfqq_end_wr(struct bfq_queue *bfqq)
|
|
{
|
|
/*
|
|
* If bfqq has been enjoying interactive weight-raising, then
|
|
* reset soft_rt_next_start. We do it for the following
|
|
* reason. bfqq may have been conveying the I/O needed to load
|
|
* a soft real-time application. Such an application actually
|
|
* exhibits a soft real-time I/O pattern after it finishes
|
|
* loading, and finally starts doing its job. But, if bfqq has
|
|
* been receiving a lot of bandwidth so far (likely to happen
|
|
* on a fast device), then soft_rt_next_start now contains a
|
|
* high value that. So, without this reset, bfqq would be
|
|
* prevented from being possibly considered as soft_rt for a
|
|
* very long time.
|
|
*/
|
|
|
|
if (bfqq->wr_cur_max_time !=
|
|
bfqq->bfqd->bfq_wr_rt_max_time)
|
|
bfqq->soft_rt_next_start = jiffies;
|
|
|
|
if (bfq_bfqq_busy(bfqq))
|
|
bfqq->bfqd->wr_busy_queues--;
|
|
bfqq->wr_coeff = 1;
|
|
bfqq->wr_cur_max_time = 0;
|
|
bfqq->last_wr_start_finish = jiffies;
|
|
/*
|
|
* Trigger a weight change on the next invocation of
|
|
* __bfq_entity_update_weight_prio.
|
|
*/
|
|
bfqq->entity.prio_changed = 1;
|
|
}
|
|
|
|
void bfq_end_wr_async_queues(struct bfq_data *bfqd,
|
|
struct bfq_group *bfqg)
|
|
{
|
|
int i, j;
|
|
|
|
for (i = 0; i < 2; i++)
|
|
for (j = 0; j < IOPRIO_BE_NR; j++)
|
|
if (bfqg->async_bfqq[i][j])
|
|
bfq_bfqq_end_wr(bfqg->async_bfqq[i][j]);
|
|
if (bfqg->async_idle_bfqq)
|
|
bfq_bfqq_end_wr(bfqg->async_idle_bfqq);
|
|
}
|
|
|
|
static void bfq_end_wr(struct bfq_data *bfqd)
|
|
{
|
|
struct bfq_queue *bfqq;
|
|
|
|
spin_lock_irq(&bfqd->lock);
|
|
|
|
list_for_each_entry(bfqq, &bfqd->active_list, bfqq_list)
|
|
bfq_bfqq_end_wr(bfqq);
|
|
list_for_each_entry(bfqq, &bfqd->idle_list, bfqq_list)
|
|
bfq_bfqq_end_wr(bfqq);
|
|
bfq_end_wr_async(bfqd);
|
|
|
|
spin_unlock_irq(&bfqd->lock);
|
|
}
|
|
|
|
static sector_t bfq_io_struct_pos(void *io_struct, bool request)
|
|
{
|
|
if (request)
|
|
return blk_rq_pos(io_struct);
|
|
else
|
|
return ((struct bio *)io_struct)->bi_iter.bi_sector;
|
|
}
|
|
|
|
static int bfq_rq_close_to_sector(void *io_struct, bool request,
|
|
sector_t sector)
|
|
{
|
|
return abs(bfq_io_struct_pos(io_struct, request) - sector) <=
|
|
BFQQ_CLOSE_THR;
|
|
}
|
|
|
|
static struct bfq_queue *bfqq_find_close(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq,
|
|
sector_t sector)
|
|
{
|
|
struct rb_root *root = &bfq_bfqq_to_bfqg(bfqq)->rq_pos_tree;
|
|
struct rb_node *parent, *node;
|
|
struct bfq_queue *__bfqq;
|
|
|
|
if (RB_EMPTY_ROOT(root))
|
|
return NULL;
|
|
|
|
/*
|
|
* First, if we find a request starting at the end of the last
|
|
* request, choose it.
|
|
*/
|
|
__bfqq = bfq_rq_pos_tree_lookup(bfqd, root, sector, &parent, NULL);
|
|
if (__bfqq)
|
|
return __bfqq;
|
|
|
|
/*
|
|
* If the exact sector wasn't found, the parent of the NULL leaf
|
|
* will contain the closest sector (rq_pos_tree sorted by
|
|
* next_request position).
|
|
*/
|
|
__bfqq = rb_entry(parent, struct bfq_queue, pos_node);
|
|
if (bfq_rq_close_to_sector(__bfqq->next_rq, true, sector))
|
|
return __bfqq;
|
|
|
|
if (blk_rq_pos(__bfqq->next_rq) < sector)
|
|
node = rb_next(&__bfqq->pos_node);
|
|
else
|
|
node = rb_prev(&__bfqq->pos_node);
|
|
if (!node)
|
|
return NULL;
|
|
|
|
__bfqq = rb_entry(node, struct bfq_queue, pos_node);
|
|
if (bfq_rq_close_to_sector(__bfqq->next_rq, true, sector))
|
|
return __bfqq;
|
|
|
|
return NULL;
|
|
}
|
|
|
|
static struct bfq_queue *bfq_find_close_cooperator(struct bfq_data *bfqd,
|
|
struct bfq_queue *cur_bfqq,
|
|
sector_t sector)
|
|
{
|
|
struct bfq_queue *bfqq;
|
|
|
|
/*
|
|
* We shall notice if some of the queues are cooperating,
|
|
* e.g., working closely on the same area of the device. In
|
|
* that case, we can group them together and: 1) don't waste
|
|
* time idling, and 2) serve the union of their requests in
|
|
* the best possible order for throughput.
|
|
*/
|
|
bfqq = bfqq_find_close(bfqd, cur_bfqq, sector);
|
|
if (!bfqq || bfqq == cur_bfqq)
|
|
return NULL;
|
|
|
|
return bfqq;
|
|
}
|
|
|
|
static struct bfq_queue *
|
|
bfq_setup_merge(struct bfq_queue *bfqq, struct bfq_queue *new_bfqq)
|
|
{
|
|
int process_refs, new_process_refs;
|
|
struct bfq_queue *__bfqq;
|
|
|
|
/*
|
|
* If there are no process references on the new_bfqq, then it is
|
|
* unsafe to follow the ->new_bfqq chain as other bfqq's in the chain
|
|
* may have dropped their last reference (not just their last process
|
|
* reference).
|
|
*/
|
|
if (!bfqq_process_refs(new_bfqq))
|
|
return NULL;
|
|
|
|
/* Avoid a circular list and skip interim queue merges. */
|
|
while ((__bfqq = new_bfqq->new_bfqq)) {
|
|
if (__bfqq == bfqq)
|
|
return NULL;
|
|
new_bfqq = __bfqq;
|
|
}
|
|
|
|
process_refs = bfqq_process_refs(bfqq);
|
|
new_process_refs = bfqq_process_refs(new_bfqq);
|
|
/*
|
|
* If the process for the bfqq has gone away, there is no
|
|
* sense in merging the queues.
|
|
*/
|
|
if (process_refs == 0 || new_process_refs == 0)
|
|
return NULL;
|
|
|
|
bfq_log_bfqq(bfqq->bfqd, bfqq, "scheduling merge with queue %d",
|
|
new_bfqq->pid);
|
|
|
|
/*
|
|
* Merging is just a redirection: the requests of the process
|
|
* owning one of the two queues are redirected to the other queue.
|
|
* The latter queue, in its turn, is set as shared if this is the
|
|
* first time that the requests of some process are redirected to
|
|
* it.
|
|
*
|
|
* We redirect bfqq to new_bfqq and not the opposite, because
|
|
* we are in the context of the process owning bfqq, thus we
|
|
* have the io_cq of this process. So we can immediately
|
|
* configure this io_cq to redirect the requests of the
|
|
* process to new_bfqq. In contrast, the io_cq of new_bfqq is
|
|
* not available any more (new_bfqq->bic == NULL).
|
|
*
|
|
* Anyway, even in case new_bfqq coincides with the in-service
|
|
* queue, redirecting requests the in-service queue is the
|
|
* best option, as we feed the in-service queue with new
|
|
* requests close to the last request served and, by doing so,
|
|
* are likely to increase the throughput.
|
|
*/
|
|
bfqq->new_bfqq = new_bfqq;
|
|
new_bfqq->ref += process_refs;
|
|
return new_bfqq;
|
|
}
|
|
|
|
static bool bfq_may_be_close_cooperator(struct bfq_queue *bfqq,
|
|
struct bfq_queue *new_bfqq)
|
|
{
|
|
if (bfq_too_late_for_merging(new_bfqq))
|
|
return false;
|
|
|
|
if (bfq_class_idle(bfqq) || bfq_class_idle(new_bfqq) ||
|
|
(bfqq->ioprio_class != new_bfqq->ioprio_class))
|
|
return false;
|
|
|
|
/*
|
|
* If either of the queues has already been detected as seeky,
|
|
* then merging it with the other queue is unlikely to lead to
|
|
* sequential I/O.
|
|
*/
|
|
if (BFQQ_SEEKY(bfqq) || BFQQ_SEEKY(new_bfqq))
|
|
return false;
|
|
|
|
/*
|
|
* Interleaved I/O is known to be done by (some) applications
|
|
* only for reads, so it does not make sense to merge async
|
|
* queues.
|
|
*/
|
|
if (!bfq_bfqq_sync(bfqq) || !bfq_bfqq_sync(new_bfqq))
|
|
return false;
|
|
|
|
return true;
|
|
}
|
|
|
|
static bool idling_boosts_thr_without_issues(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq);
|
|
|
|
/*
|
|
* Attempt to schedule a merge of bfqq with the currently in-service
|
|
* queue or with a close queue among the scheduled queues. Return
|
|
* NULL if no merge was scheduled, a pointer to the shared bfq_queue
|
|
* structure otherwise.
|
|
*
|
|
* The OOM queue is not allowed to participate to cooperation: in fact, since
|
|
* the requests temporarily redirected to the OOM queue could be redirected
|
|
* again to dedicated queues at any time, the state needed to correctly
|
|
* handle merging with the OOM queue would be quite complex and expensive
|
|
* to maintain. Besides, in such a critical condition as an out of memory,
|
|
* the benefits of queue merging may be little relevant, or even negligible.
|
|
*
|
|
* WARNING: queue merging may impair fairness among non-weight raised
|
|
* queues, for at least two reasons: 1) the original weight of a
|
|
* merged queue may change during the merged state, 2) even being the
|
|
* weight the same, a merged queue may be bloated with many more
|
|
* requests than the ones produced by its originally-associated
|
|
* process.
|
|
*/
|
|
static struct bfq_queue *
|
|
bfq_setup_cooperator(struct bfq_data *bfqd, struct bfq_queue *bfqq,
|
|
void *io_struct, bool request, struct bfq_io_cq *bic)
|
|
{
|
|
struct bfq_queue *in_service_bfqq, *new_bfqq;
|
|
|
|
/*
|
|
* Check delayed stable merge for rotational or non-queueing
|
|
* devs. For this branch to be executed, bfqq must not be
|
|
* currently merged with some other queue (i.e., bfqq->bic
|
|
* must be non null). If we considered also merged queues,
|
|
* then we should also check whether bfqq has already been
|
|
* merged with bic->stable_merge_bfqq. But this would be
|
|
* costly and complicated.
|
|
*/
|
|
if (unlikely(!bfqd->nonrot_with_queueing)) {
|
|
if (bic->stable_merge_bfqq &&
|
|
!bfq_bfqq_just_created(bfqq) &&
|
|
time_is_after_jiffies(bfqq->split_time +
|
|
msecs_to_jiffies(200))) {
|
|
struct bfq_queue *stable_merge_bfqq =
|
|
bic->stable_merge_bfqq;
|
|
int proc_ref = min(bfqq_process_refs(bfqq),
|
|
bfqq_process_refs(stable_merge_bfqq));
|
|
|
|
/* deschedule stable merge, because done or aborted here */
|
|
bfq_put_stable_ref(stable_merge_bfqq);
|
|
|
|
bic->stable_merge_bfqq = NULL;
|
|
|
|
if (!idling_boosts_thr_without_issues(bfqd, bfqq) &&
|
|
proc_ref > 0) {
|
|
/* next function will take at least one ref */
|
|
struct bfq_queue *new_bfqq =
|
|
bfq_setup_merge(bfqq, stable_merge_bfqq);
|
|
|
|
bic->stably_merged = true;
|
|
if (new_bfqq && new_bfqq->bic)
|
|
new_bfqq->bic->stably_merged = true;
|
|
return new_bfqq;
|
|
} else
|
|
return NULL;
|
|
}
|
|
}
|
|
|
|
/*
|
|
* Do not perform queue merging if the device is non
|
|
* rotational and performs internal queueing. In fact, such a
|
|
* device reaches a high speed through internal parallelism
|
|
* and pipelining. This means that, to reach a high
|
|
* throughput, it must have many requests enqueued at the same
|
|
* time. But, in this configuration, the internal scheduling
|
|
* algorithm of the device does exactly the job of queue
|
|
* merging: it reorders requests so as to obtain as much as
|
|
* possible a sequential I/O pattern. As a consequence, with
|
|
* the workload generated by processes doing interleaved I/O,
|
|
* the throughput reached by the device is likely to be the
|
|
* same, with and without queue merging.
|
|
*
|
|
* Disabling merging also provides a remarkable benefit in
|
|
* terms of throughput. Merging tends to make many workloads
|
|
* artificially more uneven, because of shared queues
|
|
* remaining non empty for incomparably more time than
|
|
* non-merged queues. This may accentuate workload
|
|
* asymmetries. For example, if one of the queues in a set of
|
|
* merged queues has a higher weight than a normal queue, then
|
|
* the shared queue may inherit such a high weight and, by
|
|
* staying almost always active, may force BFQ to perform I/O
|
|
* plugging most of the time. This evidently makes it harder
|
|
* for BFQ to let the device reach a high throughput.
|
|
*
|
|
* Finally, the likely() macro below is not used because one
|
|
* of the two branches is more likely than the other, but to
|
|
* have the code path after the following if() executed as
|
|
* fast as possible for the case of a non rotational device
|
|
* with queueing. We want it because this is the fastest kind
|
|
* of device. On the opposite end, the likely() may lengthen
|
|
* the execution time of BFQ for the case of slower devices
|
|
* (rotational or at least without queueing). But in this case
|
|
* the execution time of BFQ matters very little, if not at
|
|
* all.
|
|
*/
|
|
if (likely(bfqd->nonrot_with_queueing))
|
|
return NULL;
|
|
|
|
/*
|
|
* Prevent bfqq from being merged if it has been created too
|
|
* long ago. The idea is that true cooperating processes, and
|
|
* thus their associated bfq_queues, are supposed to be
|
|
* created shortly after each other. This is the case, e.g.,
|
|
* for KVM/QEMU and dump I/O threads. Basing on this
|
|
* assumption, the following filtering greatly reduces the
|
|
* probability that two non-cooperating processes, which just
|
|
* happen to do close I/O for some short time interval, have
|
|
* their queues merged by mistake.
|
|
*/
|
|
if (bfq_too_late_for_merging(bfqq))
|
|
return NULL;
|
|
|
|
if (bfqq->new_bfqq)
|
|
return bfqq->new_bfqq;
|
|
|
|
if (!io_struct || unlikely(bfqq == &bfqd->oom_bfqq))
|
|
return NULL;
|
|
|
|
/* If there is only one backlogged queue, don't search. */
|
|
if (bfq_tot_busy_queues(bfqd) == 1)
|
|
return NULL;
|
|
|
|
in_service_bfqq = bfqd->in_service_queue;
|
|
|
|
if (in_service_bfqq && in_service_bfqq != bfqq &&
|
|
likely(in_service_bfqq != &bfqd->oom_bfqq) &&
|
|
bfq_rq_close_to_sector(io_struct, request,
|
|
bfqd->in_serv_last_pos) &&
|
|
bfqq->entity.parent == in_service_bfqq->entity.parent &&
|
|
bfq_may_be_close_cooperator(bfqq, in_service_bfqq)) {
|
|
new_bfqq = bfq_setup_merge(bfqq, in_service_bfqq);
|
|
if (new_bfqq)
|
|
return new_bfqq;
|
|
}
|
|
/*
|
|
* Check whether there is a cooperator among currently scheduled
|
|
* queues. The only thing we need is that the bio/request is not
|
|
* NULL, as we need it to establish whether a cooperator exists.
|
|
*/
|
|
new_bfqq = bfq_find_close_cooperator(bfqd, bfqq,
|
|
bfq_io_struct_pos(io_struct, request));
|
|
|
|
if (new_bfqq && likely(new_bfqq != &bfqd->oom_bfqq) &&
|
|
bfq_may_be_close_cooperator(bfqq, new_bfqq))
|
|
return bfq_setup_merge(bfqq, new_bfqq);
|
|
|
|
return NULL;
|
|
}
|
|
|
|
static void bfq_bfqq_save_state(struct bfq_queue *bfqq)
|
|
{
|
|
struct bfq_io_cq *bic = bfqq->bic;
|
|
|
|
/*
|
|
* If !bfqq->bic, the queue is already shared or its requests
|
|
* have already been redirected to a shared queue; both idle window
|
|
* and weight raising state have already been saved. Do nothing.
|
|
*/
|
|
if (!bic)
|
|
return;
|
|
|
|
bic->saved_last_serv_time_ns = bfqq->last_serv_time_ns;
|
|
bic->saved_inject_limit = bfqq->inject_limit;
|
|
bic->saved_decrease_time_jif = bfqq->decrease_time_jif;
|
|
|
|
bic->saved_weight = bfqq->entity.orig_weight;
|
|
bic->saved_ttime = bfqq->ttime;
|
|
bic->saved_has_short_ttime = bfq_bfqq_has_short_ttime(bfqq);
|
|
bic->saved_IO_bound = bfq_bfqq_IO_bound(bfqq);
|
|
bic->saved_io_start_time = bfqq->io_start_time;
|
|
bic->saved_tot_idle_time = bfqq->tot_idle_time;
|
|
bic->saved_in_large_burst = bfq_bfqq_in_large_burst(bfqq);
|
|
bic->was_in_burst_list = !hlist_unhashed(&bfqq->burst_list_node);
|
|
if (unlikely(bfq_bfqq_just_created(bfqq) &&
|
|
!bfq_bfqq_in_large_burst(bfqq) &&
|
|
bfqq->bfqd->low_latency)) {
|
|
/*
|
|
* bfqq being merged right after being created: bfqq
|
|
* would have deserved interactive weight raising, but
|
|
* did not make it to be set in a weight-raised state,
|
|
* because of this early merge. Store directly the
|
|
* weight-raising state that would have been assigned
|
|
* to bfqq, so that to avoid that bfqq unjustly fails
|
|
* to enjoy weight raising if split soon.
|
|
*/
|
|
bic->saved_wr_coeff = bfqq->bfqd->bfq_wr_coeff;
|
|
bic->saved_wr_start_at_switch_to_srt = bfq_smallest_from_now();
|
|
bic->saved_wr_cur_max_time = bfq_wr_duration(bfqq->bfqd);
|
|
bic->saved_last_wr_start_finish = jiffies;
|
|
} else {
|
|
bic->saved_wr_coeff = bfqq->wr_coeff;
|
|
bic->saved_wr_start_at_switch_to_srt =
|
|
bfqq->wr_start_at_switch_to_srt;
|
|
bic->saved_service_from_wr = bfqq->service_from_wr;
|
|
bic->saved_last_wr_start_finish = bfqq->last_wr_start_finish;
|
|
bic->saved_wr_cur_max_time = bfqq->wr_cur_max_time;
|
|
}
|
|
}
|
|
|
|
|
|
static void
|
|
bfq_reassign_last_bfqq(struct bfq_queue *cur_bfqq, struct bfq_queue *new_bfqq)
|
|
{
|
|
if (cur_bfqq->entity.parent &&
|
|
cur_bfqq->entity.parent->last_bfqq_created == cur_bfqq)
|
|
cur_bfqq->entity.parent->last_bfqq_created = new_bfqq;
|
|
else if (cur_bfqq->bfqd && cur_bfqq->bfqd->last_bfqq_created == cur_bfqq)
|
|
cur_bfqq->bfqd->last_bfqq_created = new_bfqq;
|
|
}
|
|
|
|
void bfq_release_process_ref(struct bfq_data *bfqd, struct bfq_queue *bfqq)
|
|
{
|
|
/*
|
|
* To prevent bfqq's service guarantees from being violated,
|
|
* bfqq may be left busy, i.e., queued for service, even if
|
|
* empty (see comments in __bfq_bfqq_expire() for
|
|
* details). But, if no process will send requests to bfqq any
|
|
* longer, then there is no point in keeping bfqq queued for
|
|
* service. In addition, keeping bfqq queued for service, but
|
|
* with no process ref any longer, may have caused bfqq to be
|
|
* freed when dequeued from service. But this is assumed to
|
|
* never happen.
|
|
*/
|
|
if (bfq_bfqq_busy(bfqq) && RB_EMPTY_ROOT(&bfqq->sort_list) &&
|
|
bfqq != bfqd->in_service_queue)
|
|
bfq_del_bfqq_busy(bfqd, bfqq, false);
|
|
|
|
bfq_reassign_last_bfqq(bfqq, NULL);
|
|
|
|
bfq_put_queue(bfqq);
|
|
}
|
|
|
|
static void
|
|
bfq_merge_bfqqs(struct bfq_data *bfqd, struct bfq_io_cq *bic,
|
|
struct bfq_queue *bfqq, struct bfq_queue *new_bfqq)
|
|
{
|
|
bfq_log_bfqq(bfqd, bfqq, "merging with queue %lu",
|
|
(unsigned long)new_bfqq->pid);
|
|
/* Save weight raising and idle window of the merged queues */
|
|
bfq_bfqq_save_state(bfqq);
|
|
bfq_bfqq_save_state(new_bfqq);
|
|
if (bfq_bfqq_IO_bound(bfqq))
|
|
bfq_mark_bfqq_IO_bound(new_bfqq);
|
|
bfq_clear_bfqq_IO_bound(bfqq);
|
|
|
|
/*
|
|
* The processes associated with bfqq are cooperators of the
|
|
* processes associated with new_bfqq. So, if bfqq has a
|
|
* waker, then assume that all these processes will be happy
|
|
* to let bfqq's waker freely inject I/O when they have no
|
|
* I/O.
|
|
*/
|
|
if (bfqq->waker_bfqq && !new_bfqq->waker_bfqq &&
|
|
bfqq->waker_bfqq != new_bfqq) {
|
|
new_bfqq->waker_bfqq = bfqq->waker_bfqq;
|
|
new_bfqq->tentative_waker_bfqq = NULL;
|
|
|
|
/*
|
|
* If the waker queue disappears, then
|
|
* new_bfqq->waker_bfqq must be reset. So insert
|
|
* new_bfqq into the woken_list of the waker. See
|
|
* bfq_check_waker for details.
|
|
*/
|
|
hlist_add_head(&new_bfqq->woken_list_node,
|
|
&new_bfqq->waker_bfqq->woken_list);
|
|
|
|
}
|
|
|
|
/*
|
|
* If bfqq is weight-raised, then let new_bfqq inherit
|
|
* weight-raising. To reduce false positives, neglect the case
|
|
* where bfqq has just been created, but has not yet made it
|
|
* to be weight-raised (which may happen because EQM may merge
|
|
* bfqq even before bfq_add_request is executed for the first
|
|
* time for bfqq). Handling this case would however be very
|
|
* easy, thanks to the flag just_created.
|
|
*/
|
|
if (new_bfqq->wr_coeff == 1 && bfqq->wr_coeff > 1) {
|
|
new_bfqq->wr_coeff = bfqq->wr_coeff;
|
|
new_bfqq->wr_cur_max_time = bfqq->wr_cur_max_time;
|
|
new_bfqq->last_wr_start_finish = bfqq->last_wr_start_finish;
|
|
new_bfqq->wr_start_at_switch_to_srt =
|
|
bfqq->wr_start_at_switch_to_srt;
|
|
if (bfq_bfqq_busy(new_bfqq))
|
|
bfqd->wr_busy_queues++;
|
|
new_bfqq->entity.prio_changed = 1;
|
|
}
|
|
|
|
if (bfqq->wr_coeff > 1) { /* bfqq has given its wr to new_bfqq */
|
|
bfqq->wr_coeff = 1;
|
|
bfqq->entity.prio_changed = 1;
|
|
if (bfq_bfqq_busy(bfqq))
|
|
bfqd->wr_busy_queues--;
|
|
}
|
|
|
|
bfq_log_bfqq(bfqd, new_bfqq, "merge_bfqqs: wr_busy %d",
|
|
bfqd->wr_busy_queues);
|
|
|
|
/*
|
|
* Merge queues (that is, let bic redirect its requests to new_bfqq)
|
|
*/
|
|
bic_set_bfqq(bic, new_bfqq, 1);
|
|
bfq_mark_bfqq_coop(new_bfqq);
|
|
/*
|
|
* new_bfqq now belongs to at least two bics (it is a shared queue):
|
|
* set new_bfqq->bic to NULL. bfqq either:
|
|
* - does not belong to any bic any more, and hence bfqq->bic must
|
|
* be set to NULL, or
|
|
* - is a queue whose owning bics have already been redirected to a
|
|
* different queue, hence the queue is destined to not belong to
|
|
* any bic soon and bfqq->bic is already NULL (therefore the next
|
|
* assignment causes no harm).
|
|
*/
|
|
new_bfqq->bic = NULL;
|
|
/*
|
|
* If the queue is shared, the pid is the pid of one of the associated
|
|
* processes. Which pid depends on the exact sequence of merge events
|
|
* the queue underwent. So printing such a pid is useless and confusing
|
|
* because it reports a random pid between those of the associated
|
|
* processes.
|
|
* We mark such a queue with a pid -1, and then print SHARED instead of
|
|
* a pid in logging messages.
|
|
*/
|
|
new_bfqq->pid = -1;
|
|
bfqq->bic = NULL;
|
|
|
|
bfq_reassign_last_bfqq(bfqq, new_bfqq);
|
|
|
|
bfq_release_process_ref(bfqd, bfqq);
|
|
}
|
|
|
|
static bool bfq_allow_bio_merge(struct request_queue *q, struct request *rq,
|
|
struct bio *bio)
|
|
{
|
|
struct bfq_data *bfqd = q->elevator->elevator_data;
|
|
bool is_sync = op_is_sync(bio->bi_opf);
|
|
struct bfq_queue *bfqq = bfqd->bio_bfqq, *new_bfqq;
|
|
|
|
/*
|
|
* Disallow merge of a sync bio into an async request.
|
|
*/
|
|
if (is_sync && !rq_is_sync(rq))
|
|
return false;
|
|
|
|
/*
|
|
* Lookup the bfqq that this bio will be queued with. Allow
|
|
* merge only if rq is queued there.
|
|
*/
|
|
if (!bfqq)
|
|
return false;
|
|
|
|
/*
|
|
* We take advantage of this function to perform an early merge
|
|
* of the queues of possible cooperating processes.
|
|
*/
|
|
new_bfqq = bfq_setup_cooperator(bfqd, bfqq, bio, false, bfqd->bio_bic);
|
|
if (new_bfqq) {
|
|
/*
|
|
* bic still points to bfqq, then it has not yet been
|
|
* redirected to some other bfq_queue, and a queue
|
|
* merge between bfqq and new_bfqq can be safely
|
|
* fulfilled, i.e., bic can be redirected to new_bfqq
|
|
* and bfqq can be put.
|
|
*/
|
|
bfq_merge_bfqqs(bfqd, bfqd->bio_bic, bfqq,
|
|
new_bfqq);
|
|
/*
|
|
* If we get here, bio will be queued into new_queue,
|
|
* so use new_bfqq to decide whether bio and rq can be
|
|
* merged.
|
|
*/
|
|
bfqq = new_bfqq;
|
|
|
|
/*
|
|
* Change also bqfd->bio_bfqq, as
|
|
* bfqd->bio_bic now points to new_bfqq, and
|
|
* this function may be invoked again (and then may
|
|
* use again bqfd->bio_bfqq).
|
|
*/
|
|
bfqd->bio_bfqq = bfqq;
|
|
}
|
|
|
|
return bfqq == RQ_BFQQ(rq);
|
|
}
|
|
|
|
/*
|
|
* Set the maximum time for the in-service queue to consume its
|
|
* budget. This prevents seeky processes from lowering the throughput.
|
|
* In practice, a time-slice service scheme is used with seeky
|
|
* processes.
|
|
*/
|
|
static void bfq_set_budget_timeout(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
unsigned int timeout_coeff;
|
|
|
|
if (bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time)
|
|
timeout_coeff = 1;
|
|
else
|
|
timeout_coeff = bfqq->entity.weight / bfqq->entity.orig_weight;
|
|
|
|
bfqd->last_budget_start = ktime_get();
|
|
|
|
bfqq->budget_timeout = jiffies +
|
|
bfqd->bfq_timeout * timeout_coeff;
|
|
}
|
|
|
|
static void __bfq_set_in_service_queue(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
if (bfqq) {
|
|
bfq_clear_bfqq_fifo_expire(bfqq);
|
|
|
|
bfqd->budgets_assigned = (bfqd->budgets_assigned * 7 + 256) / 8;
|
|
|
|
if (time_is_before_jiffies(bfqq->last_wr_start_finish) &&
|
|
bfqq->wr_coeff > 1 &&
|
|
bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time &&
|
|
time_is_before_jiffies(bfqq->budget_timeout)) {
|
|
/*
|
|
* For soft real-time queues, move the start
|
|
* of the weight-raising period forward by the
|
|
* time the queue has not received any
|
|
* service. Otherwise, a relatively long
|
|
* service delay is likely to cause the
|
|
* weight-raising period of the queue to end,
|
|
* because of the short duration of the
|
|
* weight-raising period of a soft real-time
|
|
* queue. It is worth noting that this move
|
|
* is not so dangerous for the other queues,
|
|
* because soft real-time queues are not
|
|
* greedy.
|
|
*
|
|
* To not add a further variable, we use the
|
|
* overloaded field budget_timeout to
|
|
* determine for how long the queue has not
|
|
* received service, i.e., how much time has
|
|
* elapsed since the queue expired. However,
|
|
* this is a little imprecise, because
|
|
* budget_timeout is set to jiffies if bfqq
|
|
* not only expires, but also remains with no
|
|
* request.
|
|
*/
|
|
if (time_after(bfqq->budget_timeout,
|
|
bfqq->last_wr_start_finish))
|
|
bfqq->last_wr_start_finish +=
|
|
jiffies - bfqq->budget_timeout;
|
|
else
|
|
bfqq->last_wr_start_finish = jiffies;
|
|
}
|
|
|
|
bfq_set_budget_timeout(bfqd, bfqq);
|
|
bfq_log_bfqq(bfqd, bfqq,
|
|
"set_in_service_queue, cur-budget = %d",
|
|
bfqq->entity.budget);
|
|
}
|
|
|
|
bfqd->in_service_queue = bfqq;
|
|
bfqd->in_serv_last_pos = 0;
|
|
}
|
|
|
|
/*
|
|
* Get and set a new queue for service.
|
|
*/
|
|
static struct bfq_queue *bfq_set_in_service_queue(struct bfq_data *bfqd)
|
|
{
|
|
struct bfq_queue *bfqq = bfq_get_next_queue(bfqd);
|
|
|
|
__bfq_set_in_service_queue(bfqd, bfqq);
|
|
return bfqq;
|
|
}
|
|
|
|
static void bfq_arm_slice_timer(struct bfq_data *bfqd)
|
|
{
|
|
struct bfq_queue *bfqq = bfqd->in_service_queue;
|
|
u32 sl;
|
|
|
|
bfq_mark_bfqq_wait_request(bfqq);
|
|
|
|
/*
|
|
* We don't want to idle for seeks, but we do want to allow
|
|
* fair distribution of slice time for a process doing back-to-back
|
|
* seeks. So allow a little bit of time for him to submit a new rq.
|
|
*/
|
|
sl = bfqd->bfq_slice_idle;
|
|
/*
|
|
* Unless the queue is being weight-raised or the scenario is
|
|
* asymmetric, grant only minimum idle time if the queue
|
|
* is seeky. A long idling is preserved for a weight-raised
|
|
* queue, or, more in general, in an asymmetric scenario,
|
|
* because a long idling is needed for guaranteeing to a queue
|
|
* its reserved share of the throughput (in particular, it is
|
|
* needed if the queue has a higher weight than some other
|
|
* queue).
|
|
*/
|
|
if (BFQQ_SEEKY(bfqq) && bfqq->wr_coeff == 1 &&
|
|
!bfq_asymmetric_scenario(bfqd, bfqq))
|
|
sl = min_t(u64, sl, BFQ_MIN_TT);
|
|
else if (bfqq->wr_coeff > 1)
|
|
sl = max_t(u32, sl, 20ULL * NSEC_PER_MSEC);
|
|
|
|
bfqd->last_idling_start = ktime_get();
|
|
bfqd->last_idling_start_jiffies = jiffies;
|
|
|
|
hrtimer_start(&bfqd->idle_slice_timer, ns_to_ktime(sl),
|
|
HRTIMER_MODE_REL);
|
|
bfqg_stats_set_start_idle_time(bfqq_group(bfqq));
|
|
}
|
|
|
|
/*
|
|
* In autotuning mode, max_budget is dynamically recomputed as the
|
|
* amount of sectors transferred in timeout at the estimated peak
|
|
* rate. This enables BFQ to utilize a full timeslice with a full
|
|
* budget, even if the in-service queue is served at peak rate. And
|
|
* this maximises throughput with sequential workloads.
|
|
*/
|
|
static unsigned long bfq_calc_max_budget(struct bfq_data *bfqd)
|
|
{
|
|
return (u64)bfqd->peak_rate * USEC_PER_MSEC *
|
|
jiffies_to_msecs(bfqd->bfq_timeout)>>BFQ_RATE_SHIFT;
|
|
}
|
|
|
|
/*
|
|
* Update parameters related to throughput and responsiveness, as a
|
|
* function of the estimated peak rate. See comments on
|
|
* bfq_calc_max_budget(), and on the ref_wr_duration array.
|
|
*/
|
|
static void update_thr_responsiveness_params(struct bfq_data *bfqd)
|
|
{
|
|
if (bfqd->bfq_user_max_budget == 0) {
|
|
bfqd->bfq_max_budget =
|
|
bfq_calc_max_budget(bfqd);
|
|
bfq_log(bfqd, "new max_budget = %d", bfqd->bfq_max_budget);
|
|
}
|
|
}
|
|
|
|
static void bfq_reset_rate_computation(struct bfq_data *bfqd,
|
|
struct request *rq)
|
|
{
|
|
if (rq != NULL) { /* new rq dispatch now, reset accordingly */
|
|
bfqd->last_dispatch = bfqd->first_dispatch = ktime_get_ns();
|
|
bfqd->peak_rate_samples = 1;
|
|
bfqd->sequential_samples = 0;
|
|
bfqd->tot_sectors_dispatched = bfqd->last_rq_max_size =
|
|
blk_rq_sectors(rq);
|
|
} else /* no new rq dispatched, just reset the number of samples */
|
|
bfqd->peak_rate_samples = 0; /* full re-init on next disp. */
|
|
|
|
bfq_log(bfqd,
|
|
"reset_rate_computation at end, sample %u/%u tot_sects %llu",
|
|
bfqd->peak_rate_samples, bfqd->sequential_samples,
|
|
bfqd->tot_sectors_dispatched);
|
|
}
|
|
|
|
static void bfq_update_rate_reset(struct bfq_data *bfqd, struct request *rq)
|
|
{
|
|
u32 rate, weight, divisor;
|
|
|
|
/*
|
|
* For the convergence property to hold (see comments on
|
|
* bfq_update_peak_rate()) and for the assessment to be
|
|
* reliable, a minimum number of samples must be present, and
|
|
* a minimum amount of time must have elapsed. If not so, do
|
|
* not compute new rate. Just reset parameters, to get ready
|
|
* for a new evaluation attempt.
|
|
*/
|
|
if (bfqd->peak_rate_samples < BFQ_RATE_MIN_SAMPLES ||
|
|
bfqd->delta_from_first < BFQ_RATE_MIN_INTERVAL)
|
|
goto reset_computation;
|
|
|
|
/*
|
|
* If a new request completion has occurred after last
|
|
* dispatch, then, to approximate the rate at which requests
|
|
* have been served by the device, it is more precise to
|
|
* extend the observation interval to the last completion.
|
|
*/
|
|
bfqd->delta_from_first =
|
|
max_t(u64, bfqd->delta_from_first,
|
|
bfqd->last_completion - bfqd->first_dispatch);
|
|
|
|
/*
|
|
* Rate computed in sects/usec, and not sects/nsec, for
|
|
* precision issues.
|
|
*/
|
|
rate = div64_ul(bfqd->tot_sectors_dispatched<<BFQ_RATE_SHIFT,
|
|
div_u64(bfqd->delta_from_first, NSEC_PER_USEC));
|
|
|
|
/*
|
|
* Peak rate not updated if:
|
|
* - the percentage of sequential dispatches is below 3/4 of the
|
|
* total, and rate is below the current estimated peak rate
|
|
* - rate is unreasonably high (> 20M sectors/sec)
|
|
*/
|
|
if ((bfqd->sequential_samples < (3 * bfqd->peak_rate_samples)>>2 &&
|
|
rate <= bfqd->peak_rate) ||
|
|
rate > 20<<BFQ_RATE_SHIFT)
|
|
goto reset_computation;
|
|
|
|
/*
|
|
* We have to update the peak rate, at last! To this purpose,
|
|
* we use a low-pass filter. We compute the smoothing constant
|
|
* of the filter as a function of the 'weight' of the new
|
|
* measured rate.
|
|
*
|
|
* As can be seen in next formulas, we define this weight as a
|
|
* quantity proportional to how sequential the workload is,
|
|
* and to how long the observation time interval is.
|
|
*
|
|
* The weight runs from 0 to 8. The maximum value of the
|
|
* weight, 8, yields the minimum value for the smoothing
|
|
* constant. At this minimum value for the smoothing constant,
|
|
* the measured rate contributes for half of the next value of
|
|
* the estimated peak rate.
|
|
*
|
|
* So, the first step is to compute the weight as a function
|
|
* of how sequential the workload is. Note that the weight
|
|
* cannot reach 9, because bfqd->sequential_samples cannot
|
|
* become equal to bfqd->peak_rate_samples, which, in its
|
|
* turn, holds true because bfqd->sequential_samples is not
|
|
* incremented for the first sample.
|
|
*/
|
|
weight = (9 * bfqd->sequential_samples) / bfqd->peak_rate_samples;
|
|
|
|
/*
|
|
* Second step: further refine the weight as a function of the
|
|
* duration of the observation interval.
|
|
*/
|
|
weight = min_t(u32, 8,
|
|
div_u64(weight * bfqd->delta_from_first,
|
|
BFQ_RATE_REF_INTERVAL));
|
|
|
|
/*
|
|
* Divisor ranging from 10, for minimum weight, to 2, for
|
|
* maximum weight.
|
|
*/
|
|
divisor = 10 - weight;
|
|
|
|
/*
|
|
* Finally, update peak rate:
|
|
*
|
|
* peak_rate = peak_rate * (divisor-1) / divisor + rate / divisor
|
|
*/
|
|
bfqd->peak_rate *= divisor-1;
|
|
bfqd->peak_rate /= divisor;
|
|
rate /= divisor; /* smoothing constant alpha = 1/divisor */
|
|
|
|
bfqd->peak_rate += rate;
|
|
|
|
/*
|
|
* For a very slow device, bfqd->peak_rate can reach 0 (see
|
|
* the minimum representable values reported in the comments
|
|
* on BFQ_RATE_SHIFT). Push to 1 if this happens, to avoid
|
|
* divisions by zero where bfqd->peak_rate is used as a
|
|
* divisor.
|
|
*/
|
|
bfqd->peak_rate = max_t(u32, 1, bfqd->peak_rate);
|
|
|
|
update_thr_responsiveness_params(bfqd);
|
|
|
|
reset_computation:
|
|
bfq_reset_rate_computation(bfqd, rq);
|
|
}
|
|
|
|
/*
|
|
* Update the read/write peak rate (the main quantity used for
|
|
* auto-tuning, see update_thr_responsiveness_params()).
|
|
*
|
|
* It is not trivial to estimate the peak rate (correctly): because of
|
|
* the presence of sw and hw queues between the scheduler and the
|
|
* device components that finally serve I/O requests, it is hard to
|
|
* say exactly when a given dispatched request is served inside the
|
|
* device, and for how long. As a consequence, it is hard to know
|
|
* precisely at what rate a given set of requests is actually served
|
|
* by the device.
|
|
*
|
|
* On the opposite end, the dispatch time of any request is trivially
|
|
* available, and, from this piece of information, the "dispatch rate"
|
|
* of requests can be immediately computed. So, the idea in the next
|
|
* function is to use what is known, namely request dispatch times
|
|
* (plus, when useful, request completion times), to estimate what is
|
|
* unknown, namely in-device request service rate.
|
|
*
|
|
* The main issue is that, because of the above facts, the rate at
|
|
* which a certain set of requests is dispatched over a certain time
|
|
* interval can vary greatly with respect to the rate at which the
|
|
* same requests are then served. But, since the size of any
|
|
* intermediate queue is limited, and the service scheme is lossless
|
|
* (no request is silently dropped), the following obvious convergence
|
|
* property holds: the number of requests dispatched MUST become
|
|
* closer and closer to the number of requests completed as the
|
|
* observation interval grows. This is the key property used in
|
|
* the next function to estimate the peak service rate as a function
|
|
* of the observed dispatch rate. The function assumes to be invoked
|
|
* on every request dispatch.
|
|
*/
|
|
static void bfq_update_peak_rate(struct bfq_data *bfqd, struct request *rq)
|
|
{
|
|
u64 now_ns = ktime_get_ns();
|
|
|
|
if (bfqd->peak_rate_samples == 0) { /* first dispatch */
|
|
bfq_log(bfqd, "update_peak_rate: goto reset, samples %d",
|
|
bfqd->peak_rate_samples);
|
|
bfq_reset_rate_computation(bfqd, rq);
|
|
goto update_last_values; /* will add one sample */
|
|
}
|
|
|
|
/*
|
|
* Device idle for very long: the observation interval lasting
|
|
* up to this dispatch cannot be a valid observation interval
|
|
* for computing a new peak rate (similarly to the late-
|
|
* completion event in bfq_completed_request()). Go to
|
|
* update_rate_and_reset to have the following three steps
|
|
* taken:
|
|
* - close the observation interval at the last (previous)
|
|
* request dispatch or completion
|
|
* - compute rate, if possible, for that observation interval
|
|
* - start a new observation interval with this dispatch
|
|
*/
|
|
if (now_ns - bfqd->last_dispatch > 100*NSEC_PER_MSEC &&
|
|
bfqd->rq_in_driver == 0)
|
|
goto update_rate_and_reset;
|
|
|
|
/* Update sampling information */
|
|
bfqd->peak_rate_samples++;
|
|
|
|
if ((bfqd->rq_in_driver > 0 ||
|
|
now_ns - bfqd->last_completion < BFQ_MIN_TT)
|
|
&& !BFQ_RQ_SEEKY(bfqd, bfqd->last_position, rq))
|
|
bfqd->sequential_samples++;
|
|
|
|
bfqd->tot_sectors_dispatched += blk_rq_sectors(rq);
|
|
|
|
/* Reset max observed rq size every 32 dispatches */
|
|
if (likely(bfqd->peak_rate_samples % 32))
|
|
bfqd->last_rq_max_size =
|
|
max_t(u32, blk_rq_sectors(rq), bfqd->last_rq_max_size);
|
|
else
|
|
bfqd->last_rq_max_size = blk_rq_sectors(rq);
|
|
|
|
bfqd->delta_from_first = now_ns - bfqd->first_dispatch;
|
|
|
|
/* Target observation interval not yet reached, go on sampling */
|
|
if (bfqd->delta_from_first < BFQ_RATE_REF_INTERVAL)
|
|
goto update_last_values;
|
|
|
|
update_rate_and_reset:
|
|
bfq_update_rate_reset(bfqd, rq);
|
|
update_last_values:
|
|
bfqd->last_position = blk_rq_pos(rq) + blk_rq_sectors(rq);
|
|
if (RQ_BFQQ(rq) == bfqd->in_service_queue)
|
|
bfqd->in_serv_last_pos = bfqd->last_position;
|
|
bfqd->last_dispatch = now_ns;
|
|
}
|
|
|
|
/*
|
|
* Remove request from internal lists.
|
|
*/
|
|
static void bfq_dispatch_remove(struct request_queue *q, struct request *rq)
|
|
{
|
|
struct bfq_queue *bfqq = RQ_BFQQ(rq);
|
|
|
|
/*
|
|
* For consistency, the next instruction should have been
|
|
* executed after removing the request from the queue and
|
|
* dispatching it. We execute instead this instruction before
|
|
* bfq_remove_request() (and hence introduce a temporary
|
|
* inconsistency), for efficiency. In fact, should this
|
|
* dispatch occur for a non in-service bfqq, this anticipated
|
|
* increment prevents two counters related to bfqq->dispatched
|
|
* from risking to be, first, uselessly decremented, and then
|
|
* incremented again when the (new) value of bfqq->dispatched
|
|
* happens to be taken into account.
|
|
*/
|
|
bfqq->dispatched++;
|
|
bfq_update_peak_rate(q->elevator->elevator_data, rq);
|
|
|
|
bfq_remove_request(q, rq);
|
|
}
|
|
|
|
/*
|
|
* There is a case where idling does not have to be performed for
|
|
* throughput concerns, but to preserve the throughput share of
|
|
* the process associated with bfqq.
|
|
*
|
|
* To introduce this case, we can note that allowing the drive
|
|
* to enqueue more than one request at a time, and hence
|
|
* delegating de facto final scheduling decisions to the
|
|
* drive's internal scheduler, entails loss of control on the
|
|
* actual request service order. In particular, the critical
|
|
* situation is when requests from different processes happen
|
|
* to be present, at the same time, in the internal queue(s)
|
|
* of the drive. In such a situation, the drive, by deciding
|
|
* the service order of the internally-queued requests, does
|
|
* determine also the actual throughput distribution among
|
|
* these processes. But the drive typically has no notion or
|
|
* concern about per-process throughput distribution, and
|
|
* makes its decisions only on a per-request basis. Therefore,
|
|
* the service distribution enforced by the drive's internal
|
|
* scheduler is likely to coincide with the desired throughput
|
|
* distribution only in a completely symmetric, or favorably
|
|
* skewed scenario where:
|
|
* (i-a) each of these processes must get the same throughput as
|
|
* the others,
|
|
* (i-b) in case (i-a) does not hold, it holds that the process
|
|
* associated with bfqq must receive a lower or equal
|
|
* throughput than any of the other processes;
|
|
* (ii) the I/O of each process has the same properties, in
|
|
* terms of locality (sequential or random), direction
|
|
* (reads or writes), request sizes, greediness
|
|
* (from I/O-bound to sporadic), and so on;
|
|
|
|
* In fact, in such a scenario, the drive tends to treat the requests
|
|
* of each process in about the same way as the requests of the
|
|
* others, and thus to provide each of these processes with about the
|
|
* same throughput. This is exactly the desired throughput
|
|
* distribution if (i-a) holds, or, if (i-b) holds instead, this is an
|
|
* even more convenient distribution for (the process associated with)
|
|
* bfqq.
|
|
*
|
|
* In contrast, in any asymmetric or unfavorable scenario, device
|
|
* idling (I/O-dispatch plugging) is certainly needed to guarantee
|
|
* that bfqq receives its assigned fraction of the device throughput
|
|
* (see [1] for details).
|
|
*
|
|
* The problem is that idling may significantly reduce throughput with
|
|
* certain combinations of types of I/O and devices. An important
|
|
* example is sync random I/O on flash storage with command
|
|
* queueing. So, unless bfqq falls in cases where idling also boosts
|
|
* throughput, it is important to check conditions (i-a), i(-b) and
|
|
* (ii) accurately, so as to avoid idling when not strictly needed for
|
|
* service guarantees.
|
|
*
|
|
* Unfortunately, it is extremely difficult to thoroughly check
|
|
* condition (ii). And, in case there are active groups, it becomes
|
|
* very difficult to check conditions (i-a) and (i-b) too. In fact,
|
|
* if there are active groups, then, for conditions (i-a) or (i-b) to
|
|
* become false 'indirectly', it is enough that an active group
|
|
* contains more active processes or sub-groups than some other active
|
|
* group. More precisely, for conditions (i-a) or (i-b) to become
|
|
* false because of such a group, it is not even necessary that the
|
|
* group is (still) active: it is sufficient that, even if the group
|
|
* has become inactive, some of its descendant processes still have
|
|
* some request already dispatched but still waiting for
|
|
* completion. In fact, requests have still to be guaranteed their
|
|
* share of the throughput even after being dispatched. In this
|
|
* respect, it is easy to show that, if a group frequently becomes
|
|
* inactive while still having in-flight requests, and if, when this
|
|
* happens, the group is not considered in the calculation of whether
|
|
* the scenario is asymmetric, then the group may fail to be
|
|
* guaranteed its fair share of the throughput (basically because
|
|
* idling may not be performed for the descendant processes of the
|
|
* group, but it had to be). We address this issue with the following
|
|
* bi-modal behavior, implemented in the function
|
|
* bfq_asymmetric_scenario().
|
|
*
|
|
* If there are groups with requests waiting for completion
|
|
* (as commented above, some of these groups may even be
|
|
* already inactive), then the scenario is tagged as
|
|
* asymmetric, conservatively, without checking any of the
|
|
* conditions (i-a), (i-b) or (ii). So the device is idled for bfqq.
|
|
* This behavior matches also the fact that groups are created
|
|
* exactly if controlling I/O is a primary concern (to
|
|
* preserve bandwidth and latency guarantees).
|
|
*
|
|
* On the opposite end, if there are no groups with requests waiting
|
|
* for completion, then only conditions (i-a) and (i-b) are actually
|
|
* controlled, i.e., provided that conditions (i-a) or (i-b) holds,
|
|
* idling is not performed, regardless of whether condition (ii)
|
|
* holds. In other words, only if conditions (i-a) and (i-b) do not
|
|
* hold, then idling is allowed, and the device tends to be prevented
|
|
* from queueing many requests, possibly of several processes. Since
|
|
* there are no groups with requests waiting for completion, then, to
|
|
* control conditions (i-a) and (i-b) it is enough to check just
|
|
* whether all the queues with requests waiting for completion also
|
|
* have the same weight.
|
|
*
|
|
* Not checking condition (ii) evidently exposes bfqq to the
|
|
* risk of getting less throughput than its fair share.
|
|
* However, for queues with the same weight, a further
|
|
* mechanism, preemption, mitigates or even eliminates this
|
|
* problem. And it does so without consequences on overall
|
|
* throughput. This mechanism and its benefits are explained
|
|
* in the next three paragraphs.
|
|
*
|
|
* Even if a queue, say Q, is expired when it remains idle, Q
|
|
* can still preempt the new in-service queue if the next
|
|
* request of Q arrives soon (see the comments on
|
|
* bfq_bfqq_update_budg_for_activation). If all queues and
|
|
* groups have the same weight, this form of preemption,
|
|
* combined with the hole-recovery heuristic described in the
|
|
* comments on function bfq_bfqq_update_budg_for_activation,
|
|
* are enough to preserve a correct bandwidth distribution in
|
|
* the mid term, even without idling. In fact, even if not
|
|
* idling allows the internal queues of the device to contain
|
|
* many requests, and thus to reorder requests, we can rather
|
|
* safely assume that the internal scheduler still preserves a
|
|
* minimum of mid-term fairness.
|
|
*
|
|
* More precisely, this preemption-based, idleless approach
|
|
* provides fairness in terms of IOPS, and not sectors per
|
|
* second. This can be seen with a simple example. Suppose
|
|
* that there are two queues with the same weight, but that
|
|
* the first queue receives requests of 8 sectors, while the
|
|
* second queue receives requests of 1024 sectors. In
|
|
* addition, suppose that each of the two queues contains at
|
|
* most one request at a time, which implies that each queue
|
|
* always remains idle after it is served. Finally, after
|
|
* remaining idle, each queue receives very quickly a new
|
|
* request. It follows that the two queues are served
|
|
* alternatively, preempting each other if needed. This
|
|
* implies that, although both queues have the same weight,
|
|
* the queue with large requests receives a service that is
|
|
* 1024/8 times as high as the service received by the other
|
|
* queue.
|
|
*
|
|
* The motivation for using preemption instead of idling (for
|
|
* queues with the same weight) is that, by not idling,
|
|
* service guarantees are preserved (completely or at least in
|
|
* part) without minimally sacrificing throughput. And, if
|
|
* there is no active group, then the primary expectation for
|
|
* this device is probably a high throughput.
|
|
*
|
|
* We are now left only with explaining the two sub-conditions in the
|
|
* additional compound condition that is checked below for deciding
|
|
* whether the scenario is asymmetric. To explain the first
|
|
* sub-condition, we need to add that the function
|
|
* bfq_asymmetric_scenario checks the weights of only
|
|
* non-weight-raised queues, for efficiency reasons (see comments on
|
|
* bfq_weights_tree_add()). Then the fact that bfqq is weight-raised
|
|
* is checked explicitly here. More precisely, the compound condition
|
|
* below takes into account also the fact that, even if bfqq is being
|
|
* weight-raised, the scenario is still symmetric if all queues with
|
|
* requests waiting for completion happen to be
|
|
* weight-raised. Actually, we should be even more precise here, and
|
|
* differentiate between interactive weight raising and soft real-time
|
|
* weight raising.
|
|
*
|
|
* The second sub-condition checked in the compound condition is
|
|
* whether there is a fair amount of already in-flight I/O not
|
|
* belonging to bfqq. If so, I/O dispatching is to be plugged, for the
|
|
* following reason. The drive may decide to serve in-flight
|
|
* non-bfqq's I/O requests before bfqq's ones, thereby delaying the
|
|
* arrival of new I/O requests for bfqq (recall that bfqq is sync). If
|
|
* I/O-dispatching is not plugged, then, while bfqq remains empty, a
|
|
* basically uncontrolled amount of I/O from other queues may be
|
|
* dispatched too, possibly causing the service of bfqq's I/O to be
|
|
* delayed even longer in the drive. This problem gets more and more
|
|
* serious as the speed and the queue depth of the drive grow,
|
|
* because, as these two quantities grow, the probability to find no
|
|
* queue busy but many requests in flight grows too. By contrast,
|
|
* plugging I/O dispatching minimizes the delay induced by already
|
|
* in-flight I/O, and enables bfqq to recover the bandwidth it may
|
|
* lose because of this delay.
|
|
*
|
|
* As a side note, it is worth considering that the above
|
|
* device-idling countermeasures may however fail in the following
|
|
* unlucky scenario: if I/O-dispatch plugging is (correctly) disabled
|
|
* in a time period during which all symmetry sub-conditions hold, and
|
|
* therefore the device is allowed to enqueue many requests, but at
|
|
* some later point in time some sub-condition stops to hold, then it
|
|
* may become impossible to make requests be served in the desired
|
|
* order until all the requests already queued in the device have been
|
|
* served. The last sub-condition commented above somewhat mitigates
|
|
* this problem for weight-raised queues.
|
|
*
|
|
* However, as an additional mitigation for this problem, we preserve
|
|
* plugging for a special symmetric case that may suddenly turn into
|
|
* asymmetric: the case where only bfqq is busy. In this case, not
|
|
* expiring bfqq does not cause any harm to any other queues in terms
|
|
* of service guarantees. In contrast, it avoids the following unlucky
|
|
* sequence of events: (1) bfqq is expired, (2) a new queue with a
|
|
* lower weight than bfqq becomes busy (or more queues), (3) the new
|
|
* queue is served until a new request arrives for bfqq, (4) when bfqq
|
|
* is finally served, there are so many requests of the new queue in
|
|
* the drive that the pending requests for bfqq take a lot of time to
|
|
* be served. In particular, event (2) may case even already
|
|
* dispatched requests of bfqq to be delayed, inside the drive. So, to
|
|
* avoid this series of events, the scenario is preventively declared
|
|
* as asymmetric also if bfqq is the only busy queues
|
|
*/
|
|
static bool idling_needed_for_service_guarantees(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
int tot_busy_queues = bfq_tot_busy_queues(bfqd);
|
|
|
|
/* No point in idling for bfqq if it won't get requests any longer */
|
|
if (unlikely(!bfqq_process_refs(bfqq)))
|
|
return false;
|
|
|
|
return (bfqq->wr_coeff > 1 &&
|
|
(bfqd->wr_busy_queues <
|
|
tot_busy_queues ||
|
|
bfqd->rq_in_driver >=
|
|
bfqq->dispatched + 4)) ||
|
|
bfq_asymmetric_scenario(bfqd, bfqq) ||
|
|
tot_busy_queues == 1;
|
|
}
|
|
|
|
static bool __bfq_bfqq_expire(struct bfq_data *bfqd, struct bfq_queue *bfqq,
|
|
enum bfqq_expiration reason)
|
|
{
|
|
/*
|
|
* If this bfqq is shared between multiple processes, check
|
|
* to make sure that those processes are still issuing I/Os
|
|
* within the mean seek distance. If not, it may be time to
|
|
* break the queues apart again.
|
|
*/
|
|
if (bfq_bfqq_coop(bfqq) && BFQQ_SEEKY(bfqq))
|
|
bfq_mark_bfqq_split_coop(bfqq);
|
|
|
|
/*
|
|
* Consider queues with a higher finish virtual time than
|
|
* bfqq. If idling_needed_for_service_guarantees(bfqq) returns
|
|
* true, then bfqq's bandwidth would be violated if an
|
|
* uncontrolled amount of I/O from these queues were
|
|
* dispatched while bfqq is waiting for its new I/O to
|
|
* arrive. This is exactly what may happen if this is a forced
|
|
* expiration caused by a preemption attempt, and if bfqq is
|
|
* not re-scheduled. To prevent this from happening, re-queue
|
|
* bfqq if it needs I/O-dispatch plugging, even if it is
|
|
* empty. By doing so, bfqq is granted to be served before the
|
|
* above queues (provided that bfqq is of course eligible).
|
|
*/
|
|
if (RB_EMPTY_ROOT(&bfqq->sort_list) &&
|
|
!(reason == BFQQE_PREEMPTED &&
|
|
idling_needed_for_service_guarantees(bfqd, bfqq))) {
|
|
if (bfqq->dispatched == 0)
|
|
/*
|
|
* Overloading budget_timeout field to store
|
|
* the time at which the queue remains with no
|
|
* backlog and no outstanding request; used by
|
|
* the weight-raising mechanism.
|
|
*/
|
|
bfqq->budget_timeout = jiffies;
|
|
|
|
bfq_del_bfqq_busy(bfqd, bfqq, true);
|
|
} else {
|
|
bfq_requeue_bfqq(bfqd, bfqq, true);
|
|
/*
|
|
* Resort priority tree of potential close cooperators.
|
|
* See comments on bfq_pos_tree_add_move() for the unlikely().
|
|
*/
|
|
if (unlikely(!bfqd->nonrot_with_queueing &&
|
|
!RB_EMPTY_ROOT(&bfqq->sort_list)))
|
|
bfq_pos_tree_add_move(bfqd, bfqq);
|
|
}
|
|
|
|
/*
|
|
* All in-service entities must have been properly deactivated
|
|
* or requeued before executing the next function, which
|
|
* resets all in-service entities as no more in service. This
|
|
* may cause bfqq to be freed. If this happens, the next
|
|
* function returns true.
|
|
*/
|
|
return __bfq_bfqd_reset_in_service(bfqd);
|
|
}
|
|
|
|
/**
|
|
* __bfq_bfqq_recalc_budget - try to adapt the budget to the @bfqq behavior.
|
|
* @bfqd: device data.
|
|
* @bfqq: queue to update.
|
|
* @reason: reason for expiration.
|
|
*
|
|
* Handle the feedback on @bfqq budget at queue expiration.
|
|
* See the body for detailed comments.
|
|
*/
|
|
static void __bfq_bfqq_recalc_budget(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq,
|
|
enum bfqq_expiration reason)
|
|
{
|
|
struct request *next_rq;
|
|
int budget, min_budget;
|
|
|
|
min_budget = bfq_min_budget(bfqd);
|
|
|
|
if (bfqq->wr_coeff == 1)
|
|
budget = bfqq->max_budget;
|
|
else /*
|
|
* Use a constant, low budget for weight-raised queues,
|
|
* to help achieve a low latency. Keep it slightly higher
|
|
* than the minimum possible budget, to cause a little
|
|
* bit fewer expirations.
|
|
*/
|
|
budget = 2 * min_budget;
|
|
|
|
bfq_log_bfqq(bfqd, bfqq, "recalc_budg: last budg %d, budg left %d",
|
|
bfqq->entity.budget, bfq_bfqq_budget_left(bfqq));
|
|
bfq_log_bfqq(bfqd, bfqq, "recalc_budg: last max_budg %d, min budg %d",
|
|
budget, bfq_min_budget(bfqd));
|
|
bfq_log_bfqq(bfqd, bfqq, "recalc_budg: sync %d, seeky %d",
|
|
bfq_bfqq_sync(bfqq), BFQQ_SEEKY(bfqd->in_service_queue));
|
|
|
|
if (bfq_bfqq_sync(bfqq) && bfqq->wr_coeff == 1) {
|
|
switch (reason) {
|
|
/*
|
|
* Caveat: in all the following cases we trade latency
|
|
* for throughput.
|
|
*/
|
|
case BFQQE_TOO_IDLE:
|
|
/*
|
|
* This is the only case where we may reduce
|
|
* the budget: if there is no request of the
|
|
* process still waiting for completion, then
|
|
* we assume (tentatively) that the timer has
|
|
* expired because the batch of requests of
|
|
* the process could have been served with a
|
|
* smaller budget. Hence, betting that
|
|
* process will behave in the same way when it
|
|
* becomes backlogged again, we reduce its
|
|
* next budget. As long as we guess right,
|
|
* this budget cut reduces the latency
|
|
* experienced by the process.
|
|
*
|
|
* However, if there are still outstanding
|
|
* requests, then the process may have not yet
|
|
* issued its next request just because it is
|
|
* still waiting for the completion of some of
|
|
* the still outstanding ones. So in this
|
|
* subcase we do not reduce its budget, on the
|
|
* contrary we increase it to possibly boost
|
|
* the throughput, as discussed in the
|
|
* comments to the BUDGET_TIMEOUT case.
|
|
*/
|
|
if (bfqq->dispatched > 0) /* still outstanding reqs */
|
|
budget = min(budget * 2, bfqd->bfq_max_budget);
|
|
else {
|
|
if (budget > 5 * min_budget)
|
|
budget -= 4 * min_budget;
|
|
else
|
|
budget = min_budget;
|
|
}
|
|
break;
|
|
case BFQQE_BUDGET_TIMEOUT:
|
|
/*
|
|
* We double the budget here because it gives
|
|
* the chance to boost the throughput if this
|
|
* is not a seeky process (and has bumped into
|
|
* this timeout because of, e.g., ZBR).
|
|
*/
|
|
budget = min(budget * 2, bfqd->bfq_max_budget);
|
|
break;
|
|
case BFQQE_BUDGET_EXHAUSTED:
|
|
/*
|
|
* The process still has backlog, and did not
|
|
* let either the budget timeout or the disk
|
|
* idling timeout expire. Hence it is not
|
|
* seeky, has a short thinktime and may be
|
|
* happy with a higher budget too. So
|
|
* definitely increase the budget of this good
|
|
* candidate to boost the disk throughput.
|
|
*/
|
|
budget = min(budget * 4, bfqd->bfq_max_budget);
|
|
break;
|
|
case BFQQE_NO_MORE_REQUESTS:
|
|
/*
|
|
* For queues that expire for this reason, it
|
|
* is particularly important to keep the
|
|
* budget close to the actual service they
|
|
* need. Doing so reduces the timestamp
|
|
* misalignment problem described in the
|
|
* comments in the body of
|
|
* __bfq_activate_entity. In fact, suppose
|
|
* that a queue systematically expires for
|
|
* BFQQE_NO_MORE_REQUESTS and presents a
|
|
* new request in time to enjoy timestamp
|
|
* back-shifting. The larger the budget of the
|
|
* queue is with respect to the service the
|
|
* queue actually requests in each service
|
|
* slot, the more times the queue can be
|
|
* reactivated with the same virtual finish
|
|
* time. It follows that, even if this finish
|
|
* time is pushed to the system virtual time
|
|
* to reduce the consequent timestamp
|
|
* misalignment, the queue unjustly enjoys for
|
|
* many re-activations a lower finish time
|
|
* than all newly activated queues.
|
|
*
|
|
* The service needed by bfqq is measured
|
|
* quite precisely by bfqq->entity.service.
|
|
* Since bfqq does not enjoy device idling,
|
|
* bfqq->entity.service is equal to the number
|
|
* of sectors that the process associated with
|
|
* bfqq requested to read/write before waiting
|
|
* for request completions, or blocking for
|
|
* other reasons.
|
|
*/
|
|
budget = max_t(int, bfqq->entity.service, min_budget);
|
|
break;
|
|
default:
|
|
return;
|
|
}
|
|
} else if (!bfq_bfqq_sync(bfqq)) {
|
|
/*
|
|
* Async queues get always the maximum possible
|
|
* budget, as for them we do not care about latency
|
|
* (in addition, their ability to dispatch is limited
|
|
* by the charging factor).
|
|
*/
|
|
budget = bfqd->bfq_max_budget;
|
|
}
|
|
|
|
bfqq->max_budget = budget;
|
|
|
|
if (bfqd->budgets_assigned >= bfq_stats_min_budgets &&
|
|
!bfqd->bfq_user_max_budget)
|
|
bfqq->max_budget = min(bfqq->max_budget, bfqd->bfq_max_budget);
|
|
|
|
/*
|
|
* If there is still backlog, then assign a new budget, making
|
|
* sure that it is large enough for the next request. Since
|
|
* the finish time of bfqq must be kept in sync with the
|
|
* budget, be sure to call __bfq_bfqq_expire() *after* this
|
|
* update.
|
|
*
|
|
* If there is no backlog, then no need to update the budget;
|
|
* it will be updated on the arrival of a new request.
|
|
*/
|
|
next_rq = bfqq->next_rq;
|
|
if (next_rq)
|
|
bfqq->entity.budget = max_t(unsigned long, bfqq->max_budget,
|
|
bfq_serv_to_charge(next_rq, bfqq));
|
|
|
|
bfq_log_bfqq(bfqd, bfqq, "head sect: %u, new budget %d",
|
|
next_rq ? blk_rq_sectors(next_rq) : 0,
|
|
bfqq->entity.budget);
|
|
}
|
|
|
|
/*
|
|
* Return true if the process associated with bfqq is "slow". The slow
|
|
* flag is used, in addition to the budget timeout, to reduce the
|
|
* amount of service provided to seeky processes, and thus reduce
|
|
* their chances to lower the throughput. More details in the comments
|
|
* on the function bfq_bfqq_expire().
|
|
*
|
|
* An important observation is in order: as discussed in the comments
|
|
* on the function bfq_update_peak_rate(), with devices with internal
|
|
* queues, it is hard if ever possible to know when and for how long
|
|
* an I/O request is processed by the device (apart from the trivial
|
|
* I/O pattern where a new request is dispatched only after the
|
|
* previous one has been completed). This makes it hard to evaluate
|
|
* the real rate at which the I/O requests of each bfq_queue are
|
|
* served. In fact, for an I/O scheduler like BFQ, serving a
|
|
* bfq_queue means just dispatching its requests during its service
|
|
* slot (i.e., until the budget of the queue is exhausted, or the
|
|
* queue remains idle, or, finally, a timeout fires). But, during the
|
|
* service slot of a bfq_queue, around 100 ms at most, the device may
|
|
* be even still processing requests of bfq_queues served in previous
|
|
* service slots. On the opposite end, the requests of the in-service
|
|
* bfq_queue may be completed after the service slot of the queue
|
|
* finishes.
|
|
*
|
|
* Anyway, unless more sophisticated solutions are used
|
|
* (where possible), the sum of the sizes of the requests dispatched
|
|
* during the service slot of a bfq_queue is probably the only
|
|
* approximation available for the service received by the bfq_queue
|
|
* during its service slot. And this sum is the quantity used in this
|
|
* function to evaluate the I/O speed of a process.
|
|
*/
|
|
static bool bfq_bfqq_is_slow(struct bfq_data *bfqd, struct bfq_queue *bfqq,
|
|
bool compensate, enum bfqq_expiration reason,
|
|
unsigned long *delta_ms)
|
|
{
|
|
ktime_t delta_ktime;
|
|
u32 delta_usecs;
|
|
bool slow = BFQQ_SEEKY(bfqq); /* if delta too short, use seekyness */
|
|
|
|
if (!bfq_bfqq_sync(bfqq))
|
|
return false;
|
|
|
|
if (compensate)
|
|
delta_ktime = bfqd->last_idling_start;
|
|
else
|
|
delta_ktime = ktime_get();
|
|
delta_ktime = ktime_sub(delta_ktime, bfqd->last_budget_start);
|
|
delta_usecs = ktime_to_us(delta_ktime);
|
|
|
|
/* don't use too short time intervals */
|
|
if (delta_usecs < 1000) {
|
|
if (blk_queue_nonrot(bfqd->queue))
|
|
/*
|
|
* give same worst-case guarantees as idling
|
|
* for seeky
|
|
*/
|
|
*delta_ms = BFQ_MIN_TT / NSEC_PER_MSEC;
|
|
else /* charge at least one seek */
|
|
*delta_ms = bfq_slice_idle / NSEC_PER_MSEC;
|
|
|
|
return slow;
|
|
}
|
|
|
|
*delta_ms = delta_usecs / USEC_PER_MSEC;
|
|
|
|
/*
|
|
* Use only long (> 20ms) intervals to filter out excessive
|
|
* spikes in service rate estimation.
|
|
*/
|
|
if (delta_usecs > 20000) {
|
|
/*
|
|
* Caveat for rotational devices: processes doing I/O
|
|
* in the slower disk zones tend to be slow(er) even
|
|
* if not seeky. In this respect, the estimated peak
|
|
* rate is likely to be an average over the disk
|
|
* surface. Accordingly, to not be too harsh with
|
|
* unlucky processes, a process is deemed slow only if
|
|
* its rate has been lower than half of the estimated
|
|
* peak rate.
|
|
*/
|
|
slow = bfqq->entity.service < bfqd->bfq_max_budget / 2;
|
|
}
|
|
|
|
bfq_log_bfqq(bfqd, bfqq, "bfq_bfqq_is_slow: slow %d", slow);
|
|
|
|
return slow;
|
|
}
|
|
|
|
/*
|
|
* To be deemed as soft real-time, an application must meet two
|
|
* requirements. First, the application must not require an average
|
|
* bandwidth higher than the approximate bandwidth required to playback or
|
|
* record a compressed high-definition video.
|
|
* The next function is invoked on the completion of the last request of a
|
|
* batch, to compute the next-start time instant, soft_rt_next_start, such
|
|
* that, if the next request of the application does not arrive before
|
|
* soft_rt_next_start, then the above requirement on the bandwidth is met.
|
|
*
|
|
* The second requirement is that the request pattern of the application is
|
|
* isochronous, i.e., that, after issuing a request or a batch of requests,
|
|
* the application stops issuing new requests until all its pending requests
|
|
* have been completed. After that, the application may issue a new batch,
|
|
* and so on.
|
|
* For this reason the next function is invoked to compute
|
|
* soft_rt_next_start only for applications that meet this requirement,
|
|
* whereas soft_rt_next_start is set to infinity for applications that do
|
|
* not.
|
|
*
|
|
* Unfortunately, even a greedy (i.e., I/O-bound) application may
|
|
* happen to meet, occasionally or systematically, both the above
|
|
* bandwidth and isochrony requirements. This may happen at least in
|
|
* the following circumstances. First, if the CPU load is high. The
|
|
* application may stop issuing requests while the CPUs are busy
|
|
* serving other processes, then restart, then stop again for a while,
|
|
* and so on. The other circumstances are related to the storage
|
|
* device: the storage device is highly loaded or reaches a low-enough
|
|
* throughput with the I/O of the application (e.g., because the I/O
|
|
* is random and/or the device is slow). In all these cases, the
|
|
* I/O of the application may be simply slowed down enough to meet
|
|
* the bandwidth and isochrony requirements. To reduce the probability
|
|
* that greedy applications are deemed as soft real-time in these
|
|
* corner cases, a further rule is used in the computation of
|
|
* soft_rt_next_start: the return value of this function is forced to
|
|
* be higher than the maximum between the following two quantities.
|
|
*
|
|
* (a) Current time plus: (1) the maximum time for which the arrival
|
|
* of a request is waited for when a sync queue becomes idle,
|
|
* namely bfqd->bfq_slice_idle, and (2) a few extra jiffies. We
|
|
* postpone for a moment the reason for adding a few extra
|
|
* jiffies; we get back to it after next item (b). Lower-bounding
|
|
* the return value of this function with the current time plus
|
|
* bfqd->bfq_slice_idle tends to filter out greedy applications,
|
|
* because the latter issue their next request as soon as possible
|
|
* after the last one has been completed. In contrast, a soft
|
|
* real-time application spends some time processing data, after a
|
|
* batch of its requests has been completed.
|
|
*
|
|
* (b) Current value of bfqq->soft_rt_next_start. As pointed out
|
|
* above, greedy applications may happen to meet both the
|
|
* bandwidth and isochrony requirements under heavy CPU or
|
|
* storage-device load. In more detail, in these scenarios, these
|
|
* applications happen, only for limited time periods, to do I/O
|
|
* slowly enough to meet all the requirements described so far,
|
|
* including the filtering in above item (a). These slow-speed
|
|
* time intervals are usually interspersed between other time
|
|
* intervals during which these applications do I/O at a very high
|
|
* speed. Fortunately, exactly because of the high speed of the
|
|
* I/O in the high-speed intervals, the values returned by this
|
|
* function happen to be so high, near the end of any such
|
|
* high-speed interval, to be likely to fall *after* the end of
|
|
* the low-speed time interval that follows. These high values are
|
|
* stored in bfqq->soft_rt_next_start after each invocation of
|
|
* this function. As a consequence, if the last value of
|
|
* bfqq->soft_rt_next_start is constantly used to lower-bound the
|
|
* next value that this function may return, then, from the very
|
|
* beginning of a low-speed interval, bfqq->soft_rt_next_start is
|
|
* likely to be constantly kept so high that any I/O request
|
|
* issued during the low-speed interval is considered as arriving
|
|
* to soon for the application to be deemed as soft
|
|
* real-time. Then, in the high-speed interval that follows, the
|
|
* application will not be deemed as soft real-time, just because
|
|
* it will do I/O at a high speed. And so on.
|
|
*
|
|
* Getting back to the filtering in item (a), in the following two
|
|
* cases this filtering might be easily passed by a greedy
|
|
* application, if the reference quantity was just
|
|
* bfqd->bfq_slice_idle:
|
|
* 1) HZ is so low that the duration of a jiffy is comparable to or
|
|
* higher than bfqd->bfq_slice_idle. This happens, e.g., on slow
|
|
* devices with HZ=100. The time granularity may be so coarse
|
|
* that the approximation, in jiffies, of bfqd->bfq_slice_idle
|
|
* is rather lower than the exact value.
|
|
* 2) jiffies, instead of increasing at a constant rate, may stop increasing
|
|
* for a while, then suddenly 'jump' by several units to recover the lost
|
|
* increments. This seems to happen, e.g., inside virtual machines.
|
|
* To address this issue, in the filtering in (a) we do not use as a
|
|
* reference time interval just bfqd->bfq_slice_idle, but
|
|
* bfqd->bfq_slice_idle plus a few jiffies. In particular, we add the
|
|
* minimum number of jiffies for which the filter seems to be quite
|
|
* precise also in embedded systems and KVM/QEMU virtual machines.
|
|
*/
|
|
static unsigned long bfq_bfqq_softrt_next_start(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
return max3(bfqq->soft_rt_next_start,
|
|
bfqq->last_idle_bklogged +
|
|
HZ * bfqq->service_from_backlogged /
|
|
bfqd->bfq_wr_max_softrt_rate,
|
|
jiffies + nsecs_to_jiffies(bfqq->bfqd->bfq_slice_idle) + 4);
|
|
}
|
|
|
|
/**
|
|
* bfq_bfqq_expire - expire a queue.
|
|
* @bfqd: device owning the queue.
|
|
* @bfqq: the queue to expire.
|
|
* @compensate: if true, compensate for the time spent idling.
|
|
* @reason: the reason causing the expiration.
|
|
*
|
|
* If the process associated with bfqq does slow I/O (e.g., because it
|
|
* issues random requests), we charge bfqq with the time it has been
|
|
* in service instead of the service it has received (see
|
|
* bfq_bfqq_charge_time for details on how this goal is achieved). As
|
|
* a consequence, bfqq will typically get higher timestamps upon
|
|
* reactivation, and hence it will be rescheduled as if it had
|
|
* received more service than what it has actually received. In the
|
|
* end, bfqq receives less service in proportion to how slowly its
|
|
* associated process consumes its budgets (and hence how seriously it
|
|
* tends to lower the throughput). In addition, this time-charging
|
|
* strategy guarantees time fairness among slow processes. In
|
|
* contrast, if the process associated with bfqq is not slow, we
|
|
* charge bfqq exactly with the service it has received.
|
|
*
|
|
* Charging time to the first type of queues and the exact service to
|
|
* the other has the effect of using the WF2Q+ policy to schedule the
|
|
* former on a timeslice basis, without violating service domain
|
|
* guarantees among the latter.
|
|
*/
|
|
void bfq_bfqq_expire(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq,
|
|
bool compensate,
|
|
enum bfqq_expiration reason)
|
|
{
|
|
bool slow;
|
|
unsigned long delta = 0;
|
|
struct bfq_entity *entity = &bfqq->entity;
|
|
|
|
/*
|
|
* Check whether the process is slow (see bfq_bfqq_is_slow).
|
|
*/
|
|
slow = bfq_bfqq_is_slow(bfqd, bfqq, compensate, reason, &delta);
|
|
|
|
/*
|
|
* As above explained, charge slow (typically seeky) and
|
|
* timed-out queues with the time and not the service
|
|
* received, to favor sequential workloads.
|
|
*
|
|
* Processes doing I/O in the slower disk zones will tend to
|
|
* be slow(er) even if not seeky. Therefore, since the
|
|
* estimated peak rate is actually an average over the disk
|
|
* surface, these processes may timeout just for bad luck. To
|
|
* avoid punishing them, do not charge time to processes that
|
|
* succeeded in consuming at least 2/3 of their budget. This
|
|
* allows BFQ to preserve enough elasticity to still perform
|
|
* bandwidth, and not time, distribution with little unlucky
|
|
* or quasi-sequential processes.
|
|
*/
|
|
if (bfqq->wr_coeff == 1 &&
|
|
(slow ||
|
|
(reason == BFQQE_BUDGET_TIMEOUT &&
|
|
bfq_bfqq_budget_left(bfqq) >= entity->budget / 3)))
|
|
bfq_bfqq_charge_time(bfqd, bfqq, delta);
|
|
|
|
if (bfqd->low_latency && bfqq->wr_coeff == 1)
|
|
bfqq->last_wr_start_finish = jiffies;
|
|
|
|
if (bfqd->low_latency && bfqd->bfq_wr_max_softrt_rate > 0 &&
|
|
RB_EMPTY_ROOT(&bfqq->sort_list)) {
|
|
/*
|
|
* If we get here, and there are no outstanding
|
|
* requests, then the request pattern is isochronous
|
|
* (see the comments on the function
|
|
* bfq_bfqq_softrt_next_start()). Therefore we can
|
|
* compute soft_rt_next_start.
|
|
*
|
|
* If, instead, the queue still has outstanding
|
|
* requests, then we have to wait for the completion
|
|
* of all the outstanding requests to discover whether
|
|
* the request pattern is actually isochronous.
|
|
*/
|
|
if (bfqq->dispatched == 0)
|
|
bfqq->soft_rt_next_start =
|
|
bfq_bfqq_softrt_next_start(bfqd, bfqq);
|
|
else if (bfqq->dispatched > 0) {
|
|
/*
|
|
* Schedule an update of soft_rt_next_start to when
|
|
* the task may be discovered to be isochronous.
|
|
*/
|
|
bfq_mark_bfqq_softrt_update(bfqq);
|
|
}
|
|
}
|
|
|
|
bfq_log_bfqq(bfqd, bfqq,
|
|
"expire (%d, slow %d, num_disp %d, short_ttime %d)", reason,
|
|
slow, bfqq->dispatched, bfq_bfqq_has_short_ttime(bfqq));
|
|
|
|
/*
|
|
* bfqq expired, so no total service time needs to be computed
|
|
* any longer: reset state machine for measuring total service
|
|
* times.
|
|
*/
|
|
bfqd->rqs_injected = bfqd->wait_dispatch = false;
|
|
bfqd->waited_rq = NULL;
|
|
|
|
/*
|
|
* Increase, decrease or leave budget unchanged according to
|
|
* reason.
|
|
*/
|
|
__bfq_bfqq_recalc_budget(bfqd, bfqq, reason);
|
|
if (__bfq_bfqq_expire(bfqd, bfqq, reason))
|
|
/* bfqq is gone, no more actions on it */
|
|
return;
|
|
|
|
/* mark bfqq as waiting a request only if a bic still points to it */
|
|
if (!bfq_bfqq_busy(bfqq) &&
|
|
reason != BFQQE_BUDGET_TIMEOUT &&
|
|
reason != BFQQE_BUDGET_EXHAUSTED) {
|
|
bfq_mark_bfqq_non_blocking_wait_rq(bfqq);
|
|
/*
|
|
* Not setting service to 0, because, if the next rq
|
|
* arrives in time, the queue will go on receiving
|
|
* service with this same budget (as if it never expired)
|
|
*/
|
|
} else
|
|
entity->service = 0;
|
|
|
|
/*
|
|
* Reset the received-service counter for every parent entity.
|
|
* Differently from what happens with bfqq->entity.service,
|
|
* the resetting of this counter never needs to be postponed
|
|
* for parent entities. In fact, in case bfqq may have a
|
|
* chance to go on being served using the last, partially
|
|
* consumed budget, bfqq->entity.service needs to be kept,
|
|
* because if bfqq then actually goes on being served using
|
|
* the same budget, the last value of bfqq->entity.service is
|
|
* needed to properly decrement bfqq->entity.budget by the
|
|
* portion already consumed. In contrast, it is not necessary
|
|
* to keep entity->service for parent entities too, because
|
|
* the bubble up of the new value of bfqq->entity.budget will
|
|
* make sure that the budgets of parent entities are correct,
|
|
* even in case bfqq and thus parent entities go on receiving
|
|
* service with the same budget.
|
|
*/
|
|
entity = entity->parent;
|
|
for_each_entity(entity)
|
|
entity->service = 0;
|
|
}
|
|
|
|
/*
|
|
* Budget timeout is not implemented through a dedicated timer, but
|
|
* just checked on request arrivals and completions, as well as on
|
|
* idle timer expirations.
|
|
*/
|
|
static bool bfq_bfqq_budget_timeout(struct bfq_queue *bfqq)
|
|
{
|
|
return time_is_before_eq_jiffies(bfqq->budget_timeout);
|
|
}
|
|
|
|
/*
|
|
* If we expire a queue that is actively waiting (i.e., with the
|
|
* device idled) for the arrival of a new request, then we may incur
|
|
* the timestamp misalignment problem described in the body of the
|
|
* function __bfq_activate_entity. Hence we return true only if this
|
|
* condition does not hold, or if the queue is slow enough to deserve
|
|
* only to be kicked off for preserving a high throughput.
|
|
*/
|
|
static bool bfq_may_expire_for_budg_timeout(struct bfq_queue *bfqq)
|
|
{
|
|
bfq_log_bfqq(bfqq->bfqd, bfqq,
|
|
"may_budget_timeout: wait_request %d left %d timeout %d",
|
|
bfq_bfqq_wait_request(bfqq),
|
|
bfq_bfqq_budget_left(bfqq) >= bfqq->entity.budget / 3,
|
|
bfq_bfqq_budget_timeout(bfqq));
|
|
|
|
return (!bfq_bfqq_wait_request(bfqq) ||
|
|
bfq_bfqq_budget_left(bfqq) >= bfqq->entity.budget / 3)
|
|
&&
|
|
bfq_bfqq_budget_timeout(bfqq);
|
|
}
|
|
|
|
static bool idling_boosts_thr_without_issues(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
bool rot_without_queueing =
|
|
!blk_queue_nonrot(bfqd->queue) && !bfqd->hw_tag,
|
|
bfqq_sequential_and_IO_bound,
|
|
idling_boosts_thr;
|
|
|
|
/* No point in idling for bfqq if it won't get requests any longer */
|
|
if (unlikely(!bfqq_process_refs(bfqq)))
|
|
return false;
|
|
|
|
bfqq_sequential_and_IO_bound = !BFQQ_SEEKY(bfqq) &&
|
|
bfq_bfqq_IO_bound(bfqq) && bfq_bfqq_has_short_ttime(bfqq);
|
|
|
|
/*
|
|
* The next variable takes into account the cases where idling
|
|
* boosts the throughput.
|
|
*
|
|
* The value of the variable is computed considering, first, that
|
|
* idling is virtually always beneficial for the throughput if:
|
|
* (a) the device is not NCQ-capable and rotational, or
|
|
* (b) regardless of the presence of NCQ, the device is rotational and
|
|
* the request pattern for bfqq is I/O-bound and sequential, or
|
|
* (c) regardless of whether it is rotational, the device is
|
|
* not NCQ-capable and the request pattern for bfqq is
|
|
* I/O-bound and sequential.
|
|
*
|
|
* Secondly, and in contrast to the above item (b), idling an
|
|
* NCQ-capable flash-based device would not boost the
|
|
* throughput even with sequential I/O; rather it would lower
|
|
* the throughput in proportion to how fast the device
|
|
* is. Accordingly, the next variable is true if any of the
|
|
* above conditions (a), (b) or (c) is true, and, in
|
|
* particular, happens to be false if bfqd is an NCQ-capable
|
|
* flash-based device.
|
|
*/
|
|
idling_boosts_thr = rot_without_queueing ||
|
|
((!blk_queue_nonrot(bfqd->queue) || !bfqd->hw_tag) &&
|
|
bfqq_sequential_and_IO_bound);
|
|
|
|
/*
|
|
* The return value of this function is equal to that of
|
|
* idling_boosts_thr, unless a special case holds. In this
|
|
* special case, described below, idling may cause problems to
|
|
* weight-raised queues.
|
|
*
|
|
* When the request pool is saturated (e.g., in the presence
|
|
* of write hogs), if the processes associated with
|
|
* non-weight-raised queues ask for requests at a lower rate,
|
|
* then processes associated with weight-raised queues have a
|
|
* higher probability to get a request from the pool
|
|
* immediately (or at least soon) when they need one. Thus
|
|
* they have a higher probability to actually get a fraction
|
|
* of the device throughput proportional to their high
|
|
* weight. This is especially true with NCQ-capable drives,
|
|
* which enqueue several requests in advance, and further
|
|
* reorder internally-queued requests.
|
|
*
|
|
* For this reason, we force to false the return value if
|
|
* there are weight-raised busy queues. In this case, and if
|
|
* bfqq is not weight-raised, this guarantees that the device
|
|
* is not idled for bfqq (if, instead, bfqq is weight-raised,
|
|
* then idling will be guaranteed by another variable, see
|
|
* below). Combined with the timestamping rules of BFQ (see
|
|
* [1] for details), this behavior causes bfqq, and hence any
|
|
* sync non-weight-raised queue, to get a lower number of
|
|
* requests served, and thus to ask for a lower number of
|
|
* requests from the request pool, before the busy
|
|
* weight-raised queues get served again. This often mitigates
|
|
* starvation problems in the presence of heavy write
|
|
* workloads and NCQ, thereby guaranteeing a higher
|
|
* application and system responsiveness in these hostile
|
|
* scenarios.
|
|
*/
|
|
return idling_boosts_thr &&
|
|
bfqd->wr_busy_queues == 0;
|
|
}
|
|
|
|
/*
|
|
* For a queue that becomes empty, device idling is allowed only if
|
|
* this function returns true for that queue. As a consequence, since
|
|
* device idling plays a critical role for both throughput boosting
|
|
* and service guarantees, the return value of this function plays a
|
|
* critical role as well.
|
|
*
|
|
* In a nutshell, this function returns true only if idling is
|
|
* beneficial for throughput or, even if detrimental for throughput,
|
|
* idling is however necessary to preserve service guarantees (low
|
|
* latency, desired throughput distribution, ...). In particular, on
|
|
* NCQ-capable devices, this function tries to return false, so as to
|
|
* help keep the drives' internal queues full, whenever this helps the
|
|
* device boost the throughput without causing any service-guarantee
|
|
* issue.
|
|
*
|
|
* Most of the issues taken into account to get the return value of
|
|
* this function are not trivial. We discuss these issues in the two
|
|
* functions providing the main pieces of information needed by this
|
|
* function.
|
|
*/
|
|
static bool bfq_better_to_idle(struct bfq_queue *bfqq)
|
|
{
|
|
struct bfq_data *bfqd = bfqq->bfqd;
|
|
bool idling_boosts_thr_with_no_issue, idling_needed_for_service_guar;
|
|
|
|
/* No point in idling for bfqq if it won't get requests any longer */
|
|
if (unlikely(!bfqq_process_refs(bfqq)))
|
|
return false;
|
|
|
|
if (unlikely(bfqd->strict_guarantees))
|
|
return true;
|
|
|
|
/*
|
|
* Idling is performed only if slice_idle > 0. In addition, we
|
|
* do not idle if
|
|
* (a) bfqq is async
|
|
* (b) bfqq is in the idle io prio class: in this case we do
|
|
* not idle because we want to minimize the bandwidth that
|
|
* queues in this class can steal to higher-priority queues
|
|
*/
|
|
if (bfqd->bfq_slice_idle == 0 || !bfq_bfqq_sync(bfqq) ||
|
|
bfq_class_idle(bfqq))
|
|
return false;
|
|
|
|
idling_boosts_thr_with_no_issue =
|
|
idling_boosts_thr_without_issues(bfqd, bfqq);
|
|
|
|
idling_needed_for_service_guar =
|
|
idling_needed_for_service_guarantees(bfqd, bfqq);
|
|
|
|
/*
|
|
* We have now the two components we need to compute the
|
|
* return value of the function, which is true only if idling
|
|
* either boosts the throughput (without issues), or is
|
|
* necessary to preserve service guarantees.
|
|
*/
|
|
return idling_boosts_thr_with_no_issue ||
|
|
idling_needed_for_service_guar;
|
|
}
|
|
|
|
/*
|
|
* If the in-service queue is empty but the function bfq_better_to_idle
|
|
* returns true, then:
|
|
* 1) the queue must remain in service and cannot be expired, and
|
|
* 2) the device must be idled to wait for the possible arrival of a new
|
|
* request for the queue.
|
|
* See the comments on the function bfq_better_to_idle for the reasons
|
|
* why performing device idling is the best choice to boost the throughput
|
|
* and preserve service guarantees when bfq_better_to_idle itself
|
|
* returns true.
|
|
*/
|
|
static bool bfq_bfqq_must_idle(struct bfq_queue *bfqq)
|
|
{
|
|
return RB_EMPTY_ROOT(&bfqq->sort_list) && bfq_better_to_idle(bfqq);
|
|
}
|
|
|
|
/*
|
|
* This function chooses the queue from which to pick the next extra
|
|
* I/O request to inject, if it finds a compatible queue. See the
|
|
* comments on bfq_update_inject_limit() for details on the injection
|
|
* mechanism, and for the definitions of the quantities mentioned
|
|
* below.
|
|
*/
|
|
static struct bfq_queue *
|
|
bfq_choose_bfqq_for_injection(struct bfq_data *bfqd)
|
|
{
|
|
struct bfq_queue *bfqq, *in_serv_bfqq = bfqd->in_service_queue;
|
|
unsigned int limit = in_serv_bfqq->inject_limit;
|
|
/*
|
|
* If
|
|
* - bfqq is not weight-raised and therefore does not carry
|
|
* time-critical I/O,
|
|
* or
|
|
* - regardless of whether bfqq is weight-raised, bfqq has
|
|
* however a long think time, during which it can absorb the
|
|
* effect of an appropriate number of extra I/O requests
|
|
* from other queues (see bfq_update_inject_limit for
|
|
* details on the computation of this number);
|
|
* then injection can be performed without restrictions.
|
|
*/
|
|
bool in_serv_always_inject = in_serv_bfqq->wr_coeff == 1 ||
|
|
!bfq_bfqq_has_short_ttime(in_serv_bfqq);
|
|
|
|
/*
|
|
* If
|
|
* - the baseline total service time could not be sampled yet,
|
|
* so the inject limit happens to be still 0, and
|
|
* - a lot of time has elapsed since the plugging of I/O
|
|
* dispatching started, so drive speed is being wasted
|
|
* significantly;
|
|
* then temporarily raise inject limit to one request.
|
|
*/
|
|
if (limit == 0 && in_serv_bfqq->last_serv_time_ns == 0 &&
|
|
bfq_bfqq_wait_request(in_serv_bfqq) &&
|
|
time_is_before_eq_jiffies(bfqd->last_idling_start_jiffies +
|
|
bfqd->bfq_slice_idle)
|
|
)
|
|
limit = 1;
|
|
|
|
if (bfqd->rq_in_driver >= limit)
|
|
return NULL;
|
|
|
|
/*
|
|
* Linear search of the source queue for injection; but, with
|
|
* a high probability, very few steps are needed to find a
|
|
* candidate queue, i.e., a queue with enough budget left for
|
|
* its next request. In fact:
|
|
* - BFQ dynamically updates the budget of every queue so as
|
|
* to accommodate the expected backlog of the queue;
|
|
* - if a queue gets all its requests dispatched as injected
|
|
* service, then the queue is removed from the active list
|
|
* (and re-added only if it gets new requests, but then it
|
|
* is assigned again enough budget for its new backlog).
|
|
*/
|
|
list_for_each_entry(bfqq, &bfqd->active_list, bfqq_list)
|
|
if (!RB_EMPTY_ROOT(&bfqq->sort_list) &&
|
|
(in_serv_always_inject || bfqq->wr_coeff > 1) &&
|
|
bfq_serv_to_charge(bfqq->next_rq, bfqq) <=
|
|
bfq_bfqq_budget_left(bfqq)) {
|
|
/*
|
|
* Allow for only one large in-flight request
|
|
* on non-rotational devices, for the
|
|
* following reason. On non-rotationl drives,
|
|
* large requests take much longer than
|
|
* smaller requests to be served. In addition,
|
|
* the drive prefers to serve large requests
|
|
* w.r.t. to small ones, if it can choose. So,
|
|
* having more than one large requests queued
|
|
* in the drive may easily make the next first
|
|
* request of the in-service queue wait for so
|
|
* long to break bfqq's service guarantees. On
|
|
* the bright side, large requests let the
|
|
* drive reach a very high throughput, even if
|
|
* there is only one in-flight large request
|
|
* at a time.
|
|
*/
|
|
if (blk_queue_nonrot(bfqd->queue) &&
|
|
blk_rq_sectors(bfqq->next_rq) >=
|
|
BFQQ_SECT_THR_NONROT)
|
|
limit = min_t(unsigned int, 1, limit);
|
|
else
|
|
limit = in_serv_bfqq->inject_limit;
|
|
|
|
if (bfqd->rq_in_driver < limit) {
|
|
bfqd->rqs_injected = true;
|
|
return bfqq;
|
|
}
|
|
}
|
|
|
|
return NULL;
|
|
}
|
|
|
|
/*
|
|
* Select a queue for service. If we have a current queue in service,
|
|
* check whether to continue servicing it, or retrieve and set a new one.
|
|
*/
|
|
static struct bfq_queue *bfq_select_queue(struct bfq_data *bfqd)
|
|
{
|
|
struct bfq_queue *bfqq;
|
|
struct request *next_rq;
|
|
enum bfqq_expiration reason = BFQQE_BUDGET_TIMEOUT;
|
|
|
|
bfqq = bfqd->in_service_queue;
|
|
if (!bfqq)
|
|
goto new_queue;
|
|
|
|
bfq_log_bfqq(bfqd, bfqq, "select_queue: already in-service queue");
|
|
|
|
/*
|
|
* Do not expire bfqq for budget timeout if bfqq may be about
|
|
* to enjoy device idling. The reason why, in this case, we
|
|
* prevent bfqq from expiring is the same as in the comments
|
|
* on the case where bfq_bfqq_must_idle() returns true, in
|
|
* bfq_completed_request().
|
|
*/
|
|
if (bfq_may_expire_for_budg_timeout(bfqq) &&
|
|
!bfq_bfqq_must_idle(bfqq))
|
|
goto expire;
|
|
|
|
check_queue:
|
|
/*
|
|
* This loop is rarely executed more than once. Even when it
|
|
* happens, it is much more convenient to re-execute this loop
|
|
* than to return NULL and trigger a new dispatch to get a
|
|
* request served.
|
|
*/
|
|
next_rq = bfqq->next_rq;
|
|
/*
|
|
* If bfqq has requests queued and it has enough budget left to
|
|
* serve them, keep the queue, otherwise expire it.
|
|
*/
|
|
if (next_rq) {
|
|
if (bfq_serv_to_charge(next_rq, bfqq) >
|
|
bfq_bfqq_budget_left(bfqq)) {
|
|
/*
|
|
* Expire the queue for budget exhaustion,
|
|
* which makes sure that the next budget is
|
|
* enough to serve the next request, even if
|
|
* it comes from the fifo expired path.
|
|
*/
|
|
reason = BFQQE_BUDGET_EXHAUSTED;
|
|
goto expire;
|
|
} else {
|
|
/*
|
|
* The idle timer may be pending because we may
|
|
* not disable disk idling even when a new request
|
|
* arrives.
|
|
*/
|
|
if (bfq_bfqq_wait_request(bfqq)) {
|
|
/*
|
|
* If we get here: 1) at least a new request
|
|
* has arrived but we have not disabled the
|
|
* timer because the request was too small,
|
|
* 2) then the block layer has unplugged
|
|
* the device, causing the dispatch to be
|
|
* invoked.
|
|
*
|
|
* Since the device is unplugged, now the
|
|
* requests are probably large enough to
|
|
* provide a reasonable throughput.
|
|
* So we disable idling.
|
|
*/
|
|
bfq_clear_bfqq_wait_request(bfqq);
|
|
hrtimer_try_to_cancel(&bfqd->idle_slice_timer);
|
|
}
|
|
goto keep_queue;
|
|
}
|
|
}
|
|
|
|
/*
|
|
* No requests pending. However, if the in-service queue is idling
|
|
* for a new request, or has requests waiting for a completion and
|
|
* may idle after their completion, then keep it anyway.
|
|
*
|
|
* Yet, inject service from other queues if it boosts
|
|
* throughput and is possible.
|
|
*/
|
|
if (bfq_bfqq_wait_request(bfqq) ||
|
|
(bfqq->dispatched != 0 && bfq_better_to_idle(bfqq))) {
|
|
struct bfq_queue *async_bfqq =
|
|
bfqq->bic && bfqq->bic->bfqq[0] &&
|
|
bfq_bfqq_busy(bfqq->bic->bfqq[0]) &&
|
|
bfqq->bic->bfqq[0]->next_rq ?
|
|
bfqq->bic->bfqq[0] : NULL;
|
|
struct bfq_queue *blocked_bfqq =
|
|
!hlist_empty(&bfqq->woken_list) ?
|
|
container_of(bfqq->woken_list.first,
|
|
struct bfq_queue,
|
|
woken_list_node)
|
|
: NULL;
|
|
|
|
/*
|
|
* The next four mutually-exclusive ifs decide
|
|
* whether to try injection, and choose the queue to
|
|
* pick an I/O request from.
|
|
*
|
|
* The first if checks whether the process associated
|
|
* with bfqq has also async I/O pending. If so, it
|
|
* injects such I/O unconditionally. Injecting async
|
|
* I/O from the same process can cause no harm to the
|
|
* process. On the contrary, it can only increase
|
|
* bandwidth and reduce latency for the process.
|
|
*
|
|
* The second if checks whether there happens to be a
|
|
* non-empty waker queue for bfqq, i.e., a queue whose
|
|
* I/O needs to be completed for bfqq to receive new
|
|
* I/O. This happens, e.g., if bfqq is associated with
|
|
* a process that does some sync. A sync generates
|
|
* extra blocking I/O, which must be completed before
|
|
* the process associated with bfqq can go on with its
|
|
* I/O. If the I/O of the waker queue is not served,
|
|
* then bfqq remains empty, and no I/O is dispatched,
|
|
* until the idle timeout fires for bfqq. This is
|
|
* likely to result in lower bandwidth and higher
|
|
* latencies for bfqq, and in a severe loss of total
|
|
* throughput. The best action to take is therefore to
|
|
* serve the waker queue as soon as possible. So do it
|
|
* (without relying on the third alternative below for
|
|
* eventually serving waker_bfqq's I/O; see the last
|
|
* paragraph for further details). This systematic
|
|
* injection of I/O from the waker queue does not
|
|
* cause any delay to bfqq's I/O. On the contrary,
|
|
* next bfqq's I/O is brought forward dramatically,
|
|
* for it is not blocked for milliseconds.
|
|
*
|
|
* The third if checks whether there is a queue woken
|
|
* by bfqq, and currently with pending I/O. Such a
|
|
* woken queue does not steal bandwidth from bfqq,
|
|
* because it remains soon without I/O if bfqq is not
|
|
* served. So there is virtually no risk of loss of
|
|
* bandwidth for bfqq if this woken queue has I/O
|
|
* dispatched while bfqq is waiting for new I/O.
|
|
*
|
|
* The fourth if checks whether bfqq is a queue for
|
|
* which it is better to avoid injection. It is so if
|
|
* bfqq delivers more throughput when served without
|
|
* any further I/O from other queues in the middle, or
|
|
* if the service times of bfqq's I/O requests both
|
|
* count more than overall throughput, and may be
|
|
* easily increased by injection (this happens if bfqq
|
|
* has a short think time). If none of these
|
|
* conditions holds, then a candidate queue for
|
|
* injection is looked for through
|
|
* bfq_choose_bfqq_for_injection(). Note that the
|
|
* latter may return NULL (for example if the inject
|
|
* limit for bfqq is currently 0).
|
|
*
|
|
* NOTE: motivation for the second alternative
|
|
*
|
|
* Thanks to the way the inject limit is updated in
|
|
* bfq_update_has_short_ttime(), it is rather likely
|
|
* that, if I/O is being plugged for bfqq and the
|
|
* waker queue has pending I/O requests that are
|
|
* blocking bfqq's I/O, then the fourth alternative
|
|
* above lets the waker queue get served before the
|
|
* I/O-plugging timeout fires. So one may deem the
|
|
* second alternative superfluous. It is not, because
|
|
* the fourth alternative may be way less effective in
|
|
* case of a synchronization. For two main
|
|
* reasons. First, throughput may be low because the
|
|
* inject limit may be too low to guarantee the same
|
|
* amount of injected I/O, from the waker queue or
|
|
* other queues, that the second alternative
|
|
* guarantees (the second alternative unconditionally
|
|
* injects a pending I/O request of the waker queue
|
|
* for each bfq_dispatch_request()). Second, with the
|
|
* fourth alternative, the duration of the plugging,
|
|
* i.e., the time before bfqq finally receives new I/O,
|
|
* may not be minimized, because the waker queue may
|
|
* happen to be served only after other queues.
|
|
*/
|
|
if (async_bfqq &&
|
|
icq_to_bic(async_bfqq->next_rq->elv.icq) == bfqq->bic &&
|
|
bfq_serv_to_charge(async_bfqq->next_rq, async_bfqq) <=
|
|
bfq_bfqq_budget_left(async_bfqq))
|
|
bfqq = bfqq->bic->bfqq[0];
|
|
else if (bfqq->waker_bfqq &&
|
|
bfq_bfqq_busy(bfqq->waker_bfqq) &&
|
|
bfqq->waker_bfqq->next_rq &&
|
|
bfq_serv_to_charge(bfqq->waker_bfqq->next_rq,
|
|
bfqq->waker_bfqq) <=
|
|
bfq_bfqq_budget_left(bfqq->waker_bfqq)
|
|
)
|
|
bfqq = bfqq->waker_bfqq;
|
|
else if (blocked_bfqq &&
|
|
bfq_bfqq_busy(blocked_bfqq) &&
|
|
blocked_bfqq->next_rq &&
|
|
bfq_serv_to_charge(blocked_bfqq->next_rq,
|
|
blocked_bfqq) <=
|
|
bfq_bfqq_budget_left(blocked_bfqq)
|
|
)
|
|
bfqq = blocked_bfqq;
|
|
else if (!idling_boosts_thr_without_issues(bfqd, bfqq) &&
|
|
(bfqq->wr_coeff == 1 || bfqd->wr_busy_queues > 1 ||
|
|
!bfq_bfqq_has_short_ttime(bfqq)))
|
|
bfqq = bfq_choose_bfqq_for_injection(bfqd);
|
|
else
|
|
bfqq = NULL;
|
|
|
|
goto keep_queue;
|
|
}
|
|
|
|
reason = BFQQE_NO_MORE_REQUESTS;
|
|
expire:
|
|
bfq_bfqq_expire(bfqd, bfqq, false, reason);
|
|
new_queue:
|
|
bfqq = bfq_set_in_service_queue(bfqd);
|
|
if (bfqq) {
|
|
bfq_log_bfqq(bfqd, bfqq, "select_queue: checking new queue");
|
|
goto check_queue;
|
|
}
|
|
keep_queue:
|
|
if (bfqq)
|
|
bfq_log_bfqq(bfqd, bfqq, "select_queue: returned this queue");
|
|
else
|
|
bfq_log(bfqd, "select_queue: no queue returned");
|
|
|
|
return bfqq;
|
|
}
|
|
|
|
static void bfq_update_wr_data(struct bfq_data *bfqd, struct bfq_queue *bfqq)
|
|
{
|
|
struct bfq_entity *entity = &bfqq->entity;
|
|
|
|
if (bfqq->wr_coeff > 1) { /* queue is being weight-raised */
|
|
bfq_log_bfqq(bfqd, bfqq,
|
|
"raising period dur %u/%u msec, old coeff %u, w %d(%d)",
|
|
jiffies_to_msecs(jiffies - bfqq->last_wr_start_finish),
|
|
jiffies_to_msecs(bfqq->wr_cur_max_time),
|
|
bfqq->wr_coeff,
|
|
bfqq->entity.weight, bfqq->entity.orig_weight);
|
|
|
|
if (entity->prio_changed)
|
|
bfq_log_bfqq(bfqd, bfqq, "WARN: pending prio change");
|
|
|
|
/*
|
|
* If the queue was activated in a burst, or too much
|
|
* time has elapsed from the beginning of this
|
|
* weight-raising period, then end weight raising.
|
|
*/
|
|
if (bfq_bfqq_in_large_burst(bfqq))
|
|
bfq_bfqq_end_wr(bfqq);
|
|
else if (time_is_before_jiffies(bfqq->last_wr_start_finish +
|
|
bfqq->wr_cur_max_time)) {
|
|
if (bfqq->wr_cur_max_time != bfqd->bfq_wr_rt_max_time ||
|
|
time_is_before_jiffies(bfqq->wr_start_at_switch_to_srt +
|
|
bfq_wr_duration(bfqd))) {
|
|
/*
|
|
* Either in interactive weight
|
|
* raising, or in soft_rt weight
|
|
* raising with the
|
|
* interactive-weight-raising period
|
|
* elapsed (so no switch back to
|
|
* interactive weight raising).
|
|
*/
|
|
bfq_bfqq_end_wr(bfqq);
|
|
} else { /*
|
|
* soft_rt finishing while still in
|
|
* interactive period, switch back to
|
|
* interactive weight raising
|
|
*/
|
|
switch_back_to_interactive_wr(bfqq, bfqd);
|
|
bfqq->entity.prio_changed = 1;
|
|
}
|
|
}
|
|
if (bfqq->wr_coeff > 1 &&
|
|
bfqq->wr_cur_max_time != bfqd->bfq_wr_rt_max_time &&
|
|
bfqq->service_from_wr > max_service_from_wr) {
|
|
/* see comments on max_service_from_wr */
|
|
bfq_bfqq_end_wr(bfqq);
|
|
}
|
|
}
|
|
/*
|
|
* To improve latency (for this or other queues), immediately
|
|
* update weight both if it must be raised and if it must be
|
|
* lowered. Since, entity may be on some active tree here, and
|
|
* might have a pending change of its ioprio class, invoke
|
|
* next function with the last parameter unset (see the
|
|
* comments on the function).
|
|
*/
|
|
if ((entity->weight > entity->orig_weight) != (bfqq->wr_coeff > 1))
|
|
__bfq_entity_update_weight_prio(bfq_entity_service_tree(entity),
|
|
entity, false);
|
|
}
|
|
|
|
/*
|
|
* Dispatch next request from bfqq.
|
|
*/
|
|
static struct request *bfq_dispatch_rq_from_bfqq(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
struct request *rq = bfqq->next_rq;
|
|
unsigned long service_to_charge;
|
|
|
|
service_to_charge = bfq_serv_to_charge(rq, bfqq);
|
|
|
|
bfq_bfqq_served(bfqq, service_to_charge);
|
|
|
|
if (bfqq == bfqd->in_service_queue && bfqd->wait_dispatch) {
|
|
bfqd->wait_dispatch = false;
|
|
bfqd->waited_rq = rq;
|
|
}
|
|
|
|
bfq_dispatch_remove(bfqd->queue, rq);
|
|
|
|
if (bfqq != bfqd->in_service_queue)
|
|
goto return_rq;
|
|
|
|
/*
|
|
* If weight raising has to terminate for bfqq, then next
|
|
* function causes an immediate update of bfqq's weight,
|
|
* without waiting for next activation. As a consequence, on
|
|
* expiration, bfqq will be timestamped as if has never been
|
|
* weight-raised during this service slot, even if it has
|
|
* received part or even most of the service as a
|
|
* weight-raised queue. This inflates bfqq's timestamps, which
|
|
* is beneficial, as bfqq is then more willing to leave the
|
|
* device immediately to possible other weight-raised queues.
|
|
*/
|
|
bfq_update_wr_data(bfqd, bfqq);
|
|
|
|
/*
|
|
* Expire bfqq, pretending that its budget expired, if bfqq
|
|
* belongs to CLASS_IDLE and other queues are waiting for
|
|
* service.
|
|
*/
|
|
if (!(bfq_tot_busy_queues(bfqd) > 1 && bfq_class_idle(bfqq)))
|
|
goto return_rq;
|
|
|
|
bfq_bfqq_expire(bfqd, bfqq, false, BFQQE_BUDGET_EXHAUSTED);
|
|
|
|
return_rq:
|
|
return rq;
|
|
}
|
|
|
|
static bool bfq_has_work(struct blk_mq_hw_ctx *hctx)
|
|
{
|
|
struct bfq_data *bfqd = hctx->queue->elevator->elevator_data;
|
|
|
|
/*
|
|
* Avoiding lock: a race on bfqd->busy_queues should cause at
|
|
* most a call to dispatch for nothing
|
|
*/
|
|
return !list_empty_careful(&bfqd->dispatch) ||
|
|
bfq_tot_busy_queues(bfqd) > 0;
|
|
}
|
|
|
|
static struct request *__bfq_dispatch_request(struct blk_mq_hw_ctx *hctx)
|
|
{
|
|
struct bfq_data *bfqd = hctx->queue->elevator->elevator_data;
|
|
struct request *rq = NULL;
|
|
struct bfq_queue *bfqq = NULL;
|
|
|
|
if (!list_empty(&bfqd->dispatch)) {
|
|
rq = list_first_entry(&bfqd->dispatch, struct request,
|
|
queuelist);
|
|
list_del_init(&rq->queuelist);
|
|
|
|
bfqq = RQ_BFQQ(rq);
|
|
|
|
if (bfqq) {
|
|
/*
|
|
* Increment counters here, because this
|
|
* dispatch does not follow the standard
|
|
* dispatch flow (where counters are
|
|
* incremented)
|
|
*/
|
|
bfqq->dispatched++;
|
|
|
|
goto inc_in_driver_start_rq;
|
|
}
|
|
|
|
/*
|
|
* We exploit the bfq_finish_requeue_request hook to
|
|
* decrement rq_in_driver, but
|
|
* bfq_finish_requeue_request will not be invoked on
|
|
* this request. So, to avoid unbalance, just start
|
|
* this request, without incrementing rq_in_driver. As
|
|
* a negative consequence, rq_in_driver is deceptively
|
|
* lower than it should be while this request is in
|
|
* service. This may cause bfq_schedule_dispatch to be
|
|
* invoked uselessly.
|
|
*
|
|
* As for implementing an exact solution, the
|
|
* bfq_finish_requeue_request hook, if defined, is
|
|
* probably invoked also on this request. So, by
|
|
* exploiting this hook, we could 1) increment
|
|
* rq_in_driver here, and 2) decrement it in
|
|
* bfq_finish_requeue_request. Such a solution would
|
|
* let the value of the counter be always accurate,
|
|
* but it would entail using an extra interface
|
|
* function. This cost seems higher than the benefit,
|
|
* being the frequency of non-elevator-private
|
|
* requests very low.
|
|
*/
|
|
goto start_rq;
|
|
}
|
|
|
|
bfq_log(bfqd, "dispatch requests: %d busy queues",
|
|
bfq_tot_busy_queues(bfqd));
|
|
|
|
if (bfq_tot_busy_queues(bfqd) == 0)
|
|
goto exit;
|
|
|
|
/*
|
|
* Force device to serve one request at a time if
|
|
* strict_guarantees is true. Forcing this service scheme is
|
|
* currently the ONLY way to guarantee that the request
|
|
* service order enforced by the scheduler is respected by a
|
|
* queueing device. Otherwise the device is free even to make
|
|
* some unlucky request wait for as long as the device
|
|
* wishes.
|
|
*
|
|
* Of course, serving one request at a time may cause loss of
|
|
* throughput.
|
|
*/
|
|
if (bfqd->strict_guarantees && bfqd->rq_in_driver > 0)
|
|
goto exit;
|
|
|
|
bfqq = bfq_select_queue(bfqd);
|
|
if (!bfqq)
|
|
goto exit;
|
|
|
|
rq = bfq_dispatch_rq_from_bfqq(bfqd, bfqq);
|
|
|
|
if (rq) {
|
|
inc_in_driver_start_rq:
|
|
bfqd->rq_in_driver++;
|
|
start_rq:
|
|
rq->rq_flags |= RQF_STARTED;
|
|
}
|
|
exit:
|
|
return rq;
|
|
}
|
|
|
|
#ifdef CONFIG_BFQ_CGROUP_DEBUG
|
|
static void bfq_update_dispatch_stats(struct request_queue *q,
|
|
struct request *rq,
|
|
struct bfq_queue *in_serv_queue,
|
|
bool idle_timer_disabled)
|
|
{
|
|
struct bfq_queue *bfqq = rq ? RQ_BFQQ(rq) : NULL;
|
|
|
|
if (!idle_timer_disabled && !bfqq)
|
|
return;
|
|
|
|
/*
|
|
* rq and bfqq are guaranteed to exist until this function
|
|
* ends, for the following reasons. First, rq can be
|
|
* dispatched to the device, and then can be completed and
|
|
* freed, only after this function ends. Second, rq cannot be
|
|
* merged (and thus freed because of a merge) any longer,
|
|
* because it has already started. Thus rq cannot be freed
|
|
* before this function ends, and, since rq has a reference to
|
|
* bfqq, the same guarantee holds for bfqq too.
|
|
*
|
|
* In addition, the following queue lock guarantees that
|
|
* bfqq_group(bfqq) exists as well.
|
|
*/
|
|
spin_lock_irq(&q->queue_lock);
|
|
if (idle_timer_disabled)
|
|
/*
|
|
* Since the idle timer has been disabled,
|
|
* in_serv_queue contained some request when
|
|
* __bfq_dispatch_request was invoked above, which
|
|
* implies that rq was picked exactly from
|
|
* in_serv_queue. Thus in_serv_queue == bfqq, and is
|
|
* therefore guaranteed to exist because of the above
|
|
* arguments.
|
|
*/
|
|
bfqg_stats_update_idle_time(bfqq_group(in_serv_queue));
|
|
if (bfqq) {
|
|
struct bfq_group *bfqg = bfqq_group(bfqq);
|
|
|
|
bfqg_stats_update_avg_queue_size(bfqg);
|
|
bfqg_stats_set_start_empty_time(bfqg);
|
|
bfqg_stats_update_io_remove(bfqg, rq->cmd_flags);
|
|
}
|
|
spin_unlock_irq(&q->queue_lock);
|
|
}
|
|
#else
|
|
static inline void bfq_update_dispatch_stats(struct request_queue *q,
|
|
struct request *rq,
|
|
struct bfq_queue *in_serv_queue,
|
|
bool idle_timer_disabled) {}
|
|
#endif /* CONFIG_BFQ_CGROUP_DEBUG */
|
|
|
|
static struct request *bfq_dispatch_request(struct blk_mq_hw_ctx *hctx)
|
|
{
|
|
struct bfq_data *bfqd = hctx->queue->elevator->elevator_data;
|
|
struct request *rq;
|
|
struct bfq_queue *in_serv_queue;
|
|
bool waiting_rq, idle_timer_disabled;
|
|
|
|
spin_lock_irq(&bfqd->lock);
|
|
|
|
in_serv_queue = bfqd->in_service_queue;
|
|
waiting_rq = in_serv_queue && bfq_bfqq_wait_request(in_serv_queue);
|
|
|
|
rq = __bfq_dispatch_request(hctx);
|
|
|
|
idle_timer_disabled =
|
|
waiting_rq && !bfq_bfqq_wait_request(in_serv_queue);
|
|
|
|
spin_unlock_irq(&bfqd->lock);
|
|
|
|
bfq_update_dispatch_stats(hctx->queue, rq, in_serv_queue,
|
|
idle_timer_disabled);
|
|
|
|
return rq;
|
|
}
|
|
|
|
/*
|
|
* Task holds one reference to the queue, dropped when task exits. Each rq
|
|
* in-flight on this queue also holds a reference, dropped when rq is freed.
|
|
*
|
|
* Scheduler lock must be held here. Recall not to use bfqq after calling
|
|
* this function on it.
|
|
*/
|
|
void bfq_put_queue(struct bfq_queue *bfqq)
|
|
{
|
|
struct bfq_queue *item;
|
|
struct hlist_node *n;
|
|
struct bfq_group *bfqg = bfqq_group(bfqq);
|
|
|
|
if (bfqq->bfqd)
|
|
bfq_log_bfqq(bfqq->bfqd, bfqq, "put_queue: %p %d",
|
|
bfqq, bfqq->ref);
|
|
|
|
bfqq->ref--;
|
|
if (bfqq->ref)
|
|
return;
|
|
|
|
if (!hlist_unhashed(&bfqq->burst_list_node)) {
|
|
hlist_del_init(&bfqq->burst_list_node);
|
|
/*
|
|
* Decrement also burst size after the removal, if the
|
|
* process associated with bfqq is exiting, and thus
|
|
* does not contribute to the burst any longer. This
|
|
* decrement helps filter out false positives of large
|
|
* bursts, when some short-lived process (often due to
|
|
* the execution of commands by some service) happens
|
|
* to start and exit while a complex application is
|
|
* starting, and thus spawning several processes that
|
|
* do I/O (and that *must not* be treated as a large
|
|
* burst, see comments on bfq_handle_burst).
|
|
*
|
|
* In particular, the decrement is performed only if:
|
|
* 1) bfqq is not a merged queue, because, if it is,
|
|
* then this free of bfqq is not triggered by the exit
|
|
* of the process bfqq is associated with, but exactly
|
|
* by the fact that bfqq has just been merged.
|
|
* 2) burst_size is greater than 0, to handle
|
|
* unbalanced decrements. Unbalanced decrements may
|
|
* happen in te following case: bfqq is inserted into
|
|
* the current burst list--without incrementing
|
|
* bust_size--because of a split, but the current
|
|
* burst list is not the burst list bfqq belonged to
|
|
* (see comments on the case of a split in
|
|
* bfq_set_request).
|
|
*/
|
|
if (bfqq->bic && bfqq->bfqd->burst_size > 0)
|
|
bfqq->bfqd->burst_size--;
|
|
}
|
|
|
|
/*
|
|
* bfqq does not exist any longer, so it cannot be woken by
|
|
* any other queue, and cannot wake any other queue. Then bfqq
|
|
* must be removed from the woken list of its possible waker
|
|
* queue, and all queues in the woken list of bfqq must stop
|
|
* having a waker queue. Strictly speaking, these updates
|
|
* should be performed when bfqq remains with no I/O source
|
|
* attached to it, which happens before bfqq gets freed. In
|
|
* particular, this happens when the last process associated
|
|
* with bfqq exits or gets associated with a different
|
|
* queue. However, both events lead to bfqq being freed soon,
|
|
* and dangling references would come out only after bfqq gets
|
|
* freed. So these updates are done here, as a simple and safe
|
|
* way to handle all cases.
|
|
*/
|
|
/* remove bfqq from woken list */
|
|
if (!hlist_unhashed(&bfqq->woken_list_node))
|
|
hlist_del_init(&bfqq->woken_list_node);
|
|
|
|
/* reset waker for all queues in woken list */
|
|
hlist_for_each_entry_safe(item, n, &bfqq->woken_list,
|
|
woken_list_node) {
|
|
item->waker_bfqq = NULL;
|
|
hlist_del_init(&item->woken_list_node);
|
|
}
|
|
|
|
if (bfqq->bfqd && bfqq->bfqd->last_completed_rq_bfqq == bfqq)
|
|
bfqq->bfqd->last_completed_rq_bfqq = NULL;
|
|
|
|
kmem_cache_free(bfq_pool, bfqq);
|
|
bfqg_and_blkg_put(bfqg);
|
|
}
|
|
|
|
static void bfq_put_stable_ref(struct bfq_queue *bfqq)
|
|
{
|
|
bfqq->stable_ref--;
|
|
bfq_put_queue(bfqq);
|
|
}
|
|
|
|
static void bfq_put_cooperator(struct bfq_queue *bfqq)
|
|
{
|
|
struct bfq_queue *__bfqq, *next;
|
|
|
|
/*
|
|
* If this queue was scheduled to merge with another queue, be
|
|
* sure to drop the reference taken on that queue (and others in
|
|
* the merge chain). See bfq_setup_merge and bfq_merge_bfqqs.
|
|
*/
|
|
__bfqq = bfqq->new_bfqq;
|
|
while (__bfqq) {
|
|
if (__bfqq == bfqq)
|
|
break;
|
|
next = __bfqq->new_bfqq;
|
|
bfq_put_queue(__bfqq);
|
|
__bfqq = next;
|
|
}
|
|
}
|
|
|
|
static void bfq_exit_bfqq(struct bfq_data *bfqd, struct bfq_queue *bfqq)
|
|
{
|
|
if (bfqq == bfqd->in_service_queue) {
|
|
__bfq_bfqq_expire(bfqd, bfqq, BFQQE_BUDGET_TIMEOUT);
|
|
bfq_schedule_dispatch(bfqd);
|
|
}
|
|
|
|
bfq_log_bfqq(bfqd, bfqq, "exit_bfqq: %p, %d", bfqq, bfqq->ref);
|
|
|
|
bfq_put_cooperator(bfqq);
|
|
|
|
bfq_release_process_ref(bfqd, bfqq);
|
|
}
|
|
|
|
static void bfq_exit_icq_bfqq(struct bfq_io_cq *bic, bool is_sync)
|
|
{
|
|
struct bfq_queue *bfqq = bic_to_bfqq(bic, is_sync);
|
|
struct bfq_data *bfqd;
|
|
|
|
if (bfqq)
|
|
bfqd = bfqq->bfqd; /* NULL if scheduler already exited */
|
|
|
|
if (bfqq && bfqd) {
|
|
unsigned long flags;
|
|
|
|
spin_lock_irqsave(&bfqd->lock, flags);
|
|
bfqq->bic = NULL;
|
|
bfq_exit_bfqq(bfqd, bfqq);
|
|
bic_set_bfqq(bic, NULL, is_sync);
|
|
spin_unlock_irqrestore(&bfqd->lock, flags);
|
|
}
|
|
}
|
|
|
|
static void bfq_exit_icq(struct io_cq *icq)
|
|
{
|
|
struct bfq_io_cq *bic = icq_to_bic(icq);
|
|
|
|
if (bic->stable_merge_bfqq) {
|
|
struct bfq_data *bfqd = bic->stable_merge_bfqq->bfqd;
|
|
|
|
/*
|
|
* bfqd is NULL if scheduler already exited, and in
|
|
* that case this is the last time bfqq is accessed.
|
|
*/
|
|
if (bfqd) {
|
|
unsigned long flags;
|
|
|
|
spin_lock_irqsave(&bfqd->lock, flags);
|
|
bfq_put_stable_ref(bic->stable_merge_bfqq);
|
|
spin_unlock_irqrestore(&bfqd->lock, flags);
|
|
} else {
|
|
bfq_put_stable_ref(bic->stable_merge_bfqq);
|
|
}
|
|
}
|
|
|
|
bfq_exit_icq_bfqq(bic, true);
|
|
bfq_exit_icq_bfqq(bic, false);
|
|
}
|
|
|
|
/*
|
|
* Update the entity prio values; note that the new values will not
|
|
* be used until the next (re)activation.
|
|
*/
|
|
static void
|
|
bfq_set_next_ioprio_data(struct bfq_queue *bfqq, struct bfq_io_cq *bic)
|
|
{
|
|
struct task_struct *tsk = current;
|
|
int ioprio_class;
|
|
struct bfq_data *bfqd = bfqq->bfqd;
|
|
|
|
if (!bfqd)
|
|
return;
|
|
|
|
ioprio_class = IOPRIO_PRIO_CLASS(bic->ioprio);
|
|
switch (ioprio_class) {
|
|
default:
|
|
pr_err("bdi %s: bfq: bad prio class %d\n",
|
|
bdi_dev_name(bfqq->bfqd->queue->backing_dev_info),
|
|
ioprio_class);
|
|
fallthrough;
|
|
case IOPRIO_CLASS_NONE:
|
|
/*
|
|
* No prio set, inherit CPU scheduling settings.
|
|
*/
|
|
bfqq->new_ioprio = task_nice_ioprio(tsk);
|
|
bfqq->new_ioprio_class = task_nice_ioclass(tsk);
|
|
break;
|
|
case IOPRIO_CLASS_RT:
|
|
bfqq->new_ioprio = IOPRIO_PRIO_DATA(bic->ioprio);
|
|
bfqq->new_ioprio_class = IOPRIO_CLASS_RT;
|
|
break;
|
|
case IOPRIO_CLASS_BE:
|
|
bfqq->new_ioprio = IOPRIO_PRIO_DATA(bic->ioprio);
|
|
bfqq->new_ioprio_class = IOPRIO_CLASS_BE;
|
|
break;
|
|
case IOPRIO_CLASS_IDLE:
|
|
bfqq->new_ioprio_class = IOPRIO_CLASS_IDLE;
|
|
bfqq->new_ioprio = 7;
|
|
break;
|
|
}
|
|
|
|
if (bfqq->new_ioprio >= IOPRIO_BE_NR) {
|
|
pr_crit("bfq_set_next_ioprio_data: new_ioprio %d\n",
|
|
bfqq->new_ioprio);
|
|
bfqq->new_ioprio = IOPRIO_BE_NR;
|
|
}
|
|
|
|
bfqq->entity.new_weight = bfq_ioprio_to_weight(bfqq->new_ioprio);
|
|
bfq_log_bfqq(bfqd, bfqq, "new_ioprio %d new_weight %d",
|
|
bfqq->new_ioprio, bfqq->entity.new_weight);
|
|
bfqq->entity.prio_changed = 1;
|
|
}
|
|
|
|
static struct bfq_queue *bfq_get_queue(struct bfq_data *bfqd,
|
|
struct bio *bio, bool is_sync,
|
|
struct bfq_io_cq *bic,
|
|
bool respawn);
|
|
|
|
static void bfq_check_ioprio_change(struct bfq_io_cq *bic, struct bio *bio)
|
|
{
|
|
struct bfq_data *bfqd = bic_to_bfqd(bic);
|
|
struct bfq_queue *bfqq;
|
|
int ioprio = bic->icq.ioc->ioprio;
|
|
|
|
/*
|
|
* This condition may trigger on a newly created bic, be sure to
|
|
* drop the lock before returning.
|
|
*/
|
|
if (unlikely(!bfqd) || likely(bic->ioprio == ioprio))
|
|
return;
|
|
|
|
bic->ioprio = ioprio;
|
|
|
|
bfqq = bic_to_bfqq(bic, false);
|
|
if (bfqq) {
|
|
bfq_release_process_ref(bfqd, bfqq);
|
|
bfqq = bfq_get_queue(bfqd, bio, BLK_RW_ASYNC, bic, true);
|
|
bic_set_bfqq(bic, bfqq, false);
|
|
}
|
|
|
|
bfqq = bic_to_bfqq(bic, true);
|
|
if (bfqq)
|
|
bfq_set_next_ioprio_data(bfqq, bic);
|
|
}
|
|
|
|
static void bfq_init_bfqq(struct bfq_data *bfqd, struct bfq_queue *bfqq,
|
|
struct bfq_io_cq *bic, pid_t pid, int is_sync)
|
|
{
|
|
u64 now_ns = ktime_get_ns();
|
|
|
|
RB_CLEAR_NODE(&bfqq->entity.rb_node);
|
|
INIT_LIST_HEAD(&bfqq->fifo);
|
|
INIT_HLIST_NODE(&bfqq->burst_list_node);
|
|
INIT_HLIST_NODE(&bfqq->woken_list_node);
|
|
INIT_HLIST_HEAD(&bfqq->woken_list);
|
|
|
|
bfqq->ref = 0;
|
|
bfqq->bfqd = bfqd;
|
|
|
|
if (bic)
|
|
bfq_set_next_ioprio_data(bfqq, bic);
|
|
|
|
if (is_sync) {
|
|
/*
|
|
* No need to mark as has_short_ttime if in
|
|
* idle_class, because no device idling is performed
|
|
* for queues in idle class
|
|
*/
|
|
if (!bfq_class_idle(bfqq))
|
|
/* tentatively mark as has_short_ttime */
|
|
bfq_mark_bfqq_has_short_ttime(bfqq);
|
|
bfq_mark_bfqq_sync(bfqq);
|
|
bfq_mark_bfqq_just_created(bfqq);
|
|
} else
|
|
bfq_clear_bfqq_sync(bfqq);
|
|
|
|
/* set end request to minus infinity from now */
|
|
bfqq->ttime.last_end_request = now_ns + 1;
|
|
|
|
bfqq->creation_time = jiffies;
|
|
|
|
bfqq->io_start_time = now_ns;
|
|
|
|
bfq_mark_bfqq_IO_bound(bfqq);
|
|
|
|
bfqq->pid = pid;
|
|
|
|
/* Tentative initial value to trade off between thr and lat */
|
|
bfqq->max_budget = (2 * bfq_max_budget(bfqd)) / 3;
|
|
bfqq->budget_timeout = bfq_smallest_from_now();
|
|
|
|
bfqq->wr_coeff = 1;
|
|
bfqq->last_wr_start_finish = jiffies;
|
|
bfqq->wr_start_at_switch_to_srt = bfq_smallest_from_now();
|
|
bfqq->split_time = bfq_smallest_from_now();
|
|
|
|
/*
|
|
* To not forget the possibly high bandwidth consumed by a
|
|
* process/queue in the recent past,
|
|
* bfq_bfqq_softrt_next_start() returns a value at least equal
|
|
* to the current value of bfqq->soft_rt_next_start (see
|
|
* comments on bfq_bfqq_softrt_next_start). Set
|
|
* soft_rt_next_start to now, to mean that bfqq has consumed
|
|
* no bandwidth so far.
|
|
*/
|
|
bfqq->soft_rt_next_start = jiffies;
|
|
|
|
/* first request is almost certainly seeky */
|
|
bfqq->seek_history = 1;
|
|
}
|
|
|
|
static struct bfq_queue **bfq_async_queue_prio(struct bfq_data *bfqd,
|
|
struct bfq_group *bfqg,
|
|
int ioprio_class, int ioprio)
|
|
{
|
|
switch (ioprio_class) {
|
|
case IOPRIO_CLASS_RT:
|
|
return &bfqg->async_bfqq[0][ioprio];
|
|
case IOPRIO_CLASS_NONE:
|
|
ioprio = IOPRIO_NORM;
|
|
fallthrough;
|
|
case IOPRIO_CLASS_BE:
|
|
return &bfqg->async_bfqq[1][ioprio];
|
|
case IOPRIO_CLASS_IDLE:
|
|
return &bfqg->async_idle_bfqq;
|
|
default:
|
|
return NULL;
|
|
}
|
|
}
|
|
|
|
static struct bfq_queue *
|
|
bfq_do_early_stable_merge(struct bfq_data *bfqd, struct bfq_queue *bfqq,
|
|
struct bfq_io_cq *bic,
|
|
struct bfq_queue *last_bfqq_created)
|
|
{
|
|
struct bfq_queue *new_bfqq =
|
|
bfq_setup_merge(bfqq, last_bfqq_created);
|
|
|
|
if (!new_bfqq)
|
|
return bfqq;
|
|
|
|
if (new_bfqq->bic)
|
|
new_bfqq->bic->stably_merged = true;
|
|
bic->stably_merged = true;
|
|
|
|
/*
|
|
* Reusing merge functions. This implies that
|
|
* bfqq->bic must be set too, for
|
|
* bfq_merge_bfqqs to correctly save bfqq's
|
|
* state before killing it.
|
|
*/
|
|
bfqq->bic = bic;
|
|
bfq_merge_bfqqs(bfqd, bic, bfqq, new_bfqq);
|
|
|
|
return new_bfqq;
|
|
}
|
|
|
|
/*
|
|
* Many throughput-sensitive workloads are made of several parallel
|
|
* I/O flows, with all flows generated by the same application, or
|
|
* more generically by the same task (e.g., system boot). The most
|
|
* counterproductive action with these workloads is plugging I/O
|
|
* dispatch when one of the bfq_queues associated with these flows
|
|
* remains temporarily empty.
|
|
*
|
|
* To avoid this plugging, BFQ has been using a burst-handling
|
|
* mechanism for years now. This mechanism has proven effective for
|
|
* throughput, and not detrimental for service guarantees. The
|
|
* following function pushes this mechanism a little bit further,
|
|
* basing on the following two facts.
|
|
*
|
|
* First, all the I/O flows of a the same application or task
|
|
* contribute to the execution/completion of that common application
|
|
* or task. So the performance figures that matter are total
|
|
* throughput of the flows and task-wide I/O latency. In particular,
|
|
* these flows do not need to be protected from each other, in terms
|
|
* of individual bandwidth or latency.
|
|
*
|
|
* Second, the above fact holds regardless of the number of flows.
|
|
*
|
|
* Putting these two facts together, this commits merges stably the
|
|
* bfq_queues associated with these I/O flows, i.e., with the
|
|
* processes that generate these IO/ flows, regardless of how many the
|
|
* involved processes are.
|
|
*
|
|
* To decide whether a set of bfq_queues is actually associated with
|
|
* the I/O flows of a common application or task, and to merge these
|
|
* queues stably, this function operates as follows: given a bfq_queue,
|
|
* say Q2, currently being created, and the last bfq_queue, say Q1,
|
|
* created before Q2, Q2 is merged stably with Q1 if
|
|
* - very little time has elapsed since when Q1 was created
|
|
* - Q2 has the same ioprio as Q1
|
|
* - Q2 belongs to the same group as Q1
|
|
*
|
|
* Merging bfq_queues also reduces scheduling overhead. A fio test
|
|
* with ten random readers on /dev/nullb shows a throughput boost of
|
|
* 40%, with a quadcore. Since BFQ's execution time amounts to ~50% of
|
|
* the total per-request processing time, the above throughput boost
|
|
* implies that BFQ's overhead is reduced by more than 50%.
|
|
*
|
|
* This new mechanism most certainly obsoletes the current
|
|
* burst-handling heuristics. We keep those heuristics for the moment.
|
|
*/
|
|
static struct bfq_queue *bfq_do_or_sched_stable_merge(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq,
|
|
struct bfq_io_cq *bic)
|
|
{
|
|
struct bfq_queue **source_bfqq = bfqq->entity.parent ?
|
|
&bfqq->entity.parent->last_bfqq_created :
|
|
&bfqd->last_bfqq_created;
|
|
|
|
struct bfq_queue *last_bfqq_created = *source_bfqq;
|
|
|
|
/*
|
|
* If last_bfqq_created has not been set yet, then init it. If
|
|
* it has been set already, but too long ago, then move it
|
|
* forward to bfqq. Finally, move also if bfqq belongs to a
|
|
* different group than last_bfqq_created, or if bfqq has a
|
|
* different ioprio or ioprio_class. If none of these
|
|
* conditions holds true, then try an early stable merge or
|
|
* schedule a delayed stable merge.
|
|
*
|
|
* A delayed merge is scheduled (instead of performing an
|
|
* early merge), in case bfqq might soon prove to be more
|
|
* throughput-beneficial if not merged. Currently this is
|
|
* possible only if bfqd is rotational with no queueing. For
|
|
* such a drive, not merging bfqq is better for throughput if
|
|
* bfqq happens to contain sequential I/O. So, we wait a
|
|
* little bit for enough I/O to flow through bfqq. After that,
|
|
* if such an I/O is sequential, then the merge is
|
|
* canceled. Otherwise the merge is finally performed.
|
|
*/
|
|
if (!last_bfqq_created ||
|
|
time_before(last_bfqq_created->creation_time +
|
|
bfqd->bfq_burst_interval,
|
|
bfqq->creation_time) ||
|
|
bfqq->entity.parent != last_bfqq_created->entity.parent ||
|
|
bfqq->ioprio != last_bfqq_created->ioprio ||
|
|
bfqq->ioprio_class != last_bfqq_created->ioprio_class)
|
|
*source_bfqq = bfqq;
|
|
else if (time_after_eq(last_bfqq_created->creation_time +
|
|
bfqd->bfq_burst_interval,
|
|
bfqq->creation_time)) {
|
|
if (likely(bfqd->nonrot_with_queueing))
|
|
/*
|
|
* With this type of drive, leaving
|
|
* bfqq alone may provide no
|
|
* throughput benefits compared with
|
|
* merging bfqq. So merge bfqq now.
|
|
*/
|
|
bfqq = bfq_do_early_stable_merge(bfqd, bfqq,
|
|
bic,
|
|
last_bfqq_created);
|
|
else { /* schedule tentative stable merge */
|
|
/*
|
|
* get reference on last_bfqq_created,
|
|
* to prevent it from being freed,
|
|
* until we decide whether to merge
|
|
*/
|
|
last_bfqq_created->ref++;
|
|
/*
|
|
* need to keep track of stable refs, to
|
|
* compute process refs correctly
|
|
*/
|
|
last_bfqq_created->stable_ref++;
|
|
/*
|
|
* Record the bfqq to merge to.
|
|
*/
|
|
bic->stable_merge_bfqq = last_bfqq_created;
|
|
}
|
|
}
|
|
|
|
return bfqq;
|
|
}
|
|
|
|
|
|
static struct bfq_queue *bfq_get_queue(struct bfq_data *bfqd,
|
|
struct bio *bio, bool is_sync,
|
|
struct bfq_io_cq *bic,
|
|
bool respawn)
|
|
{
|
|
const int ioprio = IOPRIO_PRIO_DATA(bic->ioprio);
|
|
const int ioprio_class = IOPRIO_PRIO_CLASS(bic->ioprio);
|
|
struct bfq_queue **async_bfqq = NULL;
|
|
struct bfq_queue *bfqq;
|
|
struct bfq_group *bfqg;
|
|
|
|
rcu_read_lock();
|
|
|
|
bfqg = bfq_find_set_group(bfqd, __bio_blkcg(bio));
|
|
if (!bfqg) {
|
|
bfqq = &bfqd->oom_bfqq;
|
|
goto out;
|
|
}
|
|
|
|
if (!is_sync) {
|
|
async_bfqq = bfq_async_queue_prio(bfqd, bfqg, ioprio_class,
|
|
ioprio);
|
|
bfqq = *async_bfqq;
|
|
if (bfqq)
|
|
goto out;
|
|
}
|
|
|
|
bfqq = kmem_cache_alloc_node(bfq_pool,
|
|
GFP_NOWAIT | __GFP_ZERO | __GFP_NOWARN,
|
|
bfqd->queue->node);
|
|
|
|
if (bfqq) {
|
|
bfq_init_bfqq(bfqd, bfqq, bic, current->pid,
|
|
is_sync);
|
|
bfq_init_entity(&bfqq->entity, bfqg);
|
|
bfq_log_bfqq(bfqd, bfqq, "allocated");
|
|
} else {
|
|
bfqq = &bfqd->oom_bfqq;
|
|
bfq_log_bfqq(bfqd, bfqq, "using oom bfqq");
|
|
goto out;
|
|
}
|
|
|
|
/*
|
|
* Pin the queue now that it's allocated, scheduler exit will
|
|
* prune it.
|
|
*/
|
|
if (async_bfqq) {
|
|
bfqq->ref++; /*
|
|
* Extra group reference, w.r.t. sync
|
|
* queue. This extra reference is removed
|
|
* only if bfqq->bfqg disappears, to
|
|
* guarantee that this queue is not freed
|
|
* until its group goes away.
|
|
*/
|
|
bfq_log_bfqq(bfqd, bfqq, "get_queue, bfqq not in async: %p, %d",
|
|
bfqq, bfqq->ref);
|
|
*async_bfqq = bfqq;
|
|
}
|
|
|
|
out:
|
|
bfqq->ref++; /* get a process reference to this queue */
|
|
|
|
if (bfqq != &bfqd->oom_bfqq && is_sync && !respawn)
|
|
bfqq = bfq_do_or_sched_stable_merge(bfqd, bfqq, bic);
|
|
|
|
rcu_read_unlock();
|
|
return bfqq;
|
|
}
|
|
|
|
static void bfq_update_io_thinktime(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
struct bfq_ttime *ttime = &bfqq->ttime;
|
|
u64 elapsed;
|
|
|
|
/*
|
|
* We are really interested in how long it takes for the queue to
|
|
* become busy when there is no outstanding IO for this queue. So
|
|
* ignore cases when the bfq queue has already IO queued.
|
|
*/
|
|
if (bfqq->dispatched || bfq_bfqq_busy(bfqq))
|
|
return;
|
|
elapsed = ktime_get_ns() - bfqq->ttime.last_end_request;
|
|
elapsed = min_t(u64, elapsed, 2ULL * bfqd->bfq_slice_idle);
|
|
|
|
ttime->ttime_samples = (7*ttime->ttime_samples + 256) / 8;
|
|
ttime->ttime_total = div_u64(7*ttime->ttime_total + 256*elapsed, 8);
|
|
ttime->ttime_mean = div64_ul(ttime->ttime_total + 128,
|
|
ttime->ttime_samples);
|
|
}
|
|
|
|
static void
|
|
bfq_update_io_seektime(struct bfq_data *bfqd, struct bfq_queue *bfqq,
|
|
struct request *rq)
|
|
{
|
|
bfqq->seek_history <<= 1;
|
|
bfqq->seek_history |= BFQ_RQ_SEEKY(bfqd, bfqq->last_request_pos, rq);
|
|
|
|
if (bfqq->wr_coeff > 1 &&
|
|
bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time &&
|
|
BFQQ_TOTALLY_SEEKY(bfqq)) {
|
|
if (time_is_before_jiffies(bfqq->wr_start_at_switch_to_srt +
|
|
bfq_wr_duration(bfqd))) {
|
|
/*
|
|
* In soft_rt weight raising with the
|
|
* interactive-weight-raising period
|
|
* elapsed (so no switch back to
|
|
* interactive weight raising).
|
|
*/
|
|
bfq_bfqq_end_wr(bfqq);
|
|
} else { /*
|
|
* stopping soft_rt weight raising
|
|
* while still in interactive period,
|
|
* switch back to interactive weight
|
|
* raising
|
|
*/
|
|
switch_back_to_interactive_wr(bfqq, bfqd);
|
|
bfqq->entity.prio_changed = 1;
|
|
}
|
|
}
|
|
}
|
|
|
|
static void bfq_update_has_short_ttime(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq,
|
|
struct bfq_io_cq *bic)
|
|
{
|
|
bool has_short_ttime = true, state_changed;
|
|
|
|
/*
|
|
* No need to update has_short_ttime if bfqq is async or in
|
|
* idle io prio class, or if bfq_slice_idle is zero, because
|
|
* no device idling is performed for bfqq in this case.
|
|
*/
|
|
if (!bfq_bfqq_sync(bfqq) || bfq_class_idle(bfqq) ||
|
|
bfqd->bfq_slice_idle == 0)
|
|
return;
|
|
|
|
/* Idle window just restored, statistics are meaningless. */
|
|
if (time_is_after_eq_jiffies(bfqq->split_time +
|
|
bfqd->bfq_wr_min_idle_time))
|
|
return;
|
|
|
|
/* Think time is infinite if no process is linked to
|
|
* bfqq. Otherwise check average think time to decide whether
|
|
* to mark as has_short_ttime. To this goal, compare average
|
|
* think time with half the I/O-plugging timeout.
|
|
*/
|
|
if (atomic_read(&bic->icq.ioc->active_ref) == 0 ||
|
|
(bfq_sample_valid(bfqq->ttime.ttime_samples) &&
|
|
bfqq->ttime.ttime_mean > bfqd->bfq_slice_idle>>1))
|
|
has_short_ttime = false;
|
|
|
|
state_changed = has_short_ttime != bfq_bfqq_has_short_ttime(bfqq);
|
|
|
|
if (has_short_ttime)
|
|
bfq_mark_bfqq_has_short_ttime(bfqq);
|
|
else
|
|
bfq_clear_bfqq_has_short_ttime(bfqq);
|
|
|
|
/*
|
|
* Until the base value for the total service time gets
|
|
* finally computed for bfqq, the inject limit does depend on
|
|
* the think-time state (short|long). In particular, the limit
|
|
* is 0 or 1 if the think time is deemed, respectively, as
|
|
* short or long (details in the comments in
|
|
* bfq_update_inject_limit()). Accordingly, the next
|
|
* instructions reset the inject limit if the think-time state
|
|
* has changed and the above base value is still to be
|
|
* computed.
|
|
*
|
|
* However, the reset is performed only if more than 100 ms
|
|
* have elapsed since the last update of the inject limit, or
|
|
* (inclusive) if the change is from short to long think
|
|
* time. The reason for this waiting is as follows.
|
|
*
|
|
* bfqq may have a long think time because of a
|
|
* synchronization with some other queue, i.e., because the
|
|
* I/O of some other queue may need to be completed for bfqq
|
|
* to receive new I/O. Details in the comments on the choice
|
|
* of the queue for injection in bfq_select_queue().
|
|
*
|
|
* As stressed in those comments, if such a synchronization is
|
|
* actually in place, then, without injection on bfqq, the
|
|
* blocking I/O cannot happen to served while bfqq is in
|
|
* service. As a consequence, if bfqq is granted
|
|
* I/O-dispatch-plugging, then bfqq remains empty, and no I/O
|
|
* is dispatched, until the idle timeout fires. This is likely
|
|
* to result in lower bandwidth and higher latencies for bfqq,
|
|
* and in a severe loss of total throughput.
|
|
*
|
|
* On the opposite end, a non-zero inject limit may allow the
|
|
* I/O that blocks bfqq to be executed soon, and therefore
|
|
* bfqq to receive new I/O soon.
|
|
*
|
|
* But, if the blocking gets actually eliminated, then the
|
|
* next think-time sample for bfqq may be very low. This in
|
|
* turn may cause bfqq's think time to be deemed
|
|
* short. Without the 100 ms barrier, this new state change
|
|
* would cause the body of the next if to be executed
|
|
* immediately. But this would set to 0 the inject
|
|
* limit. Without injection, the blocking I/O would cause the
|
|
* think time of bfqq to become long again, and therefore the
|
|
* inject limit to be raised again, and so on. The only effect
|
|
* of such a steady oscillation between the two think-time
|
|
* states would be to prevent effective injection on bfqq.
|
|
*
|
|
* In contrast, if the inject limit is not reset during such a
|
|
* long time interval as 100 ms, then the number of short
|
|
* think time samples can grow significantly before the reset
|
|
* is performed. As a consequence, the think time state can
|
|
* become stable before the reset. Therefore there will be no
|
|
* state change when the 100 ms elapse, and no reset of the
|
|
* inject limit. The inject limit remains steadily equal to 1
|
|
* both during and after the 100 ms. So injection can be
|
|
* performed at all times, and throughput gets boosted.
|
|
*
|
|
* An inject limit equal to 1 is however in conflict, in
|
|
* general, with the fact that the think time of bfqq is
|
|
* short, because injection may be likely to delay bfqq's I/O
|
|
* (as explained in the comments in
|
|
* bfq_update_inject_limit()). But this does not happen in
|
|
* this special case, because bfqq's low think time is due to
|
|
* an effective handling of a synchronization, through
|
|
* injection. In this special case, bfqq's I/O does not get
|
|
* delayed by injection; on the contrary, bfqq's I/O is
|
|
* brought forward, because it is not blocked for
|
|
* milliseconds.
|
|
*
|
|
* In addition, serving the blocking I/O much sooner, and much
|
|
* more frequently than once per I/O-plugging timeout, makes
|
|
* it much quicker to detect a waker queue (the concept of
|
|
* waker queue is defined in the comments in
|
|
* bfq_add_request()). This makes it possible to start sooner
|
|
* to boost throughput more effectively, by injecting the I/O
|
|
* of the waker queue unconditionally on every
|
|
* bfq_dispatch_request().
|
|
*
|
|
* One last, important benefit of not resetting the inject
|
|
* limit before 100 ms is that, during this time interval, the
|
|
* base value for the total service time is likely to get
|
|
* finally computed for bfqq, freeing the inject limit from
|
|
* its relation with the think time.
|
|
*/
|
|
if (state_changed && bfqq->last_serv_time_ns == 0 &&
|
|
(time_is_before_eq_jiffies(bfqq->decrease_time_jif +
|
|
msecs_to_jiffies(100)) ||
|
|
!has_short_ttime))
|
|
bfq_reset_inject_limit(bfqd, bfqq);
|
|
}
|
|
|
|
/*
|
|
* Called when a new fs request (rq) is added to bfqq. Check if there's
|
|
* something we should do about it.
|
|
*/
|
|
static void bfq_rq_enqueued(struct bfq_data *bfqd, struct bfq_queue *bfqq,
|
|
struct request *rq)
|
|
{
|
|
if (rq->cmd_flags & REQ_META)
|
|
bfqq->meta_pending++;
|
|
|
|
bfqq->last_request_pos = blk_rq_pos(rq) + blk_rq_sectors(rq);
|
|
|
|
if (bfqq == bfqd->in_service_queue && bfq_bfqq_wait_request(bfqq)) {
|
|
bool small_req = bfqq->queued[rq_is_sync(rq)] == 1 &&
|
|
blk_rq_sectors(rq) < 32;
|
|
bool budget_timeout = bfq_bfqq_budget_timeout(bfqq);
|
|
|
|
/*
|
|
* There is just this request queued: if
|
|
* - the request is small, and
|
|
* - we are idling to boost throughput, and
|
|
* - the queue is not to be expired,
|
|
* then just exit.
|
|
*
|
|
* In this way, if the device is being idled to wait
|
|
* for a new request from the in-service queue, we
|
|
* avoid unplugging the device and committing the
|
|
* device to serve just a small request. In contrast
|
|
* we wait for the block layer to decide when to
|
|
* unplug the device: hopefully, new requests will be
|
|
* merged to this one quickly, then the device will be
|
|
* unplugged and larger requests will be dispatched.
|
|
*/
|
|
if (small_req && idling_boosts_thr_without_issues(bfqd, bfqq) &&
|
|
!budget_timeout)
|
|
return;
|
|
|
|
/*
|
|
* A large enough request arrived, or idling is being
|
|
* performed to preserve service guarantees, or
|
|
* finally the queue is to be expired: in all these
|
|
* cases disk idling is to be stopped, so clear
|
|
* wait_request flag and reset timer.
|
|
*/
|
|
bfq_clear_bfqq_wait_request(bfqq);
|
|
hrtimer_try_to_cancel(&bfqd->idle_slice_timer);
|
|
|
|
/*
|
|
* The queue is not empty, because a new request just
|
|
* arrived. Hence we can safely expire the queue, in
|
|
* case of budget timeout, without risking that the
|
|
* timestamps of the queue are not updated correctly.
|
|
* See [1] for more details.
|
|
*/
|
|
if (budget_timeout)
|
|
bfq_bfqq_expire(bfqd, bfqq, false,
|
|
BFQQE_BUDGET_TIMEOUT);
|
|
}
|
|
}
|
|
|
|
/* returns true if it causes the idle timer to be disabled */
|
|
static bool __bfq_insert_request(struct bfq_data *bfqd, struct request *rq)
|
|
{
|
|
struct bfq_queue *bfqq = RQ_BFQQ(rq),
|
|
*new_bfqq = bfq_setup_cooperator(bfqd, bfqq, rq, true,
|
|
RQ_BIC(rq));
|
|
bool waiting, idle_timer_disabled = false;
|
|
|
|
if (new_bfqq) {
|
|
/*
|
|
* Release the request's reference to the old bfqq
|
|
* and make sure one is taken to the shared queue.
|
|
*/
|
|
new_bfqq->allocated++;
|
|
bfqq->allocated--;
|
|
new_bfqq->ref++;
|
|
/*
|
|
* If the bic associated with the process
|
|
* issuing this request still points to bfqq
|
|
* (and thus has not been already redirected
|
|
* to new_bfqq or even some other bfq_queue),
|
|
* then complete the merge and redirect it to
|
|
* new_bfqq.
|
|
*/
|
|
if (bic_to_bfqq(RQ_BIC(rq), 1) == bfqq)
|
|
bfq_merge_bfqqs(bfqd, RQ_BIC(rq),
|
|
bfqq, new_bfqq);
|
|
|
|
bfq_clear_bfqq_just_created(bfqq);
|
|
/*
|
|
* rq is about to be enqueued into new_bfqq,
|
|
* release rq reference on bfqq
|
|
*/
|
|
bfq_put_queue(bfqq);
|
|
rq->elv.priv[1] = new_bfqq;
|
|
bfqq = new_bfqq;
|
|
}
|
|
|
|
bfq_update_io_thinktime(bfqd, bfqq);
|
|
bfq_update_has_short_ttime(bfqd, bfqq, RQ_BIC(rq));
|
|
bfq_update_io_seektime(bfqd, bfqq, rq);
|
|
|
|
waiting = bfqq && bfq_bfqq_wait_request(bfqq);
|
|
bfq_add_request(rq);
|
|
idle_timer_disabled = waiting && !bfq_bfqq_wait_request(bfqq);
|
|
|
|
rq->fifo_time = ktime_get_ns() + bfqd->bfq_fifo_expire[rq_is_sync(rq)];
|
|
list_add_tail(&rq->queuelist, &bfqq->fifo);
|
|
|
|
bfq_rq_enqueued(bfqd, bfqq, rq);
|
|
|
|
return idle_timer_disabled;
|
|
}
|
|
|
|
#ifdef CONFIG_BFQ_CGROUP_DEBUG
|
|
static void bfq_update_insert_stats(struct request_queue *q,
|
|
struct bfq_queue *bfqq,
|
|
bool idle_timer_disabled,
|
|
unsigned int cmd_flags)
|
|
{
|
|
if (!bfqq)
|
|
return;
|
|
|
|
/*
|
|
* bfqq still exists, because it can disappear only after
|
|
* either it is merged with another queue, or the process it
|
|
* is associated with exits. But both actions must be taken by
|
|
* the same process currently executing this flow of
|
|
* instructions.
|
|
*
|
|
* In addition, the following queue lock guarantees that
|
|
* bfqq_group(bfqq) exists as well.
|
|
*/
|
|
spin_lock_irq(&q->queue_lock);
|
|
bfqg_stats_update_io_add(bfqq_group(bfqq), bfqq, cmd_flags);
|
|
if (idle_timer_disabled)
|
|
bfqg_stats_update_idle_time(bfqq_group(bfqq));
|
|
spin_unlock_irq(&q->queue_lock);
|
|
}
|
|
#else
|
|
static inline void bfq_update_insert_stats(struct request_queue *q,
|
|
struct bfq_queue *bfqq,
|
|
bool idle_timer_disabled,
|
|
unsigned int cmd_flags) {}
|
|
#endif /* CONFIG_BFQ_CGROUP_DEBUG */
|
|
|
|
static void bfq_insert_request(struct blk_mq_hw_ctx *hctx, struct request *rq,
|
|
bool at_head)
|
|
{
|
|
struct request_queue *q = hctx->queue;
|
|
struct bfq_data *bfqd = q->elevator->elevator_data;
|
|
struct bfq_queue *bfqq;
|
|
bool idle_timer_disabled = false;
|
|
unsigned int cmd_flags;
|
|
|
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED
|
|
if (!cgroup_subsys_on_dfl(io_cgrp_subsys) && rq->bio)
|
|
bfqg_stats_update_legacy_io(q, rq);
|
|
#endif
|
|
spin_lock_irq(&bfqd->lock);
|
|
if (blk_mq_sched_try_insert_merge(q, rq)) {
|
|
spin_unlock_irq(&bfqd->lock);
|
|
return;
|
|
}
|
|
|
|
spin_unlock_irq(&bfqd->lock);
|
|
|
|
trace_block_rq_insert(rq);
|
|
|
|
spin_lock_irq(&bfqd->lock);
|
|
bfqq = bfq_init_rq(rq);
|
|
|
|
/*
|
|
* Reqs with at_head or passthrough flags set are to be put
|
|
* directly into dispatch list. Additional case for putting rq
|
|
* directly into the dispatch queue: the only active
|
|
* bfq_queues are bfqq and either its waker bfq_queue or one
|
|
* of its woken bfq_queues. The rationale behind this
|
|
* additional condition is as follows:
|
|
* - consider a bfq_queue, say Q1, detected as a waker of
|
|
* another bfq_queue, say Q2
|
|
* - by definition of a waker, Q1 blocks the I/O of Q2, i.e.,
|
|
* some I/O of Q1 needs to be completed for new I/O of Q2
|
|
* to arrive. A notable example of waker is journald
|
|
* - so, Q1 and Q2 are in any respect the queues of two
|
|
* cooperating processes (or of two cooperating sets of
|
|
* processes): the goal of Q1's I/O is doing what needs to
|
|
* be done so that new Q2's I/O can finally be
|
|
* issued. Therefore, if the service of Q1's I/O is delayed,
|
|
* then Q2's I/O is delayed too. Conversely, if Q2's I/O is
|
|
* delayed, the goal of Q1's I/O is hindered.
|
|
* - as a consequence, if some I/O of Q1/Q2 arrives while
|
|
* Q2/Q1 is the only queue in service, there is absolutely
|
|
* no point in delaying the service of such an I/O. The
|
|
* only possible result is a throughput loss
|
|
* - so, when the above condition holds, the best option is to
|
|
* have the new I/O dispatched as soon as possible
|
|
* - the most effective and efficient way to attain the above
|
|
* goal is to put the new I/O directly in the dispatch
|
|
* list
|
|
* - as an additional restriction, Q1 and Q2 must be the only
|
|
* busy queues for this commit to put the I/O of Q2/Q1 in
|
|
* the dispatch list. This is necessary, because, if also
|
|
* other queues are waiting for service, then putting new
|
|
* I/O directly in the dispatch list may evidently cause a
|
|
* violation of service guarantees for the other queues
|
|
*/
|
|
if (!bfqq ||
|
|
(bfqq != bfqd->in_service_queue &&
|
|
bfqd->in_service_queue != NULL &&
|
|
bfq_tot_busy_queues(bfqd) == 1 + bfq_bfqq_busy(bfqq) &&
|
|
(bfqq->waker_bfqq == bfqd->in_service_queue ||
|
|
bfqd->in_service_queue->waker_bfqq == bfqq)) || at_head) {
|
|
if (at_head)
|
|
list_add(&rq->queuelist, &bfqd->dispatch);
|
|
else
|
|
list_add_tail(&rq->queuelist, &bfqd->dispatch);
|
|
} else {
|
|
idle_timer_disabled = __bfq_insert_request(bfqd, rq);
|
|
/*
|
|
* Update bfqq, because, if a queue merge has occurred
|
|
* in __bfq_insert_request, then rq has been
|
|
* redirected into a new queue.
|
|
*/
|
|
bfqq = RQ_BFQQ(rq);
|
|
|
|
if (rq_mergeable(rq)) {
|
|
elv_rqhash_add(q, rq);
|
|
if (!q->last_merge)
|
|
q->last_merge = rq;
|
|
}
|
|
}
|
|
|
|
/*
|
|
* Cache cmd_flags before releasing scheduler lock, because rq
|
|
* may disappear afterwards (for example, because of a request
|
|
* merge).
|
|
*/
|
|
cmd_flags = rq->cmd_flags;
|
|
|
|
spin_unlock_irq(&bfqd->lock);
|
|
|
|
bfq_update_insert_stats(q, bfqq, idle_timer_disabled,
|
|
cmd_flags);
|
|
}
|
|
|
|
static void bfq_insert_requests(struct blk_mq_hw_ctx *hctx,
|
|
struct list_head *list, bool at_head)
|
|
{
|
|
while (!list_empty(list)) {
|
|
struct request *rq;
|
|
|
|
rq = list_first_entry(list, struct request, queuelist);
|
|
list_del_init(&rq->queuelist);
|
|
bfq_insert_request(hctx, rq, at_head);
|
|
}
|
|
}
|
|
|
|
static void bfq_update_hw_tag(struct bfq_data *bfqd)
|
|
{
|
|
struct bfq_queue *bfqq = bfqd->in_service_queue;
|
|
|
|
bfqd->max_rq_in_driver = max_t(int, bfqd->max_rq_in_driver,
|
|
bfqd->rq_in_driver);
|
|
|
|
if (bfqd->hw_tag == 1)
|
|
return;
|
|
|
|
/*
|
|
* This sample is valid if the number of outstanding requests
|
|
* is large enough to allow a queueing behavior. Note that the
|
|
* sum is not exact, as it's not taking into account deactivated
|
|
* requests.
|
|
*/
|
|
if (bfqd->rq_in_driver + bfqd->queued <= BFQ_HW_QUEUE_THRESHOLD)
|
|
return;
|
|
|
|
/*
|
|
* If active queue hasn't enough requests and can idle, bfq might not
|
|
* dispatch sufficient requests to hardware. Don't zero hw_tag in this
|
|
* case
|
|
*/
|
|
if (bfqq && bfq_bfqq_has_short_ttime(bfqq) &&
|
|
bfqq->dispatched + bfqq->queued[0] + bfqq->queued[1] <
|
|
BFQ_HW_QUEUE_THRESHOLD &&
|
|
bfqd->rq_in_driver < BFQ_HW_QUEUE_THRESHOLD)
|
|
return;
|
|
|
|
if (bfqd->hw_tag_samples++ < BFQ_HW_QUEUE_SAMPLES)
|
|
return;
|
|
|
|
bfqd->hw_tag = bfqd->max_rq_in_driver > BFQ_HW_QUEUE_THRESHOLD;
|
|
bfqd->max_rq_in_driver = 0;
|
|
bfqd->hw_tag_samples = 0;
|
|
|
|
bfqd->nonrot_with_queueing =
|
|
blk_queue_nonrot(bfqd->queue) && bfqd->hw_tag;
|
|
}
|
|
|
|
static void bfq_completed_request(struct bfq_queue *bfqq, struct bfq_data *bfqd)
|
|
{
|
|
u64 now_ns;
|
|
u32 delta_us;
|
|
|
|
bfq_update_hw_tag(bfqd);
|
|
|
|
bfqd->rq_in_driver--;
|
|
bfqq->dispatched--;
|
|
|
|
if (!bfqq->dispatched && !bfq_bfqq_busy(bfqq)) {
|
|
/*
|
|
* Set budget_timeout (which we overload to store the
|
|
* time at which the queue remains with no backlog and
|
|
* no outstanding request; used by the weight-raising
|
|
* mechanism).
|
|
*/
|
|
bfqq->budget_timeout = jiffies;
|
|
|
|
bfq_weights_tree_remove(bfqd, bfqq);
|
|
}
|
|
|
|
now_ns = ktime_get_ns();
|
|
|
|
bfqq->ttime.last_end_request = now_ns;
|
|
|
|
/*
|
|
* Using us instead of ns, to get a reasonable precision in
|
|
* computing rate in next check.
|
|
*/
|
|
delta_us = div_u64(now_ns - bfqd->last_completion, NSEC_PER_USEC);
|
|
|
|
/*
|
|
* If the request took rather long to complete, and, according
|
|
* to the maximum request size recorded, this completion latency
|
|
* implies that the request was certainly served at a very low
|
|
* rate (less than 1M sectors/sec), then the whole observation
|
|
* interval that lasts up to this time instant cannot be a
|
|
* valid time interval for computing a new peak rate. Invoke
|
|
* bfq_update_rate_reset to have the following three steps
|
|
* taken:
|
|
* - close the observation interval at the last (previous)
|
|
* request dispatch or completion
|
|
* - compute rate, if possible, for that observation interval
|
|
* - reset to zero samples, which will trigger a proper
|
|
* re-initialization of the observation interval on next
|
|
* dispatch
|
|
*/
|
|
if (delta_us > BFQ_MIN_TT/NSEC_PER_USEC &&
|
|
(bfqd->last_rq_max_size<<BFQ_RATE_SHIFT)/delta_us <
|
|
1UL<<(BFQ_RATE_SHIFT - 10))
|
|
bfq_update_rate_reset(bfqd, NULL);
|
|
bfqd->last_completion = now_ns;
|
|
/*
|
|
* Shared queues are likely to receive I/O at a high
|
|
* rate. This may deceptively let them be considered as wakers
|
|
* of other queues. But a false waker will unjustly steal
|
|
* bandwidth to its supposedly woken queue. So considering
|
|
* also shared queues in the waking mechanism may cause more
|
|
* control troubles than throughput benefits. Then do not set
|
|
* last_completed_rq_bfqq to bfqq if bfqq is a shared queue.
|
|
*/
|
|
if (!bfq_bfqq_coop(bfqq))
|
|
bfqd->last_completed_rq_bfqq = bfqq;
|
|
|
|
/*
|
|
* If we are waiting to discover whether the request pattern
|
|
* of the task associated with the queue is actually
|
|
* isochronous, and both requisites for this condition to hold
|
|
* are now satisfied, then compute soft_rt_next_start (see the
|
|
* comments on the function bfq_bfqq_softrt_next_start()). We
|
|
* do not compute soft_rt_next_start if bfqq is in interactive
|
|
* weight raising (see the comments in bfq_bfqq_expire() for
|
|
* an explanation). We schedule this delayed update when bfqq
|
|
* expires, if it still has in-flight requests.
|
|
*/
|
|
if (bfq_bfqq_softrt_update(bfqq) && bfqq->dispatched == 0 &&
|
|
RB_EMPTY_ROOT(&bfqq->sort_list) &&
|
|
bfqq->wr_coeff != bfqd->bfq_wr_coeff)
|
|
bfqq->soft_rt_next_start =
|
|
bfq_bfqq_softrt_next_start(bfqd, bfqq);
|
|
|
|
/*
|
|
* If this is the in-service queue, check if it needs to be expired,
|
|
* or if we want to idle in case it has no pending requests.
|
|
*/
|
|
if (bfqd->in_service_queue == bfqq) {
|
|
if (bfq_bfqq_must_idle(bfqq)) {
|
|
if (bfqq->dispatched == 0)
|
|
bfq_arm_slice_timer(bfqd);
|
|
/*
|
|
* If we get here, we do not expire bfqq, even
|
|
* if bfqq was in budget timeout or had no
|
|
* more requests (as controlled in the next
|
|
* conditional instructions). The reason for
|
|
* not expiring bfqq is as follows.
|
|
*
|
|
* Here bfqq->dispatched > 0 holds, but
|
|
* bfq_bfqq_must_idle() returned true. This
|
|
* implies that, even if no request arrives
|
|
* for bfqq before bfqq->dispatched reaches 0,
|
|
* bfqq will, however, not be expired on the
|
|
* completion event that causes bfqq->dispatch
|
|
* to reach zero. In contrast, on this event,
|
|
* bfqq will start enjoying device idling
|
|
* (I/O-dispatch plugging).
|
|
*
|
|
* But, if we expired bfqq here, bfqq would
|
|
* not have the chance to enjoy device idling
|
|
* when bfqq->dispatched finally reaches
|
|
* zero. This would expose bfqq to violation
|
|
* of its reserved service guarantees.
|
|
*/
|
|
return;
|
|
} else if (bfq_may_expire_for_budg_timeout(bfqq))
|
|
bfq_bfqq_expire(bfqd, bfqq, false,
|
|
BFQQE_BUDGET_TIMEOUT);
|
|
else if (RB_EMPTY_ROOT(&bfqq->sort_list) &&
|
|
(bfqq->dispatched == 0 ||
|
|
!bfq_better_to_idle(bfqq)))
|
|
bfq_bfqq_expire(bfqd, bfqq, false,
|
|
BFQQE_NO_MORE_REQUESTS);
|
|
}
|
|
|
|
if (!bfqd->rq_in_driver)
|
|
bfq_schedule_dispatch(bfqd);
|
|
}
|
|
|
|
static void bfq_finish_requeue_request_body(struct bfq_queue *bfqq)
|
|
{
|
|
bfqq->allocated--;
|
|
|
|
bfq_put_queue(bfqq);
|
|
}
|
|
|
|
/*
|
|
* The processes associated with bfqq may happen to generate their
|
|
* cumulative I/O at a lower rate than the rate at which the device
|
|
* could serve the same I/O. This is rather probable, e.g., if only
|
|
* one process is associated with bfqq and the device is an SSD. It
|
|
* results in bfqq becoming often empty while in service. In this
|
|
* respect, if BFQ is allowed to switch to another queue when bfqq
|
|
* remains empty, then the device goes on being fed with I/O requests,
|
|
* and the throughput is not affected. In contrast, if BFQ is not
|
|
* allowed to switch to another queue---because bfqq is sync and
|
|
* I/O-dispatch needs to be plugged while bfqq is temporarily
|
|
* empty---then, during the service of bfqq, there will be frequent
|
|
* "service holes", i.e., time intervals during which bfqq gets empty
|
|
* and the device can only consume the I/O already queued in its
|
|
* hardware queues. During service holes, the device may even get to
|
|
* remaining idle. In the end, during the service of bfqq, the device
|
|
* is driven at a lower speed than the one it can reach with the kind
|
|
* of I/O flowing through bfqq.
|
|
*
|
|
* To counter this loss of throughput, BFQ implements a "request
|
|
* injection mechanism", which tries to fill the above service holes
|
|
* with I/O requests taken from other queues. The hard part in this
|
|
* mechanism is finding the right amount of I/O to inject, so as to
|
|
* both boost throughput and not break bfqq's bandwidth and latency
|
|
* guarantees. In this respect, the mechanism maintains a per-queue
|
|
* inject limit, computed as below. While bfqq is empty, the injection
|
|
* mechanism dispatches extra I/O requests only until the total number
|
|
* of I/O requests in flight---i.e., already dispatched but not yet
|
|
* completed---remains lower than this limit.
|
|
*
|
|
* A first definition comes in handy to introduce the algorithm by
|
|
* which the inject limit is computed. We define as first request for
|
|
* bfqq, an I/O request for bfqq that arrives while bfqq is in
|
|
* service, and causes bfqq to switch from empty to non-empty. The
|
|
* algorithm updates the limit as a function of the effect of
|
|
* injection on the service times of only the first requests of
|
|
* bfqq. The reason for this restriction is that these are the
|
|
* requests whose service time is affected most, because they are the
|
|
* first to arrive after injection possibly occurred.
|
|
*
|
|
* To evaluate the effect of injection, the algorithm measures the
|
|
* "total service time" of first requests. We define as total service
|
|
* time of an I/O request, the time that elapses since when the
|
|
* request is enqueued into bfqq, to when it is completed. This
|
|
* quantity allows the whole effect of injection to be measured. It is
|
|
* easy to see why. Suppose that some requests of other queues are
|
|
* actually injected while bfqq is empty, and that a new request R
|
|
* then arrives for bfqq. If the device does start to serve all or
|
|
* part of the injected requests during the service hole, then,
|
|
* because of this extra service, it may delay the next invocation of
|
|
* the dispatch hook of BFQ. Then, even after R gets eventually
|
|
* dispatched, the device may delay the actual service of R if it is
|
|
* still busy serving the extra requests, or if it decides to serve,
|
|
* before R, some extra request still present in its queues. As a
|
|
* conclusion, the cumulative extra delay caused by injection can be
|
|
* easily evaluated by just comparing the total service time of first
|
|
* requests with and without injection.
|
|
*
|
|
* The limit-update algorithm works as follows. On the arrival of a
|
|
* first request of bfqq, the algorithm measures the total time of the
|
|
* request only if one of the three cases below holds, and, for each
|
|
* case, it updates the limit as described below:
|
|
*
|
|
* (1) If there is no in-flight request. This gives a baseline for the
|
|
* total service time of the requests of bfqq. If the baseline has
|
|
* not been computed yet, then, after computing it, the limit is
|
|
* set to 1, to start boosting throughput, and to prepare the
|
|
* ground for the next case. If the baseline has already been
|
|
* computed, then it is updated, in case it results to be lower
|
|
* than the previous value.
|
|
*
|
|
* (2) If the limit is higher than 0 and there are in-flight
|
|
* requests. By comparing the total service time in this case with
|
|
* the above baseline, it is possible to know at which extent the
|
|
* current value of the limit is inflating the total service
|
|
* time. If the inflation is below a certain threshold, then bfqq
|
|
* is assumed to be suffering from no perceivable loss of its
|
|
* service guarantees, and the limit is even tentatively
|
|
* increased. If the inflation is above the threshold, then the
|
|
* limit is decreased. Due to the lack of any hysteresis, this
|
|
* logic makes the limit oscillate even in steady workload
|
|
* conditions. Yet we opted for it, because it is fast in reaching
|
|
* the best value for the limit, as a function of the current I/O
|
|
* workload. To reduce oscillations, this step is disabled for a
|
|
* short time interval after the limit happens to be decreased.
|
|
*
|
|
* (3) Periodically, after resetting the limit, to make sure that the
|
|
* limit eventually drops in case the workload changes. This is
|
|
* needed because, after the limit has gone safely up for a
|
|
* certain workload, it is impossible to guess whether the
|
|
* baseline total service time may have changed, without measuring
|
|
* it again without injection. A more effective version of this
|
|
* step might be to just sample the baseline, by interrupting
|
|
* injection only once, and then to reset/lower the limit only if
|
|
* the total service time with the current limit does happen to be
|
|
* too large.
|
|
*
|
|
* More details on each step are provided in the comments on the
|
|
* pieces of code that implement these steps: the branch handling the
|
|
* transition from empty to non empty in bfq_add_request(), the branch
|
|
* handling injection in bfq_select_queue(), and the function
|
|
* bfq_choose_bfqq_for_injection(). These comments also explain some
|
|
* exceptions, made by the injection mechanism in some special cases.
|
|
*/
|
|
static void bfq_update_inject_limit(struct bfq_data *bfqd,
|
|
struct bfq_queue *bfqq)
|
|
{
|
|
u64 tot_time_ns = ktime_get_ns() - bfqd->last_empty_occupied_ns;
|
|
unsigned int old_limit = bfqq->inject_limit;
|
|
|
|
if (bfqq->last_serv_time_ns > 0 && bfqd->rqs_injected) {
|
|
u64 threshold = (bfqq->last_serv_time_ns * 3)>>1;
|
|
|
|
if (tot_time_ns >= threshold && old_limit > 0) {
|
|
bfqq->inject_limit--;
|
|
bfqq->decrease_time_jif = jiffies;
|
|
} else if (tot_time_ns < threshold &&
|
|
old_limit <= bfqd->max_rq_in_driver)
|
|
bfqq->inject_limit++;
|
|
}
|
|
|
|
/*
|
|
* Either we still have to compute the base value for the
|
|
* total service time, and there seem to be the right
|
|
* conditions to do it, or we can lower the last base value
|
|
* computed.
|
|
*
|
|
* NOTE: (bfqd->rq_in_driver == 1) means that there is no I/O
|
|
* request in flight, because this function is in the code
|
|
* path that handles the completion of a request of bfqq, and,
|
|
* in particular, this function is executed before
|
|
* bfqd->rq_in_driver is decremented in such a code path.
|
|
*/
|
|
if ((bfqq->last_serv_time_ns == 0 && bfqd->rq_in_driver == 1) ||
|
|
tot_time_ns < bfqq->last_serv_time_ns) {
|
|
if (bfqq->last_serv_time_ns == 0) {
|
|
/*
|
|
* Now we certainly have a base value: make sure we
|
|
* start trying injection.
|
|
*/
|
|
bfqq->inject_limit = max_t(unsigned int, 1, old_limit);
|
|
}
|
|
bfqq->last_serv_time_ns = tot_time_ns;
|
|
} else if (!bfqd->rqs_injected && bfqd->rq_in_driver == 1)
|
|
/*
|
|
* No I/O injected and no request still in service in
|
|
* the drive: these are the exact conditions for
|
|
* computing the base value of the total service time
|
|
* for bfqq. So let's update this value, because it is
|
|
* rather variable. For example, it varies if the size
|
|
* or the spatial locality of the I/O requests in bfqq
|
|
* change.
|
|
*/
|
|
bfqq->last_serv_time_ns = tot_time_ns;
|
|
|
|
|
|
/* update complete, not waiting for any request completion any longer */
|
|
bfqd->waited_rq = NULL;
|
|
bfqd->rqs_injected = false;
|
|
}
|
|
|
|
/*
|
|
* Handle either a requeue or a finish for rq. The things to do are
|
|
* the same in both cases: all references to rq are to be dropped. In
|
|
* particular, rq is considered completed from the point of view of
|
|
* the scheduler.
|
|
*/
|
|
static void bfq_finish_requeue_request(struct request *rq)
|
|
{
|
|
struct bfq_queue *bfqq = RQ_BFQQ(rq);
|
|
struct bfq_data *bfqd;
|
|
|
|
/*
|
|
* rq either is not associated with any icq, or is an already
|
|
* requeued request that has not (yet) been re-inserted into
|
|
* a bfq_queue.
|
|
*/
|
|
if (!rq->elv.icq || !bfqq)
|
|
return;
|
|
|
|
bfqd = bfqq->bfqd;
|
|
|
|
if (rq->rq_flags & RQF_STARTED)
|
|
bfqg_stats_update_completion(bfqq_group(bfqq),
|
|
rq->start_time_ns,
|
|
rq->io_start_time_ns,
|
|
rq->cmd_flags);
|
|
|
|
if (likely(rq->rq_flags & RQF_STARTED)) {
|
|
unsigned long flags;
|
|
|
|
spin_lock_irqsave(&bfqd->lock, flags);
|
|
|
|
if (rq == bfqd->waited_rq)
|
|
bfq_update_inject_limit(bfqd, bfqq);
|
|
|
|
bfq_completed_request(bfqq, bfqd);
|
|
bfq_finish_requeue_request_body(bfqq);
|
|
|
|
spin_unlock_irqrestore(&bfqd->lock, flags);
|
|
} else {
|
|
/*
|
|
* Request rq may be still/already in the scheduler,
|
|
* in which case we need to remove it (this should
|
|
* never happen in case of requeue). And we cannot
|
|
* defer such a check and removal, to avoid
|
|
* inconsistencies in the time interval from the end
|
|
* of this function to the start of the deferred work.
|
|
* This situation seems to occur only in process
|
|
* context, as a consequence of a merge. In the
|
|
* current version of the code, this implies that the
|
|
* lock is held.
|
|
*/
|
|
|
|
if (!RB_EMPTY_NODE(&rq->rb_node)) {
|
|
bfq_remove_request(rq->q, rq);
|
|
bfqg_stats_update_io_remove(bfqq_group(bfqq),
|
|
rq->cmd_flags);
|
|
}
|
|
bfq_finish_requeue_request_body(bfqq);
|
|
}
|
|
|
|
/*
|
|
* Reset private fields. In case of a requeue, this allows
|
|
* this function to correctly do nothing if it is spuriously
|
|
* invoked again on this same request (see the check at the
|
|
* beginning of the function). Probably, a better general
|
|
* design would be to prevent blk-mq from invoking the requeue
|
|
* or finish hooks of an elevator, for a request that is not
|
|
* referred by that elevator.
|
|
*
|
|
* Resetting the following fields would break the
|
|
* request-insertion logic if rq is re-inserted into a bfq
|
|
* internal queue, without a re-preparation. Here we assume
|
|
* that re-insertions of requeued requests, without
|
|
* re-preparation, can happen only for pass_through or at_head
|
|
* requests (which are not re-inserted into bfq internal
|
|
* queues).
|
|
*/
|
|
rq->elv.priv[0] = NULL;
|
|
rq->elv.priv[1] = NULL;
|
|
}
|
|
|
|
/*
|
|
* Removes the association between the current task and bfqq, assuming
|
|
* that bic points to the bfq iocontext of the task.
|
|
* Returns NULL if a new bfqq should be allocated, or the old bfqq if this
|
|
* was the last process referring to that bfqq.
|
|
*/
|
|
static struct bfq_queue *
|
|
bfq_split_bfqq(struct bfq_io_cq *bic, struct bfq_queue *bfqq)
|
|
{
|
|
bfq_log_bfqq(bfqq->bfqd, bfqq, "splitting queue");
|
|
|
|
if (bfqq_process_refs(bfqq) == 1) {
|
|
bfqq->pid = current->pid;
|
|
bfq_clear_bfqq_coop(bfqq);
|
|
bfq_clear_bfqq_split_coop(bfqq);
|
|
return bfqq;
|
|
}
|
|
|
|
bic_set_bfqq(bic, NULL, 1);
|
|
|
|
bfq_put_cooperator(bfqq);
|
|
|
|
bfq_release_process_ref(bfqq->bfqd, bfqq);
|
|
return NULL;
|
|
}
|
|
|
|
static struct bfq_queue *bfq_get_bfqq_handle_split(struct bfq_data *bfqd,
|
|
struct bfq_io_cq *bic,
|
|
struct bio *bio,
|
|
bool split, bool is_sync,
|
|
bool *new_queue)
|
|
{
|
|
struct bfq_queue *bfqq = bic_to_bfqq(bic, is_sync);
|
|
|
|
if (likely(bfqq && bfqq != &bfqd->oom_bfqq))
|
|
return bfqq;
|
|
|
|
if (new_queue)
|
|
*new_queue = true;
|
|
|
|
if (bfqq)
|
|
bfq_put_queue(bfqq);
|
|
bfqq = bfq_get_queue(bfqd, bio, is_sync, bic, split);
|
|
|
|
bic_set_bfqq(bic, bfqq, is_sync);
|
|
if (split && is_sync) {
|
|
if ((bic->was_in_burst_list && bfqd->large_burst) ||
|
|
bic->saved_in_large_burst)
|
|
bfq_mark_bfqq_in_large_burst(bfqq);
|
|
else {
|
|
bfq_clear_bfqq_in_large_burst(bfqq);
|
|
if (bic->was_in_burst_list)
|
|
/*
|
|
* If bfqq was in the current
|
|
* burst list before being
|
|
* merged, then we have to add
|
|
* it back. And we do not need
|
|
* to increase burst_size, as
|
|
* we did not decrement
|
|
* burst_size when we removed
|
|
* bfqq from the burst list as
|
|
* a consequence of a merge
|
|
* (see comments in
|
|
* bfq_put_queue). In this
|
|
* respect, it would be rather
|
|
* costly to know whether the
|
|
* current burst list is still
|
|
* the same burst list from
|
|
* which bfqq was removed on
|
|
* the merge. To avoid this
|
|
* cost, if bfqq was in a
|
|
* burst list, then we add
|
|
* bfqq to the current burst
|
|
* list without any further
|
|
* check. This can cause
|
|
* inappropriate insertions,
|
|
* but rarely enough to not
|
|
* harm the detection of large
|
|
* bursts significantly.
|
|
*/
|
|
hlist_add_head(&bfqq->burst_list_node,
|
|
&bfqd->burst_list);
|
|
}
|
|
bfqq->split_time = jiffies;
|
|
}
|
|
|
|
return bfqq;
|
|
}
|
|
|
|
/*
|
|
* Only reset private fields. The actual request preparation will be
|
|
* performed by bfq_init_rq, when rq is either inserted or merged. See
|
|
* comments on bfq_init_rq for the reason behind this delayed
|
|
* preparation.
|
|
*/
|
|
static void bfq_prepare_request(struct request *rq)
|
|
{
|
|
/*
|
|
* Regardless of whether we have an icq attached, we have to
|
|
* clear the scheduler pointers, as they might point to
|
|
* previously allocated bic/bfqq structs.
|
|
*/
|
|
rq->elv.priv[0] = rq->elv.priv[1] = NULL;
|
|
}
|
|
|
|
/*
|
|
* If needed, init rq, allocate bfq data structures associated with
|
|
* rq, and increment reference counters in the destination bfq_queue
|
|
* for rq. Return the destination bfq_queue for rq, or NULL is rq is
|
|
* not associated with any bfq_queue.
|
|
*
|
|
* This function is invoked by the functions that perform rq insertion
|
|
* or merging. One may have expected the above preparation operations
|
|
* to be performed in bfq_prepare_request, and not delayed to when rq
|
|
* is inserted or merged. The rationale behind this delayed
|
|
* preparation is that, after the prepare_request hook is invoked for
|
|
* rq, rq may still be transformed into a request with no icq, i.e., a
|
|
* request not associated with any queue. No bfq hook is invoked to
|
|
* signal this transformation. As a consequence, should these
|
|
* preparation operations be performed when the prepare_request hook
|
|
* is invoked, and should rq be transformed one moment later, bfq
|
|
* would end up in an inconsistent state, because it would have
|
|
* incremented some queue counters for an rq destined to
|
|
* transformation, without any chance to correctly lower these
|
|
* counters back. In contrast, no transformation can still happen for
|
|
* rq after rq has been inserted or merged. So, it is safe to execute
|
|
* these preparation operations when rq is finally inserted or merged.
|
|
*/
|
|
static struct bfq_queue *bfq_init_rq(struct request *rq)
|
|
{
|
|
struct request_queue *q = rq->q;
|
|
struct bio *bio = rq->bio;
|
|
struct bfq_data *bfqd = q->elevator->elevator_data;
|
|
struct bfq_io_cq *bic;
|
|
const int is_sync = rq_is_sync(rq);
|
|
struct bfq_queue *bfqq;
|
|
bool new_queue = false;
|
|
bool bfqq_already_existing = false, split = false;
|
|
|
|
if (unlikely(!rq->elv.icq))
|
|
return NULL;
|
|
|
|
/*
|
|
* Assuming that elv.priv[1] is set only if everything is set
|
|
* for this rq. This holds true, because this function is
|
|
* invoked only for insertion or merging, and, after such
|
|
* events, a request cannot be manipulated any longer before
|
|
* being removed from bfq.
|
|
*/
|
|
if (rq->elv.priv[1])
|
|
return rq->elv.priv[1];
|
|
|
|
bic = icq_to_bic(rq->elv.icq);
|
|
|
|
bfq_check_ioprio_change(bic, bio);
|
|
|
|
bfq_bic_update_cgroup(bic, bio);
|
|
|
|
bfqq = bfq_get_bfqq_handle_split(bfqd, bic, bio, false, is_sync,
|
|
&new_queue);
|
|
|
|
if (likely(!new_queue)) {
|
|
/* If the queue was seeky for too long, break it apart. */
|
|
if (bfq_bfqq_coop(bfqq) && bfq_bfqq_split_coop(bfqq) &&
|
|
!bic->stably_merged) {
|
|
struct bfq_queue *old_bfqq = bfqq;
|
|
|
|
/* Update bic before losing reference to bfqq */
|
|
if (bfq_bfqq_in_large_burst(bfqq))
|
|
bic->saved_in_large_burst = true;
|
|
|
|
bfqq = bfq_split_bfqq(bic, bfqq);
|
|
split = true;
|
|
|
|
if (!bfqq) {
|
|
bfqq = bfq_get_bfqq_handle_split(bfqd, bic, bio,
|
|
true, is_sync,
|
|
NULL);
|
|
bfqq->waker_bfqq = old_bfqq->waker_bfqq;
|
|
bfqq->tentative_waker_bfqq = NULL;
|
|
|
|
/*
|
|
* If the waker queue disappears, then
|
|
* new_bfqq->waker_bfqq must be
|
|
* reset. So insert new_bfqq into the
|
|
* woken_list of the waker. See
|
|
* bfq_check_waker for details.
|
|
*/
|
|
if (bfqq->waker_bfqq)
|
|
hlist_add_head(&bfqq->woken_list_node,
|
|
&bfqq->waker_bfqq->woken_list);
|
|
} else
|
|
bfqq_already_existing = true;
|
|
}
|
|
}
|
|
|
|
bfqq->allocated++;
|
|
bfqq->ref++;
|
|
bfq_log_bfqq(bfqd, bfqq, "get_request %p: bfqq %p, %d",
|
|
rq, bfqq, bfqq->ref);
|
|
|
|
rq->elv.priv[0] = bic;
|
|
rq->elv.priv[1] = bfqq;
|
|
|
|
/*
|
|
* If a bfq_queue has only one process reference, it is owned
|
|
* by only this bic: we can then set bfqq->bic = bic. in
|
|
* addition, if the queue has also just been split, we have to
|
|
* resume its state.
|
|
*/
|
|
if (likely(bfqq != &bfqd->oom_bfqq) && bfqq_process_refs(bfqq) == 1) {
|
|
bfqq->bic = bic;
|
|
if (split) {
|
|
/*
|
|
* The queue has just been split from a shared
|
|
* queue: restore the idle window and the
|
|
* possible weight raising period.
|
|
*/
|
|
bfq_bfqq_resume_state(bfqq, bfqd, bic,
|
|
bfqq_already_existing);
|
|
}
|
|
}
|
|
|
|
/*
|
|
* Consider bfqq as possibly belonging to a burst of newly
|
|
* created queues only if:
|
|
* 1) A burst is actually happening (bfqd->burst_size > 0)
|
|
* or
|
|
* 2) There is no other active queue. In fact, if, in
|
|
* contrast, there are active queues not belonging to the
|
|
* possible burst bfqq may belong to, then there is no gain
|
|
* in considering bfqq as belonging to a burst, and
|
|
* therefore in not weight-raising bfqq. See comments on
|
|
* bfq_handle_burst().
|
|
*
|
|
* This filtering also helps eliminating false positives,
|
|
* occurring when bfqq does not belong to an actual large
|
|
* burst, but some background task (e.g., a service) happens
|
|
* to trigger the creation of new queues very close to when
|
|
* bfqq and its possible companion queues are created. See
|
|
* comments on bfq_handle_burst() for further details also on
|
|
* this issue.
|
|
*/
|
|
if (unlikely(bfq_bfqq_just_created(bfqq) &&
|
|
(bfqd->burst_size > 0 ||
|
|
bfq_tot_busy_queues(bfqd) == 0)))
|
|
bfq_handle_burst(bfqd, bfqq);
|
|
|
|
return bfqq;
|
|
}
|
|
|
|
static void
|
|
bfq_idle_slice_timer_body(struct bfq_data *bfqd, struct bfq_queue *bfqq)
|
|
{
|
|
enum bfqq_expiration reason;
|
|
unsigned long flags;
|
|
|
|
spin_lock_irqsave(&bfqd->lock, flags);
|
|
|
|
/*
|
|
* Considering that bfqq may be in race, we should firstly check
|
|
* whether bfqq is in service before doing something on it. If
|
|
* the bfqq in race is not in service, it has already been expired
|
|
* through __bfq_bfqq_expire func and its wait_request flags has
|
|
* been cleared in __bfq_bfqd_reset_in_service func.
|
|
*/
|
|
if (bfqq != bfqd->in_service_queue) {
|
|
spin_unlock_irqrestore(&bfqd->lock, flags);
|
|
return;
|
|
}
|
|
|
|
bfq_clear_bfqq_wait_request(bfqq);
|
|
|
|
if (bfq_bfqq_budget_timeout(bfqq))
|
|
/*
|
|
* Also here the queue can be safely expired
|
|
* for budget timeout without wasting
|
|
* guarantees
|
|
*/
|
|
reason = BFQQE_BUDGET_TIMEOUT;
|
|
else if (bfqq->queued[0] == 0 && bfqq->queued[1] == 0)
|
|
/*
|
|
* The queue may not be empty upon timer expiration,
|
|
* because we may not disable the timer when the
|
|
* first request of the in-service queue arrives
|
|
* during disk idling.
|
|
*/
|
|
reason = BFQQE_TOO_IDLE;
|
|
else
|
|
goto schedule_dispatch;
|
|
|
|
bfq_bfqq_expire(bfqd, bfqq, true, reason);
|
|
|
|
schedule_dispatch:
|
|
spin_unlock_irqrestore(&bfqd->lock, flags);
|
|
bfq_schedule_dispatch(bfqd);
|
|
}
|
|
|
|
/*
|
|
* Handler of the expiration of the timer running if the in-service queue
|
|
* is idling inside its time slice.
|
|
*/
|
|
static enum hrtimer_restart bfq_idle_slice_timer(struct hrtimer *timer)
|
|
{
|
|
struct bfq_data *bfqd = container_of(timer, struct bfq_data,
|
|
idle_slice_timer);
|
|
struct bfq_queue *bfqq = bfqd->in_service_queue;
|
|
|
|
/*
|
|
* Theoretical race here: the in-service queue can be NULL or
|
|
* different from the queue that was idling if a new request
|
|
* arrives for the current queue and there is a full dispatch
|
|
* cycle that changes the in-service queue. This can hardly
|
|
* happen, but in the worst case we just expire a queue too
|
|
* early.
|
|
*/
|
|
if (bfqq)
|
|
bfq_idle_slice_timer_body(bfqd, bfqq);
|
|
|
|
return HRTIMER_NORESTART;
|
|
}
|
|
|
|
static void __bfq_put_async_bfqq(struct bfq_data *bfqd,
|
|
struct bfq_queue **bfqq_ptr)
|
|
{
|
|
struct bfq_queue *bfqq = *bfqq_ptr;
|
|
|
|
bfq_log(bfqd, "put_async_bfqq: %p", bfqq);
|
|
if (bfqq) {
|
|
bfq_bfqq_move(bfqd, bfqq, bfqd->root_group);
|
|
|
|
bfq_log_bfqq(bfqd, bfqq, "put_async_bfqq: putting %p, %d",
|
|
bfqq, bfqq->ref);
|
|
bfq_put_queue(bfqq);
|
|
*bfqq_ptr = NULL;
|
|
}
|
|
}
|
|
|
|
/*
|
|
* Release all the bfqg references to its async queues. If we are
|
|
* deallocating the group these queues may still contain requests, so
|
|
* we reparent them to the root cgroup (i.e., the only one that will
|
|
* exist for sure until all the requests on a device are gone).
|
|
*/
|
|
void bfq_put_async_queues(struct bfq_data *bfqd, struct bfq_group *bfqg)
|
|
{
|
|
int i, j;
|
|
|
|
for (i = 0; i < 2; i++)
|
|
for (j = 0; j < IOPRIO_BE_NR; j++)
|
|
__bfq_put_async_bfqq(bfqd, &bfqg->async_bfqq[i][j]);
|
|
|
|
__bfq_put_async_bfqq(bfqd, &bfqg->async_idle_bfqq);
|
|
}
|
|
|
|
/*
|
|
* See the comments on bfq_limit_depth for the purpose of
|
|
* the depths set in the function. Return minimum shallow depth we'll use.
|
|
*/
|
|
static unsigned int bfq_update_depths(struct bfq_data *bfqd,
|
|
struct sbitmap_queue *bt)
|
|
{
|
|
unsigned int i, j, min_shallow = UINT_MAX;
|
|
|
|
/*
|
|
* In-word depths if no bfq_queue is being weight-raised:
|
|
* leaving 25% of tags only for sync reads.
|
|
*
|
|
* In next formulas, right-shift the value
|
|
* (1U<<bt->sb.shift), instead of computing directly
|
|
* (1U<<(bt->sb.shift - something)), to be robust against
|
|
* any possible value of bt->sb.shift, without having to
|
|
* limit 'something'.
|
|
*/
|
|
/* no more than 50% of tags for async I/O */
|
|
bfqd->word_depths[0][0] = max((1U << bt->sb.shift) >> 1, 1U);
|
|
/*
|
|
* no more than 75% of tags for sync writes (25% extra tags
|
|
* w.r.t. async I/O, to prevent async I/O from starving sync
|
|
* writes)
|
|
*/
|
|
bfqd->word_depths[0][1] = max(((1U << bt->sb.shift) * 3) >> 2, 1U);
|
|
|
|
/*
|
|
* In-word depths in case some bfq_queue is being weight-
|
|
* raised: leaving ~63% of tags for sync reads. This is the
|
|
* highest percentage for which, in our tests, application
|
|
* start-up times didn't suffer from any regression due to tag
|
|
* shortage.
|
|
*/
|
|
/* no more than ~18% of tags for async I/O */
|
|
bfqd->word_depths[1][0] = max(((1U << bt->sb.shift) * 3) >> 4, 1U);
|
|
/* no more than ~37% of tags for sync writes (~20% extra tags) */
|
|
bfqd->word_depths[1][1] = max(((1U << bt->sb.shift) * 6) >> 4, 1U);
|
|
|
|
for (i = 0; i < 2; i++)
|
|
for (j = 0; j < 2; j++)
|
|
min_shallow = min(min_shallow, bfqd->word_depths[i][j]);
|
|
|
|
return min_shallow;
|
|
}
|
|
|
|
static void bfq_depth_updated(struct blk_mq_hw_ctx *hctx)
|
|
{
|
|
struct bfq_data *bfqd = hctx->queue->elevator->elevator_data;
|
|
struct blk_mq_tags *tags = hctx->sched_tags;
|
|
unsigned int min_shallow;
|
|
|
|
min_shallow = bfq_update_depths(bfqd, tags->bitmap_tags);
|
|
sbitmap_queue_min_shallow_depth(tags->bitmap_tags, min_shallow);
|
|
}
|
|
|
|
static int bfq_init_hctx(struct blk_mq_hw_ctx *hctx, unsigned int index)
|
|
{
|
|
bfq_depth_updated(hctx);
|
|
return 0;
|
|
}
|
|
|
|
static void bfq_exit_queue(struct elevator_queue *e)
|
|
{
|
|
struct bfq_data *bfqd = e->elevator_data;
|
|
struct bfq_queue *bfqq, *n;
|
|
|
|
hrtimer_cancel(&bfqd->idle_slice_timer);
|
|
|
|
spin_lock_irq(&bfqd->lock);
|
|
list_for_each_entry_safe(bfqq, n, &bfqd->idle_list, bfqq_list)
|
|
bfq_deactivate_bfqq(bfqd, bfqq, false, false);
|
|
spin_unlock_irq(&bfqd->lock);
|
|
|
|
hrtimer_cancel(&bfqd->idle_slice_timer);
|
|
|
|
/* release oom-queue reference to root group */
|
|
bfqg_and_blkg_put(bfqd->root_group);
|
|
|
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED
|
|
blkcg_deactivate_policy(bfqd->queue, &blkcg_policy_bfq);
|
|
#else
|
|
spin_lock_irq(&bfqd->lock);
|
|
bfq_put_async_queues(bfqd, bfqd->root_group);
|
|
kfree(bfqd->root_group);
|
|
spin_unlock_irq(&bfqd->lock);
|
|
#endif
|
|
|
|
kfree(bfqd);
|
|
}
|
|
|
|
static void bfq_init_root_group(struct bfq_group *root_group,
|
|
struct bfq_data *bfqd)
|
|
{
|
|
int i;
|
|
|
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED
|
|
root_group->entity.parent = NULL;
|
|
root_group->my_entity = NULL;
|
|
root_group->bfqd = bfqd;
|
|
#endif
|
|
root_group->rq_pos_tree = RB_ROOT;
|
|
for (i = 0; i < BFQ_IOPRIO_CLASSES; i++)
|
|
root_group->sched_data.service_tree[i] = BFQ_SERVICE_TREE_INIT;
|
|
root_group->sched_data.bfq_class_idle_last_service = jiffies;
|
|
}
|
|
|
|
static int bfq_init_queue(struct request_queue *q, struct elevator_type *e)
|
|
{
|
|
struct bfq_data *bfqd;
|
|
struct elevator_queue *eq;
|
|
|
|
eq = elevator_alloc(q, e);
|
|
if (!eq)
|
|
return -ENOMEM;
|
|
|
|
bfqd = kzalloc_node(sizeof(*bfqd), GFP_KERNEL, q->node);
|
|
if (!bfqd) {
|
|
kobject_put(&eq->kobj);
|
|
return -ENOMEM;
|
|
}
|
|
eq->elevator_data = bfqd;
|
|
|
|
spin_lock_irq(&q->queue_lock);
|
|
q->elevator = eq;
|
|
spin_unlock_irq(&q->queue_lock);
|
|
|
|
/*
|
|
* Our fallback bfqq if bfq_find_alloc_queue() runs into OOM issues.
|
|
* Grab a permanent reference to it, so that the normal code flow
|
|
* will not attempt to free it.
|
|
*/
|
|
bfq_init_bfqq(bfqd, &bfqd->oom_bfqq, NULL, 1, 0);
|
|
bfqd->oom_bfqq.ref++;
|
|
bfqd->oom_bfqq.new_ioprio = BFQ_DEFAULT_QUEUE_IOPRIO;
|
|
bfqd->oom_bfqq.new_ioprio_class = IOPRIO_CLASS_BE;
|
|
bfqd->oom_bfqq.entity.new_weight =
|
|
bfq_ioprio_to_weight(bfqd->oom_bfqq.new_ioprio);
|
|
|
|
/* oom_bfqq does not participate to bursts */
|
|
bfq_clear_bfqq_just_created(&bfqd->oom_bfqq);
|
|
|
|
/*
|
|
* Trigger weight initialization, according to ioprio, at the
|
|
* oom_bfqq's first activation. The oom_bfqq's ioprio and ioprio
|
|
* class won't be changed any more.
|
|
*/
|
|
bfqd->oom_bfqq.entity.prio_changed = 1;
|
|
|
|
bfqd->queue = q;
|
|
|
|
INIT_LIST_HEAD(&bfqd->dispatch);
|
|
|
|
hrtimer_init(&bfqd->idle_slice_timer, CLOCK_MONOTONIC,
|
|
HRTIMER_MODE_REL);
|
|
bfqd->idle_slice_timer.function = bfq_idle_slice_timer;
|
|
|
|
bfqd->queue_weights_tree = RB_ROOT_CACHED;
|
|
bfqd->num_groups_with_pending_reqs = 0;
|
|
|
|
INIT_LIST_HEAD(&bfqd->active_list);
|
|
INIT_LIST_HEAD(&bfqd->idle_list);
|
|
INIT_HLIST_HEAD(&bfqd->burst_list);
|
|
|
|
bfqd->hw_tag = -1;
|
|
bfqd->nonrot_with_queueing = blk_queue_nonrot(bfqd->queue);
|
|
|
|
bfqd->bfq_max_budget = bfq_default_max_budget;
|
|
|
|
bfqd->bfq_fifo_expire[0] = bfq_fifo_expire[0];
|
|
bfqd->bfq_fifo_expire[1] = bfq_fifo_expire[1];
|
|
bfqd->bfq_back_max = bfq_back_max;
|
|
bfqd->bfq_back_penalty = bfq_back_penalty;
|
|
bfqd->bfq_slice_idle = bfq_slice_idle;
|
|
bfqd->bfq_timeout = bfq_timeout;
|
|
|
|
bfqd->bfq_large_burst_thresh = 8;
|
|
bfqd->bfq_burst_interval = msecs_to_jiffies(180);
|
|
|
|
bfqd->low_latency = true;
|
|
|
|
/*
|
|
* Trade-off between responsiveness and fairness.
|
|
*/
|
|
bfqd->bfq_wr_coeff = 30;
|
|
bfqd->bfq_wr_rt_max_time = msecs_to_jiffies(300);
|
|
bfqd->bfq_wr_max_time = 0;
|
|
bfqd->bfq_wr_min_idle_time = msecs_to_jiffies(2000);
|
|
bfqd->bfq_wr_min_inter_arr_async = msecs_to_jiffies(500);
|
|
bfqd->bfq_wr_max_softrt_rate = 7000; /*
|
|
* Approximate rate required
|
|
* to playback or record a
|
|
* high-definition compressed
|
|
* video.
|
|
*/
|
|
bfqd->wr_busy_queues = 0;
|
|
|
|
/*
|
|
* Begin by assuming, optimistically, that the device peak
|
|
* rate is equal to 2/3 of the highest reference rate.
|
|
*/
|
|
bfqd->rate_dur_prod = ref_rate[blk_queue_nonrot(bfqd->queue)] *
|
|
ref_wr_duration[blk_queue_nonrot(bfqd->queue)];
|
|
bfqd->peak_rate = ref_rate[blk_queue_nonrot(bfqd->queue)] * 2 / 3;
|
|
|
|
spin_lock_init(&bfqd->lock);
|
|
|
|
/*
|
|
* The invocation of the next bfq_create_group_hierarchy
|
|
* function is the head of a chain of function calls
|
|
* (bfq_create_group_hierarchy->blkcg_activate_policy->
|
|
* blk_mq_freeze_queue) that may lead to the invocation of the
|
|
* has_work hook function. For this reason,
|
|
* bfq_create_group_hierarchy is invoked only after all
|
|
* scheduler data has been initialized, apart from the fields
|
|
* that can be initialized only after invoking
|
|
* bfq_create_group_hierarchy. This, in particular, enables
|
|
* has_work to correctly return false. Of course, to avoid
|
|
* other inconsistencies, the blk-mq stack must then refrain
|
|
* from invoking further scheduler hooks before this init
|
|
* function is finished.
|
|
*/
|
|
bfqd->root_group = bfq_create_group_hierarchy(bfqd, q->node);
|
|
if (!bfqd->root_group)
|
|
goto out_free;
|
|
bfq_init_root_group(bfqd->root_group, bfqd);
|
|
bfq_init_entity(&bfqd->oom_bfqq.entity, bfqd->root_group);
|
|
|
|
wbt_disable_default(q);
|
|
return 0;
|
|
|
|
out_free:
|
|
kfree(bfqd);
|
|
kobject_put(&eq->kobj);
|
|
return -ENOMEM;
|
|
}
|
|
|
|
static void bfq_slab_kill(void)
|
|
{
|
|
kmem_cache_destroy(bfq_pool);
|
|
}
|
|
|
|
static int __init bfq_slab_setup(void)
|
|
{
|
|
bfq_pool = KMEM_CACHE(bfq_queue, 0);
|
|
if (!bfq_pool)
|
|
return -ENOMEM;
|
|
return 0;
|
|
}
|
|
|
|
static ssize_t bfq_var_show(unsigned int var, char *page)
|
|
{
|
|
return sprintf(page, "%u\n", var);
|
|
}
|
|
|
|
static int bfq_var_store(unsigned long *var, const char *page)
|
|
{
|
|
unsigned long new_val;
|
|
int ret = kstrtoul(page, 10, &new_val);
|
|
|
|
if (ret)
|
|
return ret;
|
|
*var = new_val;
|
|
return 0;
|
|
}
|
|
|
|
#define SHOW_FUNCTION(__FUNC, __VAR, __CONV) \
|
|
static ssize_t __FUNC(struct elevator_queue *e, char *page) \
|
|
{ \
|
|
struct bfq_data *bfqd = e->elevator_data; \
|
|
u64 __data = __VAR; \
|
|
if (__CONV == 1) \
|
|
__data = jiffies_to_msecs(__data); \
|
|
else if (__CONV == 2) \
|
|
__data = div_u64(__data, NSEC_PER_MSEC); \
|
|
return bfq_var_show(__data, (page)); \
|
|
}
|
|
SHOW_FUNCTION(bfq_fifo_expire_sync_show, bfqd->bfq_fifo_expire[1], 2);
|
|
SHOW_FUNCTION(bfq_fifo_expire_async_show, bfqd->bfq_fifo_expire[0], 2);
|
|
SHOW_FUNCTION(bfq_back_seek_max_show, bfqd->bfq_back_max, 0);
|
|
SHOW_FUNCTION(bfq_back_seek_penalty_show, bfqd->bfq_back_penalty, 0);
|
|
SHOW_FUNCTION(bfq_slice_idle_show, bfqd->bfq_slice_idle, 2);
|
|
SHOW_FUNCTION(bfq_max_budget_show, bfqd->bfq_user_max_budget, 0);
|
|
SHOW_FUNCTION(bfq_timeout_sync_show, bfqd->bfq_timeout, 1);
|
|
SHOW_FUNCTION(bfq_strict_guarantees_show, bfqd->strict_guarantees, 0);
|
|
SHOW_FUNCTION(bfq_low_latency_show, bfqd->low_latency, 0);
|
|
#undef SHOW_FUNCTION
|
|
|
|
#define USEC_SHOW_FUNCTION(__FUNC, __VAR) \
|
|
static ssize_t __FUNC(struct elevator_queue *e, char *page) \
|
|
{ \
|
|
struct bfq_data *bfqd = e->elevator_data; \
|
|
u64 __data = __VAR; \
|
|
__data = div_u64(__data, NSEC_PER_USEC); \
|
|
return bfq_var_show(__data, (page)); \
|
|
}
|
|
USEC_SHOW_FUNCTION(bfq_slice_idle_us_show, bfqd->bfq_slice_idle);
|
|
#undef USEC_SHOW_FUNCTION
|
|
|
|
#define STORE_FUNCTION(__FUNC, __PTR, MIN, MAX, __CONV) \
|
|
static ssize_t \
|
|
__FUNC(struct elevator_queue *e, const char *page, size_t count) \
|
|
{ \
|
|
struct bfq_data *bfqd = e->elevator_data; \
|
|
unsigned long __data, __min = (MIN), __max = (MAX); \
|
|
int ret; \
|
|
\
|
|
ret = bfq_var_store(&__data, (page)); \
|
|
if (ret) \
|
|
return ret; \
|
|
if (__data < __min) \
|
|
__data = __min; \
|
|
else if (__data > __max) \
|
|
__data = __max; \
|
|
if (__CONV == 1) \
|
|
*(__PTR) = msecs_to_jiffies(__data); \
|
|
else if (__CONV == 2) \
|
|
*(__PTR) = (u64)__data * NSEC_PER_MSEC; \
|
|
else \
|
|
*(__PTR) = __data; \
|
|
return count; \
|
|
}
|
|
STORE_FUNCTION(bfq_fifo_expire_sync_store, &bfqd->bfq_fifo_expire[1], 1,
|
|
INT_MAX, 2);
|
|
STORE_FUNCTION(bfq_fifo_expire_async_store, &bfqd->bfq_fifo_expire[0], 1,
|
|
INT_MAX, 2);
|
|
STORE_FUNCTION(bfq_back_seek_max_store, &bfqd->bfq_back_max, 0, INT_MAX, 0);
|
|
STORE_FUNCTION(bfq_back_seek_penalty_store, &bfqd->bfq_back_penalty, 1,
|
|
INT_MAX, 0);
|
|
STORE_FUNCTION(bfq_slice_idle_store, &bfqd->bfq_slice_idle, 0, INT_MAX, 2);
|
|
#undef STORE_FUNCTION
|
|
|
|
#define USEC_STORE_FUNCTION(__FUNC, __PTR, MIN, MAX) \
|
|
static ssize_t __FUNC(struct elevator_queue *e, const char *page, size_t count)\
|
|
{ \
|
|
struct bfq_data *bfqd = e->elevator_data; \
|
|
unsigned long __data, __min = (MIN), __max = (MAX); \
|
|
int ret; \
|
|
\
|
|
ret = bfq_var_store(&__data, (page)); \
|
|
if (ret) \
|
|
return ret; \
|
|
if (__data < __min) \
|
|
__data = __min; \
|
|
else if (__data > __max) \
|
|
__data = __max; \
|
|
*(__PTR) = (u64)__data * NSEC_PER_USEC; \
|
|
return count; \
|
|
}
|
|
USEC_STORE_FUNCTION(bfq_slice_idle_us_store, &bfqd->bfq_slice_idle, 0,
|
|
UINT_MAX);
|
|
#undef USEC_STORE_FUNCTION
|
|
|
|
static ssize_t bfq_max_budget_store(struct elevator_queue *e,
|
|
const char *page, size_t count)
|
|
{
|
|
struct bfq_data *bfqd = e->elevator_data;
|
|
unsigned long __data;
|
|
int ret;
|
|
|
|
ret = bfq_var_store(&__data, (page));
|
|
if (ret)
|
|
return ret;
|
|
|
|
if (__data == 0)
|
|
bfqd->bfq_max_budget = bfq_calc_max_budget(bfqd);
|
|
else {
|
|
if (__data > INT_MAX)
|
|
__data = INT_MAX;
|
|
bfqd->bfq_max_budget = __data;
|
|
}
|
|
|
|
bfqd->bfq_user_max_budget = __data;
|
|
|
|
return count;
|
|
}
|
|
|
|
/*
|
|
* Leaving this name to preserve name compatibility with cfq
|
|
* parameters, but this timeout is used for both sync and async.
|
|
*/
|
|
static ssize_t bfq_timeout_sync_store(struct elevator_queue *e,
|
|
const char *page, size_t count)
|
|
{
|
|
struct bfq_data *bfqd = e->elevator_data;
|
|
unsigned long __data;
|
|
int ret;
|
|
|
|
ret = bfq_var_store(&__data, (page));
|
|
if (ret)
|
|
return ret;
|
|
|
|
if (__data < 1)
|
|
__data = 1;
|
|
else if (__data > INT_MAX)
|
|
__data = INT_MAX;
|
|
|
|
bfqd->bfq_timeout = msecs_to_jiffies(__data);
|
|
if (bfqd->bfq_user_max_budget == 0)
|
|
bfqd->bfq_max_budget = bfq_calc_max_budget(bfqd);
|
|
|
|
return count;
|
|
}
|
|
|
|
static ssize_t bfq_strict_guarantees_store(struct elevator_queue *e,
|
|
const char *page, size_t count)
|
|
{
|
|
struct bfq_data *bfqd = e->elevator_data;
|
|
unsigned long __data;
|
|
int ret;
|
|
|
|
ret = bfq_var_store(&__data, (page));
|
|
if (ret)
|
|
return ret;
|
|
|
|
if (__data > 1)
|
|
__data = 1;
|
|
if (!bfqd->strict_guarantees && __data == 1
|
|
&& bfqd->bfq_slice_idle < 8 * NSEC_PER_MSEC)
|
|
bfqd->bfq_slice_idle = 8 * NSEC_PER_MSEC;
|
|
|
|
bfqd->strict_guarantees = __data;
|
|
|
|
return count;
|
|
}
|
|
|
|
static ssize_t bfq_low_latency_store(struct elevator_queue *e,
|
|
const char *page, size_t count)
|
|
{
|
|
struct bfq_data *bfqd = e->elevator_data;
|
|
unsigned long __data;
|
|
int ret;
|
|
|
|
ret = bfq_var_store(&__data, (page));
|
|
if (ret)
|
|
return ret;
|
|
|
|
if (__data > 1)
|
|
__data = 1;
|
|
if (__data == 0 && bfqd->low_latency != 0)
|
|
bfq_end_wr(bfqd);
|
|
bfqd->low_latency = __data;
|
|
|
|
return count;
|
|
}
|
|
|
|
#define BFQ_ATTR(name) \
|
|
__ATTR(name, 0644, bfq_##name##_show, bfq_##name##_store)
|
|
|
|
static struct elv_fs_entry bfq_attrs[] = {
|
|
BFQ_ATTR(fifo_expire_sync),
|
|
BFQ_ATTR(fifo_expire_async),
|
|
BFQ_ATTR(back_seek_max),
|
|
BFQ_ATTR(back_seek_penalty),
|
|
BFQ_ATTR(slice_idle),
|
|
BFQ_ATTR(slice_idle_us),
|
|
BFQ_ATTR(max_budget),
|
|
BFQ_ATTR(timeout_sync),
|
|
BFQ_ATTR(strict_guarantees),
|
|
BFQ_ATTR(low_latency),
|
|
__ATTR_NULL
|
|
};
|
|
|
|
static struct elevator_type iosched_bfq_mq = {
|
|
.ops = {
|
|
.limit_depth = bfq_limit_depth,
|
|
.prepare_request = bfq_prepare_request,
|
|
.requeue_request = bfq_finish_requeue_request,
|
|
.finish_request = bfq_finish_requeue_request,
|
|
.exit_icq = bfq_exit_icq,
|
|
.insert_requests = bfq_insert_requests,
|
|
.dispatch_request = bfq_dispatch_request,
|
|
.next_request = elv_rb_latter_request,
|
|
.former_request = elv_rb_former_request,
|
|
.allow_merge = bfq_allow_bio_merge,
|
|
.bio_merge = bfq_bio_merge,
|
|
.request_merge = bfq_request_merge,
|
|
.requests_merged = bfq_requests_merged,
|
|
.request_merged = bfq_request_merged,
|
|
.has_work = bfq_has_work,
|
|
.depth_updated = bfq_depth_updated,
|
|
.init_hctx = bfq_init_hctx,
|
|
.init_sched = bfq_init_queue,
|
|
.exit_sched = bfq_exit_queue,
|
|
},
|
|
|
|
.icq_size = sizeof(struct bfq_io_cq),
|
|
.icq_align = __alignof__(struct bfq_io_cq),
|
|
.elevator_attrs = bfq_attrs,
|
|
.elevator_name = "bfq",
|
|
.elevator_owner = THIS_MODULE,
|
|
};
|
|
MODULE_ALIAS("bfq-iosched");
|
|
|
|
static int __init bfq_init(void)
|
|
{
|
|
int ret;
|
|
|
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED
|
|
ret = blkcg_policy_register(&blkcg_policy_bfq);
|
|
if (ret)
|
|
return ret;
|
|
#endif
|
|
|
|
ret = -ENOMEM;
|
|
if (bfq_slab_setup())
|
|
goto err_pol_unreg;
|
|
|
|
/*
|
|
* Times to load large popular applications for the typical
|
|
* systems installed on the reference devices (see the
|
|
* comments before the definition of the next
|
|
* array). Actually, we use slightly lower values, as the
|
|
* estimated peak rate tends to be smaller than the actual
|
|
* peak rate. The reason for this last fact is that estimates
|
|
* are computed over much shorter time intervals than the long
|
|
* intervals typically used for benchmarking. Why? First, to
|
|
* adapt more quickly to variations. Second, because an I/O
|
|
* scheduler cannot rely on a peak-rate-evaluation workload to
|
|
* be run for a long time.
|
|
*/
|
|
ref_wr_duration[0] = msecs_to_jiffies(7000); /* actually 8 sec */
|
|
ref_wr_duration[1] = msecs_to_jiffies(2500); /* actually 3 sec */
|
|
|
|
ret = elv_register(&iosched_bfq_mq);
|
|
if (ret)
|
|
goto slab_kill;
|
|
|
|
return 0;
|
|
|
|
slab_kill:
|
|
bfq_slab_kill();
|
|
err_pol_unreg:
|
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED
|
|
blkcg_policy_unregister(&blkcg_policy_bfq);
|
|
#endif
|
|
return ret;
|
|
}
|
|
|
|
static void __exit bfq_exit(void)
|
|
{
|
|
elv_unregister(&iosched_bfq_mq);
|
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED
|
|
blkcg_policy_unregister(&blkcg_policy_bfq);
|
|
#endif
|
|
bfq_slab_kill();
|
|
}
|
|
|
|
module_init(bfq_init);
|
|
module_exit(bfq_exit);
|
|
|
|
MODULE_AUTHOR("Paolo Valente");
|
|
MODULE_LICENSE("GPL");
|
|
MODULE_DESCRIPTION("MQ Budget Fair Queueing I/O Scheduler");
|