linux/fs/xfs/xfs_trans.h

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// SPDX-License-Identifier: GPL-2.0
/*
* Copyright (c) 2000-2002,2005 Silicon Graphics, Inc.
* All Rights Reserved.
*/
#ifndef __XFS_TRANS_H__
#define __XFS_TRANS_H__
/* kernel only transaction subsystem defines */
struct xlog;
struct xfs_buf;
struct xfs_buftarg;
struct xfs_efd_log_item;
struct xfs_efi_log_item;
struct xfs_inode;
struct xfs_item_ops;
struct xfs_log_iovec;
struct xfs_mount;
struct xfs_trans;
struct xfs_trans_res;
struct xfs_dquot_acct;
struct xfs_rud_log_item;
struct xfs_rui_log_item;
struct xfs_btree_cur;
struct xfs_cui_log_item;
struct xfs_cud_log_item;
struct xfs_bui_log_item;
struct xfs_bud_log_item;
struct xfs_log_item {
struct list_head li_ail; /* AIL pointers */
struct list_head li_trans; /* transaction list */
xfs_lsn_t li_lsn; /* last on-disk lsn */
struct xlog *li_log;
struct xfs_ail *li_ailp; /* ptr to AIL */
uint li_type; /* item type */
unsigned long li_flags; /* misc flags */
struct xfs_buf *li_buf; /* real buffer pointer */
struct list_head li_bio_list; /* buffer item list */
const struct xfs_item_ops *li_ops; /* function list */
xfs: Introduce delayed logging core code The delayed logging code only changes in-memory structures and as such can be enabled and disabled with a mount option. Add the mount option and emit a warning that this is an experimental feature that should not be used in production yet. We also need infrastructure to track committed items that have not yet been written to the log. This is what the Committed Item List (CIL) is for. The log item also needs to be extended to track the current log vector, the associated memory buffer and it's location in the Commit Item List. Extend the log item and log vector structures to enable this tracking. To maintain the current log format for transactions with delayed logging, we need to introduce a checkpoint transaction and a context for tracking each checkpoint from initiation to transaction completion. This includes adding a log ticket for tracking space log required/used by the context checkpoint. To track all the changes we need an io vector array per log item, rather than a single array for the entire transaction. Using the new log vector structure for this requires two passes - the first to allocate the log vector structures and chain them together, and the second to fill them out. This log vector chain can then be passed to the CIL for formatting, pinning and insertion into the CIL. Formatting of the log vector chain is relatively simple - it's just a loop over the iovecs on each log vector, but it is made slightly more complex because we re-write the iovec after the copy to point back at the memory buffer we just copied into. This code also needs to pin log items. If the log item is not already tracked in this checkpoint context, then it needs to be pinned. Otherwise it is already pinned and we don't need to pin it again. The only other complexity is calculating the amount of new log space the formatting has consumed. This needs to be accounted to the transaction in progress, and the accounting is made more complex becase we need also to steal space from it for log metadata in the checkpoint transaction. Calculate all this at insert time and update all the tickets, counters, etc correctly. Once we've formatted all the log items in the transaction, attach the busy extents to the checkpoint context so the busy extents live until checkpoint completion and can be processed at that point in time. Transactions can then be freed at this point in time. Now we need to issue checkpoints - we are tracking the amount of log space used by the items in the CIL, so we can trigger background checkpoints when the space usage gets to a certain threshold. Otherwise, checkpoints need ot be triggered when a log synchronisation point is reached - a log force event. Because the log write code already handles chained log vectors, writing the transaction is trivial, too. Construct a transaction header, add it to the head of the chain and write it into the log, then issue a commit record write. Then we can release the checkpoint log ticket and attach the context to the log buffer so it can be called during Io completion to complete the checkpoint. We also need to allow for synchronising multiple in-flight checkpoints. This is needed for two things - the first is to ensure that checkpoint commit records appear in the log in the correct sequence order (so they are replayed in the correct order). The second is so that xfs_log_force_lsn() operates correctly and only flushes and/or waits for the specific sequence it was provided with. To do this we need a wait variable and a list tracking the checkpoint commits in progress. We can walk this list and wait for the checkpoints to change state or complete easily, an this provides the necessary synchronisation for correct operation in both cases. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Alex Elder <aelder@sgi.com>
2010-05-21 12:37:18 +08:00
/* delayed logging */
struct list_head li_cil; /* CIL pointers */
struct xfs_log_vec *li_lv; /* active log vector */
xfs: allocate log vector buffers outside CIL context lock One of the problems we currently have with delayed logging is that under serious memory pressure we can deadlock memory reclaim. THis occurs when memory reclaim (such as run by kswapd) is reclaiming XFS inodes and issues a log force to unpin inodes that are dirty in the CIL. The CIL is pushed, but this will only occur once it gets the CIL context lock to ensure that all committing transactions are complete and no new transactions start being committed to the CIL while the push switches to a new context. The deadlock occurs when the CIL context lock is held by a committing process that is doing memory allocation for log vector buffers, and that allocation is then blocked on memory reclaim making progress. Memory reclaim, however, is blocked waiting for a log force to make progress, and so we effectively deadlock at this point. To solve this problem, we have to move the CIL log vector buffer allocation outside of the context lock so that memory reclaim can always make progress when it needs to force the log. The problem with doing this is that a CIL push can take place while we are determining if we need to allocate a new log vector buffer for an item and hence the current log vector may go away without warning. That means we canot rely on the existing log vector being present when we finally grab the context lock and so we must have a replacement buffer ready to go at all times. To ensure this, introduce a "shadow log vector" buffer that is always guaranteed to be present when we gain the CIL context lock and format the item. This shadow buffer may or may not be used during the formatting, but if the log item does not have an existing log vector buffer or that buffer is too small for the new modifications, we swap it for the new shadow buffer and format the modifications into that new log vector buffer. The result of this is that for any object we modify more than once in a given CIL checkpoint, we double the memory required to track dirty regions in the log. For single modifications then we consume the shadow log vectorwe allocate on commit, and that gets consumed by the checkpoint. However, if we make multiple modifications, then the second transaction commit will allocate a shadow log vector and hence we will end up with double the memory usage as only one of the log vectors is consumed by the CIL checkpoint. The remaining shadow vector will be freed when th elog item is freed. This can probably be optimised in future - access to the shadow log vector is serialised by the object lock (as opposited to the active log vector, which is controlled by the CIL context lock) and so we can probably free shadow log vector from some objects when the log item is marked clean on removal from the AIL. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Brian Foster <bfoster@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2016-07-22 07:52:35 +08:00
struct xfs_log_vec *li_lv_shadow; /* standby vector */
xfs: xfs_log_force_lsn isn't passed a LSN In doing an investigation into AIL push stalls, I was looking at the log force code to see if an async CIL push could be done instead. This lead me to xfs_log_force_lsn() and looking at how it works. xfs_log_force_lsn() is only called from inode synchronisation contexts such as fsync(), and it takes the ip->i_itemp->ili_last_lsn value as the LSN to sync the log to. This gets passed to xlog_cil_force_lsn() via xfs_log_force_lsn() to flush the CIL to the journal, and then used by xfs_log_force_lsn() to flush the iclogs to the journal. The problem is that ip->i_itemp->ili_last_lsn does not store a log sequence number. What it stores is passed to it from the ->iop_committing method, which is called by xfs_log_commit_cil(). The value this passes to the iop_committing method is the CIL context sequence number that the item was committed to. As it turns out, xlog_cil_force_lsn() converts the sequence to an actual commit LSN for the related context and returns that to xfs_log_force_lsn(). xfs_log_force_lsn() overwrites it's "lsn" variable that contained a sequence with an actual LSN and then uses that to sync the iclogs. This caused me some confusion for a while, even though I originally wrote all this code a decade ago. ->iop_committing is only used by a couple of log item types, and only inode items use the sequence number it is passed. Let's clean up the API, CIL structures and inode log item to call it a sequence number, and make it clear that the high level code is using CIL sequence numbers and not on-disk LSNs for integrity synchronisation purposes. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Allison Henderson <allison.henderson@oracle.com> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-06-18 23:21:52 +08:00
xfs_csn_t li_seq; /* CIL commit seq */
uint32_t li_order_id; /* CIL commit order */
};
/*
* li_flags use the (set/test/clear)_bit atomic interfaces because updates can
* race with each other and we don't want to have to use the AIL lock to
* serialise all updates.
*/
#define XFS_LI_IN_AIL 0
#define XFS_LI_ABORTED 1
#define XFS_LI_FAILED 2
xfs: intent item whiteouts When we log modifications based on intents, we add both intent and intent done items to the modification being made. These get written to the log to ensure that the operation is re-run if the intent done is not found in the log. However, for operations that complete wholly within a single checkpoint, the change in the checkpoint is atomic and will never need replay. In this case, we don't need to actually write the intent and intent done items to the journal because log recovery will never need to manually restart this modification. Log recovery currently handles intent/intent done matching by inserting the intent into the AIL, then removing it when a matching intent done item is found. Hence for all the intent-based operations that complete within a checkpoint, we spend all that time parsing the intent/intent done items just to cancel them and do nothing with them. Hence it follows that the only time we actually need intents in the log is when the modification crosses checkpoint boundaries in the log and so may only be partially complete in the journal. Hence if we commit and intent done item to the CIL and the intent item is in the same checkpoint, we don't actually have to write them to the journal because log recovery will always cancel the intents. We've never really worried about the overhead of logging intents unnecessarily like this because the intents we log are generally very much smaller than the change being made. e.g. freeing an extent involves modifying at lease two freespace btree blocks and the AGF, so the EFI/EFD overhead is only a small increase in space and processing time compared to the overall cost of freeing an extent. However, delayed attributes change this cost equation dramatically, especially for inline attributes. In the case of adding an inline attribute, we only log the inode core and attribute fork at present. With delayed attributes, we now log the attr intent which includes the name and value, the inode core adn attr fork, and finally the attr intent done item. We increase the number of items we log from 1 to 3, and the number of log vectors (regions) goes up from 3 to 7. Hence we tripple the number of objects that the CIL has to process, and more than double the number of log vectors that need to be written to the journal. At scale, this means delayed attributes cause a non-pipelined CIL to become CPU bound processing all the extra items, resulting in a > 40% performance degradation on 16-way file+xattr create worklaods. Pipelining the CIL (as per 5.15) reduces the performance degradation to 20%, but now the limitation is the rate at which the log items can be written to the iclogs and iclogs be dispatched for IO and completed. Even log IO completion is slowed down by these intents, because it now has to process 3x the number of items in the checkpoint. Processing completed intents is especially inefficient here, because we first insert the intent into the AIL, then remove it from the AIL when the intent done is processed. IOWs, we are also doing expensive operations in log IO completion we could completely avoid if we didn't log completed intent/intent done pairs. Enter log item whiteouts. When an intent done is committed, we can check to see if the associated intent is in the same checkpoint as we are currently committing the intent done to. If so, we can mark the intent log item with a whiteout and immediately free the intent done item rather than committing it to the CIL. We can basically skip the entire formatting and CIL insertion steps for the intent done item. However, we cannot remove the intent item from the CIL at this point because the unlocked per-cpu CIL item lists do not permit removal without holding the CIL context lock exclusively. Transaction commit only holds the context lock shared, hence the best we can do is mark the intent item with a whiteout so that the CIL push can release it rather than writing it to the log. This means we never write the intent to the log if the intent done has also been committed to the same checkpoint, but we'll always write the intent if the intent done has not been committed or has been committed to a different checkpoint. This will result in correct log recovery behaviour in all cases, without the overhead of logging unnecessary intents. This intent whiteout concept is generic - we can apply it to all intent/intent done pairs that have a direct 1:1 relationship. The way deferred ops iterate and relog intents mean that all intents currently have a 1:1 relationship with their done intent, and hence we can apply this cancellation to all existing intent/intent done implementations. For delayed attributes with a 16-way 64kB xattr create workload, whiteouts reduce the amount of journalled metadata from ~2.5GB/s down to ~600MB/s and improve the creation rate from 9000/s to 14000/s. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Allison Henderson <allison.henderson@oracle.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2022-05-04 09:50:29 +08:00
#define XFS_LI_DIRTY 3
#define XFS_LI_WHITEOUT 4
xfs: event tracing support Convert the old xfs tracing support that could only be used with the out of tree kdb and xfsidbg patches to use the generic event tracer. To use it make sure CONFIG_EVENT_TRACING is enabled and then enable all xfs trace channels by: echo 1 > /sys/kernel/debug/tracing/events/xfs/enable or alternatively enable single events by just doing the same in one event subdirectory, e.g. echo 1 > /sys/kernel/debug/tracing/events/xfs/xfs_ihold/enable or set more complex filters, etc. In Documentation/trace/events.txt all this is desctribed in more detail. To reads the events do a cat /sys/kernel/debug/tracing/trace Compared to the last posting this patch converts the tracing mostly to the one tracepoint per callsite model that other users of the new tracing facility also employ. This allows a very fine-grained control of the tracing, a cleaner output of the traces and also enables the perf tool to use each tracepoint as a virtual performance counter, allowing us to e.g. count how often certain workloads git various spots in XFS. Take a look at http://lwn.net/Articles/346470/ for some examples. Also the btree tracing isn't included at all yet, as it will require additional core tracing features not in mainline yet, I plan to deliver it later. And the really nice thing about this patch is that it actually removes many lines of code while adding this nice functionality: fs/xfs/Makefile | 8 fs/xfs/linux-2.6/xfs_acl.c | 1 fs/xfs/linux-2.6/xfs_aops.c | 52 - fs/xfs/linux-2.6/xfs_aops.h | 2 fs/xfs/linux-2.6/xfs_buf.c | 117 +-- fs/xfs/linux-2.6/xfs_buf.h | 33 fs/xfs/linux-2.6/xfs_fs_subr.c | 3 fs/xfs/linux-2.6/xfs_ioctl.c | 1 fs/xfs/linux-2.6/xfs_ioctl32.c | 1 fs/xfs/linux-2.6/xfs_iops.c | 1 fs/xfs/linux-2.6/xfs_linux.h | 1 fs/xfs/linux-2.6/xfs_lrw.c | 87 -- fs/xfs/linux-2.6/xfs_lrw.h | 45 - fs/xfs/linux-2.6/xfs_super.c | 104 --- fs/xfs/linux-2.6/xfs_super.h | 7 fs/xfs/linux-2.6/xfs_sync.c | 1 fs/xfs/linux-2.6/xfs_trace.c | 75 ++ fs/xfs/linux-2.6/xfs_trace.h | 1369 +++++++++++++++++++++++++++++++++++++++++ fs/xfs/linux-2.6/xfs_vnode.h | 4 fs/xfs/quota/xfs_dquot.c | 110 --- fs/xfs/quota/xfs_dquot.h | 21 fs/xfs/quota/xfs_qm.c | 40 - fs/xfs/quota/xfs_qm_syscalls.c | 4 fs/xfs/support/ktrace.c | 323 --------- fs/xfs/support/ktrace.h | 85 -- fs/xfs/xfs.h | 16 fs/xfs/xfs_ag.h | 14 fs/xfs/xfs_alloc.c | 230 +----- fs/xfs/xfs_alloc.h | 27 fs/xfs/xfs_alloc_btree.c | 1 fs/xfs/xfs_attr.c | 107 --- fs/xfs/xfs_attr.h | 10 fs/xfs/xfs_attr_leaf.c | 14 fs/xfs/xfs_attr_sf.h | 40 - fs/xfs/xfs_bmap.c | 507 +++------------ fs/xfs/xfs_bmap.h | 49 - fs/xfs/xfs_bmap_btree.c | 6 fs/xfs/xfs_btree.c | 5 fs/xfs/xfs_btree_trace.h | 17 fs/xfs/xfs_buf_item.c | 87 -- fs/xfs/xfs_buf_item.h | 20 fs/xfs/xfs_da_btree.c | 3 fs/xfs/xfs_da_btree.h | 7 fs/xfs/xfs_dfrag.c | 2 fs/xfs/xfs_dir2.c | 8 fs/xfs/xfs_dir2_block.c | 20 fs/xfs/xfs_dir2_leaf.c | 21 fs/xfs/xfs_dir2_node.c | 27 fs/xfs/xfs_dir2_sf.c | 26 fs/xfs/xfs_dir2_trace.c | 216 ------ fs/xfs/xfs_dir2_trace.h | 72 -- fs/xfs/xfs_filestream.c | 8 fs/xfs/xfs_fsops.c | 2 fs/xfs/xfs_iget.c | 111 --- fs/xfs/xfs_inode.c | 67 -- fs/xfs/xfs_inode.h | 76 -- fs/xfs/xfs_inode_item.c | 5 fs/xfs/xfs_iomap.c | 85 -- fs/xfs/xfs_iomap.h | 8 fs/xfs/xfs_log.c | 181 +---- fs/xfs/xfs_log_priv.h | 20 fs/xfs/xfs_log_recover.c | 1 fs/xfs/xfs_mount.c | 2 fs/xfs/xfs_quota.h | 8 fs/xfs/xfs_rename.c | 1 fs/xfs/xfs_rtalloc.c | 1 fs/xfs/xfs_rw.c | 3 fs/xfs/xfs_trans.h | 47 + fs/xfs/xfs_trans_buf.c | 62 - fs/xfs/xfs_vnodeops.c | 8 70 files changed, 2151 insertions(+), 2592 deletions(-) Signed-off-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Alex Elder <aelder@sgi.com>
2009-12-15 07:14:59 +08:00
#define XFS_LI_FLAGS \
{ (1u << XFS_LI_IN_AIL), "IN_AIL" }, \
{ (1u << XFS_LI_ABORTED), "ABORTED" }, \
{ (1u << XFS_LI_FAILED), "FAILED" }, \
xfs: intent item whiteouts When we log modifications based on intents, we add both intent and intent done items to the modification being made. These get written to the log to ensure that the operation is re-run if the intent done is not found in the log. However, for operations that complete wholly within a single checkpoint, the change in the checkpoint is atomic and will never need replay. In this case, we don't need to actually write the intent and intent done items to the journal because log recovery will never need to manually restart this modification. Log recovery currently handles intent/intent done matching by inserting the intent into the AIL, then removing it when a matching intent done item is found. Hence for all the intent-based operations that complete within a checkpoint, we spend all that time parsing the intent/intent done items just to cancel them and do nothing with them. Hence it follows that the only time we actually need intents in the log is when the modification crosses checkpoint boundaries in the log and so may only be partially complete in the journal. Hence if we commit and intent done item to the CIL and the intent item is in the same checkpoint, we don't actually have to write them to the journal because log recovery will always cancel the intents. We've never really worried about the overhead of logging intents unnecessarily like this because the intents we log are generally very much smaller than the change being made. e.g. freeing an extent involves modifying at lease two freespace btree blocks and the AGF, so the EFI/EFD overhead is only a small increase in space and processing time compared to the overall cost of freeing an extent. However, delayed attributes change this cost equation dramatically, especially for inline attributes. In the case of adding an inline attribute, we only log the inode core and attribute fork at present. With delayed attributes, we now log the attr intent which includes the name and value, the inode core adn attr fork, and finally the attr intent done item. We increase the number of items we log from 1 to 3, and the number of log vectors (regions) goes up from 3 to 7. Hence we tripple the number of objects that the CIL has to process, and more than double the number of log vectors that need to be written to the journal. At scale, this means delayed attributes cause a non-pipelined CIL to become CPU bound processing all the extra items, resulting in a > 40% performance degradation on 16-way file+xattr create worklaods. Pipelining the CIL (as per 5.15) reduces the performance degradation to 20%, but now the limitation is the rate at which the log items can be written to the iclogs and iclogs be dispatched for IO and completed. Even log IO completion is slowed down by these intents, because it now has to process 3x the number of items in the checkpoint. Processing completed intents is especially inefficient here, because we first insert the intent into the AIL, then remove it from the AIL when the intent done is processed. IOWs, we are also doing expensive operations in log IO completion we could completely avoid if we didn't log completed intent/intent done pairs. Enter log item whiteouts. When an intent done is committed, we can check to see if the associated intent is in the same checkpoint as we are currently committing the intent done to. If so, we can mark the intent log item with a whiteout and immediately free the intent done item rather than committing it to the CIL. We can basically skip the entire formatting and CIL insertion steps for the intent done item. However, we cannot remove the intent item from the CIL at this point because the unlocked per-cpu CIL item lists do not permit removal without holding the CIL context lock exclusively. Transaction commit only holds the context lock shared, hence the best we can do is mark the intent item with a whiteout so that the CIL push can release it rather than writing it to the log. This means we never write the intent to the log if the intent done has also been committed to the same checkpoint, but we'll always write the intent if the intent done has not been committed or has been committed to a different checkpoint. This will result in correct log recovery behaviour in all cases, without the overhead of logging unnecessary intents. This intent whiteout concept is generic - we can apply it to all intent/intent done pairs that have a direct 1:1 relationship. The way deferred ops iterate and relog intents mean that all intents currently have a 1:1 relationship with their done intent, and hence we can apply this cancellation to all existing intent/intent done implementations. For delayed attributes with a 16-way 64kB xattr create workload, whiteouts reduce the amount of journalled metadata from ~2.5GB/s down to ~600MB/s and improve the creation rate from 9000/s to 14000/s. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Allison Henderson <allison.henderson@oracle.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2022-05-04 09:50:29 +08:00
{ (1u << XFS_LI_DIRTY), "DIRTY" }, \
{ (1u << XFS_LI_WHITEOUT), "WHITEOUT" }
xfs: event tracing support Convert the old xfs tracing support that could only be used with the out of tree kdb and xfsidbg patches to use the generic event tracer. To use it make sure CONFIG_EVENT_TRACING is enabled and then enable all xfs trace channels by: echo 1 > /sys/kernel/debug/tracing/events/xfs/enable or alternatively enable single events by just doing the same in one event subdirectory, e.g. echo 1 > /sys/kernel/debug/tracing/events/xfs/xfs_ihold/enable or set more complex filters, etc. In Documentation/trace/events.txt all this is desctribed in more detail. To reads the events do a cat /sys/kernel/debug/tracing/trace Compared to the last posting this patch converts the tracing mostly to the one tracepoint per callsite model that other users of the new tracing facility also employ. This allows a very fine-grained control of the tracing, a cleaner output of the traces and also enables the perf tool to use each tracepoint as a virtual performance counter, allowing us to e.g. count how often certain workloads git various spots in XFS. Take a look at http://lwn.net/Articles/346470/ for some examples. Also the btree tracing isn't included at all yet, as it will require additional core tracing features not in mainline yet, I plan to deliver it later. And the really nice thing about this patch is that it actually removes many lines of code while adding this nice functionality: fs/xfs/Makefile | 8 fs/xfs/linux-2.6/xfs_acl.c | 1 fs/xfs/linux-2.6/xfs_aops.c | 52 - fs/xfs/linux-2.6/xfs_aops.h | 2 fs/xfs/linux-2.6/xfs_buf.c | 117 +-- fs/xfs/linux-2.6/xfs_buf.h | 33 fs/xfs/linux-2.6/xfs_fs_subr.c | 3 fs/xfs/linux-2.6/xfs_ioctl.c | 1 fs/xfs/linux-2.6/xfs_ioctl32.c | 1 fs/xfs/linux-2.6/xfs_iops.c | 1 fs/xfs/linux-2.6/xfs_linux.h | 1 fs/xfs/linux-2.6/xfs_lrw.c | 87 -- fs/xfs/linux-2.6/xfs_lrw.h | 45 - fs/xfs/linux-2.6/xfs_super.c | 104 --- fs/xfs/linux-2.6/xfs_super.h | 7 fs/xfs/linux-2.6/xfs_sync.c | 1 fs/xfs/linux-2.6/xfs_trace.c | 75 ++ fs/xfs/linux-2.6/xfs_trace.h | 1369 +++++++++++++++++++++++++++++++++++++++++ fs/xfs/linux-2.6/xfs_vnode.h | 4 fs/xfs/quota/xfs_dquot.c | 110 --- fs/xfs/quota/xfs_dquot.h | 21 fs/xfs/quota/xfs_qm.c | 40 - fs/xfs/quota/xfs_qm_syscalls.c | 4 fs/xfs/support/ktrace.c | 323 --------- fs/xfs/support/ktrace.h | 85 -- fs/xfs/xfs.h | 16 fs/xfs/xfs_ag.h | 14 fs/xfs/xfs_alloc.c | 230 +----- fs/xfs/xfs_alloc.h | 27 fs/xfs/xfs_alloc_btree.c | 1 fs/xfs/xfs_attr.c | 107 --- fs/xfs/xfs_attr.h | 10 fs/xfs/xfs_attr_leaf.c | 14 fs/xfs/xfs_attr_sf.h | 40 - fs/xfs/xfs_bmap.c | 507 +++------------ fs/xfs/xfs_bmap.h | 49 - fs/xfs/xfs_bmap_btree.c | 6 fs/xfs/xfs_btree.c | 5 fs/xfs/xfs_btree_trace.h | 17 fs/xfs/xfs_buf_item.c | 87 -- fs/xfs/xfs_buf_item.h | 20 fs/xfs/xfs_da_btree.c | 3 fs/xfs/xfs_da_btree.h | 7 fs/xfs/xfs_dfrag.c | 2 fs/xfs/xfs_dir2.c | 8 fs/xfs/xfs_dir2_block.c | 20 fs/xfs/xfs_dir2_leaf.c | 21 fs/xfs/xfs_dir2_node.c | 27 fs/xfs/xfs_dir2_sf.c | 26 fs/xfs/xfs_dir2_trace.c | 216 ------ fs/xfs/xfs_dir2_trace.h | 72 -- fs/xfs/xfs_filestream.c | 8 fs/xfs/xfs_fsops.c | 2 fs/xfs/xfs_iget.c | 111 --- fs/xfs/xfs_inode.c | 67 -- fs/xfs/xfs_inode.h | 76 -- fs/xfs/xfs_inode_item.c | 5 fs/xfs/xfs_iomap.c | 85 -- fs/xfs/xfs_iomap.h | 8 fs/xfs/xfs_log.c | 181 +---- fs/xfs/xfs_log_priv.h | 20 fs/xfs/xfs_log_recover.c | 1 fs/xfs/xfs_mount.c | 2 fs/xfs/xfs_quota.h | 8 fs/xfs/xfs_rename.c | 1 fs/xfs/xfs_rtalloc.c | 1 fs/xfs/xfs_rw.c | 3 fs/xfs/xfs_trans.h | 47 + fs/xfs/xfs_trans_buf.c | 62 - fs/xfs/xfs_vnodeops.c | 8 70 files changed, 2151 insertions(+), 2592 deletions(-) Signed-off-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Alex Elder <aelder@sgi.com>
2009-12-15 07:14:59 +08:00
struct xfs_item_ops {
unsigned flags;
void (*iop_size)(struct xfs_log_item *, int *, int *);
void (*iop_format)(struct xfs_log_item *, struct xfs_log_vec *);
void (*iop_pin)(struct xfs_log_item *);
void (*iop_unpin)(struct xfs_log_item *, int remove);
xfs: add log item precommit operation For inodes that are dirty, we have an attached cluster buffer that we want to use to track the dirty inode through the AIL. Unfortunately, locking the cluster buffer and adding it to the transaction when the inode is first logged in a transaction leads to buffer lock ordering inversions. The specific problem is ordering against the AGI buffer. When modifying unlinked lists, the buffer lock order is AGI -> inode cluster buffer as the AGI buffer lock serialises all access to the unlinked lists. Unfortunately, functionality like xfs_droplink() logs the inode before calling xfs_iunlink(), as do various directory manipulation functions. The inode can be logged way down in the stack as far as the bmapi routines and hence, without a major rewrite of lots of APIs there's no way we can avoid the inode being logged by something until after the AGI has been logged. As we are going to be using ordered buffers for inode AIL tracking, there isn't a need to actually lock that buffer against modification as all the modifications are captured by logging the inode item itself. Hence we don't actually need to join the cluster buffer into the transaction until just before it is committed. This means we do not perturb any of the existing buffer lock orders in transactions, and the inode cluster buffer is always locked last in a transaction that doesn't otherwise touch inode cluster buffers. We do this by introducing a precommit log item method. This commit just introduces the mechanism; the inode item implementation is in followup commits. The precommit items need to be sorted into consistent order as we may be locking multiple items here. Hence if we have two dirty inodes in cluster buffers A and B, and some other transaction has two separate dirty inodes in the same cluster buffers, locking them in different orders opens us up to ABBA deadlocks. Hence we sort the items on the transaction based on the presence of a sort log item method. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Christoph Hellwig <hch@lst.de>
2022-07-14 09:47:26 +08:00
uint64_t (*iop_sort)(struct xfs_log_item *lip);
int (*iop_precommit)(struct xfs_trans *tp, struct xfs_log_item *lip);
xfs: xfs_log_force_lsn isn't passed a LSN In doing an investigation into AIL push stalls, I was looking at the log force code to see if an async CIL push could be done instead. This lead me to xfs_log_force_lsn() and looking at how it works. xfs_log_force_lsn() is only called from inode synchronisation contexts such as fsync(), and it takes the ip->i_itemp->ili_last_lsn value as the LSN to sync the log to. This gets passed to xlog_cil_force_lsn() via xfs_log_force_lsn() to flush the CIL to the journal, and then used by xfs_log_force_lsn() to flush the iclogs to the journal. The problem is that ip->i_itemp->ili_last_lsn does not store a log sequence number. What it stores is passed to it from the ->iop_committing method, which is called by xfs_log_commit_cil(). The value this passes to the iop_committing method is the CIL context sequence number that the item was committed to. As it turns out, xlog_cil_force_lsn() converts the sequence to an actual commit LSN for the related context and returns that to xfs_log_force_lsn(). xfs_log_force_lsn() overwrites it's "lsn" variable that contained a sequence with an actual LSN and then uses that to sync the iclogs. This caused me some confusion for a while, even though I originally wrote all this code a decade ago. ->iop_committing is only used by a couple of log item types, and only inode items use the sequence number it is passed. Let's clean up the API, CIL structures and inode log item to call it a sequence number, and make it clear that the high level code is using CIL sequence numbers and not on-disk LSNs for integrity synchronisation purposes. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Allison Henderson <allison.henderson@oracle.com> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
2021-06-18 23:21:52 +08:00
void (*iop_committing)(struct xfs_log_item *lip, xfs_csn_t seq);
xfs_lsn_t (*iop_committed)(struct xfs_log_item *, xfs_lsn_t);
xfs: add log item precommit operation For inodes that are dirty, we have an attached cluster buffer that we want to use to track the dirty inode through the AIL. Unfortunately, locking the cluster buffer and adding it to the transaction when the inode is first logged in a transaction leads to buffer lock ordering inversions. The specific problem is ordering against the AGI buffer. When modifying unlinked lists, the buffer lock order is AGI -> inode cluster buffer as the AGI buffer lock serialises all access to the unlinked lists. Unfortunately, functionality like xfs_droplink() logs the inode before calling xfs_iunlink(), as do various directory manipulation functions. The inode can be logged way down in the stack as far as the bmapi routines and hence, without a major rewrite of lots of APIs there's no way we can avoid the inode being logged by something until after the AGI has been logged. As we are going to be using ordered buffers for inode AIL tracking, there isn't a need to actually lock that buffer against modification as all the modifications are captured by logging the inode item itself. Hence we don't actually need to join the cluster buffer into the transaction until just before it is committed. This means we do not perturb any of the existing buffer lock orders in transactions, and the inode cluster buffer is always locked last in a transaction that doesn't otherwise touch inode cluster buffers. We do this by introducing a precommit log item method. This commit just introduces the mechanism; the inode item implementation is in followup commits. The precommit items need to be sorted into consistent order as we may be locking multiple items here. Hence if we have two dirty inodes in cluster buffers A and B, and some other transaction has two separate dirty inodes in the same cluster buffers, locking them in different orders opens us up to ABBA deadlocks. Hence we sort the items on the transaction based on the presence of a sort log item method. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Christoph Hellwig <hch@lst.de>
2022-07-14 09:47:26 +08:00
uint (*iop_push)(struct xfs_log_item *, struct list_head *);
void (*iop_release)(struct xfs_log_item *);
xfs: proper replay of deferred ops queued during log recovery When we replay unfinished intent items that have been recovered from the log, it's possible that the replay will cause the creation of more deferred work items. As outlined in commit 509955823cc9c ("xfs: log recovery should replay deferred ops in order"), later work items have an implicit ordering dependency on earlier work items. Therefore, recovery must replay the items (both recovered and created) in the same order that they would have been during normal operation. For log recovery, we enforce this ordering by using an empty transaction to collect deferred ops that get created in the process of recovering a log intent item to prevent them from being committed before the rest of the recovered intent items. After we finish committing all the recovered log items, we allocate a transaction with an enormous block reservation, splice our huge list of created deferred ops into that transaction, and commit it, thereby finishing all those ops. This is /really/ hokey -- it's the one place in XFS where we allow nested transactions; the splicing of the defer ops list is is inelegant and has to be done twice per recovery function; and the broken way we handle inode pointers and block reservations cause subtle use-after-free and allocator problems that will be fixed by this patch and the two patches after it. Therefore, replace the hokey empty transaction with a structure designed to capture each chain of deferred ops that are created as part of recovering a single unfinished log intent. Finally, refactor the loop that replays those chains to do so using one transaction per chain. Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com> Reviewed-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
2020-09-26 08:39:37 +08:00
int (*iop_recover)(struct xfs_log_item *lip,
struct list_head *capture_list);
bool (*iop_match)(struct xfs_log_item *item, uint64_t id);
xfs: periodically relog deferred intent items There's a subtle design flaw in the deferred log item code that can lead to pinning the log tail. Taking up the defer ops chain examples from the previous commit, we can get trapped in sequences like this: Caller hands us a transaction t0 with D0-D3 attached. The defer ops chain will look like the following if the transaction rolls succeed: t1: D0(t0), D1(t0), D2(t0), D3(t0) t2: d4(t1), d5(t1), D1(t0), D2(t0), D3(t0) t3: d5(t1), D1(t0), D2(t0), D3(t0) ... t9: d9(t7), D3(t0) t10: D3(t0) t11: d10(t10), d11(t10) t12: d11(t10) In transaction 9, we finish d9 and try to roll to t10 while holding onto an intent item for D3 that we logged in t0. The previous commit changed the order in which we place new defer ops in the defer ops processing chain to reduce the maximum chain length. Now make xfs_defer_finish_noroll capable of relogging the entire chain periodically so that we can always move the log tail forward. Most chains will never get relogged, except for operations that generate very long chains (large extents containing many blocks with different sharing levels) or are on filesystems with small logs and a lot of ongoing metadata updates. Callers are now required to ensure that the transaction reservation is large enough to handle logging done items and new intent items for the maximum possible chain length. Most callers are careful to keep the chain lengths low, so the overhead should be minimal. The decision to relog an intent item is made based on whether the intent was logged in a previous checkpoint, since there's no point in relogging an intent into the same checkpoint. Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com> Reviewed-by: Brian Foster <bfoster@redhat.com>
2020-09-28 07:18:13 +08:00
struct xfs_log_item *(*iop_relog)(struct xfs_log_item *intent,
struct xfs_trans *tp);
struct xfs_log_item *(*iop_intent)(struct xfs_log_item *intent_done);
};
/*
* Log item ops flags
*/
/*
* Release the log item when the journal commits instead of inserting into the
* AIL for writeback tracking and/or log tail pinning.
*/
#define XFS_ITEM_RELEASE_WHEN_COMMITTED (1 << 0)
#define XFS_ITEM_INTENT (1 << 1)
#define XFS_ITEM_INTENT_DONE (1 << 2)
static inline bool
xlog_item_is_intent(struct xfs_log_item *lip)
{
return lip->li_ops->flags & XFS_ITEM_INTENT;
}
static inline bool
xlog_item_is_intent_done(struct xfs_log_item *lip)
{
return lip->li_ops->flags & XFS_ITEM_INTENT_DONE;
}
void xfs_log_item_init(struct xfs_mount *mp, struct xfs_log_item *item,
int type, const struct xfs_item_ops *ops);
/*
* Return values for the iop_push() routines.
*/
xfs: on-stack delayed write buffer lists Queue delwri buffers on a local on-stack list instead of a per-buftarg one, and write back the buffers per-process instead of by waking up xfsbufd. This is now easily doable given that we have very few places left that write delwri buffers: - log recovery: Only done at mount time, and already forcing out the buffers synchronously using xfs_flush_buftarg - quotacheck: Same story. - dquot reclaim: Writes out dirty dquots on the LRU under memory pressure. We might want to look into doing more of this via xfsaild, but it's already more optimal than the synchronous inode reclaim that writes each buffer synchronously. - xfsaild: This is the main beneficiary of the change. By keeping a local list of buffers to write we reduce latency of writing out buffers, and more importably we can remove all the delwri list promotions which were hitting the buffer cache hard under sustained metadata loads. The implementation is very straight forward - xfs_buf_delwri_queue now gets a new list_head pointer that it adds the delwri buffers to, and all callers need to eventually submit the list using xfs_buf_delwi_submit or xfs_buf_delwi_submit_nowait. Buffers that already are on a delwri list are skipped in xfs_buf_delwri_queue, assuming they already are on another delwri list. The biggest change to pass down the buffer list was done to the AIL pushing. Now that we operate on buffers the trylock, push and pushbuf log item methods are merged into a single push routine, which tries to lock the item, and if possible add the buffer that needs writeback to the buffer list. This leads to much simpler code than the previous split but requires the individual IOP_PUSH instances to unlock and reacquire the AIL around calls to blocking routines. Given that xfsailds now also handle writing out buffers, the conditions for log forcing and the sleep times needed some small changes. The most important one is that we consider an AIL busy as long we still have buffers to push, and the other one is that we do increment the pushed LSN for buffers that are under flushing at this moment, but still count them towards the stuck items for restart purposes. Without this we could hammer on stuck items without ever forcing the log and not make progress under heavy random delete workloads on fast flash storage devices. [ Dave Chinner: - rebase on previous patches. - improved comments for XBF_DELWRI_Q handling - fix XBF_ASYNC handling in queue submission (test 106 failure) - rename delwri submit function buffer list parameters for clarity - xfs_efd_item_push() should return XFS_ITEM_PINNED ] Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Mark Tinguely <tinguely@sgi.com> Signed-off-by: Ben Myers <bpm@sgi.com>
2012-04-23 13:58:39 +08:00
#define XFS_ITEM_SUCCESS 0
#define XFS_ITEM_PINNED 1
#define XFS_ITEM_LOCKED 2
#define XFS_ITEM_FLUSHING 3
/*
* This is the structure maintained for every active transaction.
*/
typedef struct xfs_trans {
unsigned int t_magic; /* magic number */
unsigned int t_log_res; /* amt of log space resvd */
unsigned int t_log_count; /* count for perm log res */
unsigned int t_blk_res; /* # of blocks resvd */
unsigned int t_blk_res_used; /* # of resvd blocks used */
unsigned int t_rtx_res; /* # of rt extents resvd */
unsigned int t_rtx_res_used; /* # of resvd rt extents used */
unsigned int t_flags; /* misc flags */
xfs_agnumber_t t_highest_agno; /* highest AGF locked */
struct xlog_ticket *t_ticket; /* log mgr ticket */
struct xfs_mount *t_mountp; /* ptr to fs mount struct */
struct xfs_dquot_acct *t_dqinfo; /* acctg info for dquots */
int64_t t_icount_delta; /* superblock icount change */
int64_t t_ifree_delta; /* superblock ifree change */
int64_t t_fdblocks_delta; /* superblock fdblocks chg */
int64_t t_res_fdblocks_delta; /* on-disk only chg */
int64_t t_frextents_delta;/* superblock freextents chg*/
int64_t t_res_frextents_delta; /* on-disk only chg */
int64_t t_dblocks_delta;/* superblock dblocks change */
int64_t t_agcount_delta;/* superblock agcount change */
int64_t t_imaxpct_delta;/* superblock imaxpct change */
int64_t t_rextsize_delta;/* superblock rextsize chg */
int64_t t_rbmblocks_delta;/* superblock rbmblocks chg */
int64_t t_rblocks_delta;/* superblock rblocks change */
int64_t t_rextents_delta;/* superblocks rextents chg */
int64_t t_rextslog_delta;/* superblocks rextslog chg */
struct list_head t_items; /* log item descriptors */
xfs: Improve scalability of busy extent tracking When we free a metadata extent, we record it in the per-AG busy extent array so that it is not re-used before the freeing transaction hits the disk. This array is fixed size, so when it overflows we make further allocation transactions synchronous because we cannot track more freed extents until those transactions hit the disk and are completed. Under heavy mixed allocation and freeing workloads with large log buffers, we can overflow this array quite easily. Further, the array is sparsely populated, which means that inserts need to search for a free slot, and array searches often have to search many more slots that are actually used to check all the busy extents. Quite inefficient, really. To enable this aspect of extent freeing to scale better, we need a structure that can grow dynamically. While in other areas of XFS we have used radix trees, the extents being freed are at random locations on disk so are better suited to being indexed by an rbtree. So, use a per-AG rbtree indexed by block number to track busy extents. This incures a memory allocation when marking an extent busy, but should not occur too often in low memory situations. This should scale to an arbitrary number of extents so should not be a limitation for features such as in-memory aggregation of transactions. However, there are still situations where we can't avoid allocating busy extents (such as allocation from the AGFL). To minimise the overhead of such occurences, we need to avoid doing a synchronous log force while holding the AGF locked to ensure that the previous transactions are safely on disk before we use the extent. We can do this by marking the transaction doing the allocation as synchronous rather issuing a log force. Because of the locking involved and the ordering of transactions, the synchronous transaction provides the same guarantees as a synchronous log force because it ensures that all the prior transactions are already on disk when the synchronous transaction hits the disk. i.e. it preserves the free->allocate order of the extent correctly in recovery. By doing this, we avoid holding the AGF locked while log writes are in progress, hence reducing the length of time the lock is held and therefore we increase the rate at which we can allocate and free from the allocation group, thereby increasing overall throughput. The only problem with this approach is that when a metadata buffer is marked stale (e.g. a directory block is removed), then buffer remains pinned and locked until the log goes to disk. The issue here is that if that stale buffer is reallocated in a subsequent transaction, the attempt to lock that buffer in the transaction will hang waiting the log to go to disk to unlock and unpin the buffer. Hence if someone tries to lock a pinned, stale, locked buffer we need to push on the log to get it unlocked ASAP. Effectively we are trading off a guaranteed log force for a much less common trigger for log force to occur. Ideally we should not reallocate busy extents. That is a much more complex fix to the problem as it involves direct intervention in the allocation btree searches in many places. This is left to a future set of modifications. Finally, now that we track busy extents in allocated memory, we don't need the descriptors in the transaction structure to point to them. We can replace the complex busy chunk infrastructure with a simple linked list of busy extents. This allows us to remove a large chunk of code, making the overall change a net reduction in code size. Signed-off-by: Dave Chinner <david@fromorbit.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Alex Elder <aelder@sgi.com>
2010-05-21 10:07:08 +08:00
struct list_head t_busy; /* list of busy extents */
struct list_head t_dfops; /* deferred operations */
unsigned long t_pflags; /* saved process flags state */
} xfs_trans_t;
/*
* XFS transaction mechanism exported interfaces that are
* actually macros.
*/
#define xfs_trans_set_sync(tp) ((tp)->t_flags |= XFS_TRANS_SYNC)
/*
* XFS transaction mechanism exported interfaces.
*/
int xfs_trans_alloc(struct xfs_mount *mp, struct xfs_trans_res *resp,
uint blocks, uint rtextents, uint flags,
struct xfs_trans **tpp);
int xfs_trans_alloc_empty(struct xfs_mount *mp,
struct xfs_trans **tpp);
void xfs_trans_mod_sb(xfs_trans_t *, uint, int64_t);
int xfs_trans_get_buf_map(struct xfs_trans *tp, struct xfs_buftarg *target,
struct xfs_buf_map *map, int nmaps, xfs_buf_flags_t flags,
struct xfs_buf **bpp);
static inline int
xfs_trans_get_buf(
struct xfs_trans *tp,
struct xfs_buftarg *target,
xfs_daddr_t blkno,
int numblks,
2022-04-21 06:44:59 +08:00
xfs_buf_flags_t flags,
struct xfs_buf **bpp)
{
DEFINE_SINGLE_BUF_MAP(map, blkno, numblks);
return xfs_trans_get_buf_map(tp, target, &map, 1, flags, bpp);
}
int xfs_trans_read_buf_map(struct xfs_mount *mp,
struct xfs_trans *tp,
struct xfs_buftarg *target,
struct xfs_buf_map *map, int nmaps,
xfs_buf_flags_t flags,
struct xfs_buf **bpp,
const struct xfs_buf_ops *ops);
static inline int
xfs_trans_read_buf(
struct xfs_mount *mp,
struct xfs_trans *tp,
struct xfs_buftarg *target,
xfs_daddr_t blkno,
int numblks,
xfs_buf_flags_t flags,
struct xfs_buf **bpp,
const struct xfs_buf_ops *ops)
{
DEFINE_SINGLE_BUF_MAP(map, blkno, numblks);
return xfs_trans_read_buf_map(mp, tp, target, &map, 1,
flags, bpp, ops);
}
struct xfs_buf *xfs_trans_getsb(struct xfs_trans *);
void xfs_trans_brelse(xfs_trans_t *, struct xfs_buf *);
void xfs_trans_bjoin(xfs_trans_t *, struct xfs_buf *);
void xfs_trans_bhold(xfs_trans_t *, struct xfs_buf *);
void xfs_trans_bhold_release(xfs_trans_t *, struct xfs_buf *);
void xfs_trans_binval(xfs_trans_t *, struct xfs_buf *);
void xfs_trans_inode_buf(xfs_trans_t *, struct xfs_buf *);
void xfs_trans_stale_inode_buf(xfs_trans_t *, struct xfs_buf *);
xfs: disallow marking previously dirty buffers as ordered Ordered buffers are used in situations where the buffer is not physically logged but must pass through the transaction/logging pipeline for a particular transaction. As a result, ordered buffers are not unpinned and written back until the transaction commits to the log. Ordered buffers have a strict requirement that the target buffer must not be currently dirty and resident in the log pipeline at the time it is marked ordered. If a dirty+ordered buffer is committed, the buffer is reinserted to the AIL but not physically relogged at the LSN of the associated checkpoint. The buffer log item is assigned the LSN of the latest checkpoint and the AIL effectively releases the previously logged buffer content from the active log before the buffer has been written back. If the tail pushes forward and a filesystem crash occurs while in this state, an inconsistent filesystem could result. It is currently the caller responsibility to ensure an ordered buffer is not already dirty from a previous modification. This is unclear and error prone when not used in situations where it is guaranteed a buffer has not been previously modified (such as new metadata allocations). To facilitate general purpose use of ordered buffers, update xfs_trans_ordered_buf() to conditionally order the buffer based on state of the log item and return the status of the result. If the bli is dirty, do not order the buffer and return false. The caller must either physically log the buffer (having acquired the appropriate log reservation) or push it from the AIL to clean it before it can be marked ordered in the current transaction. Note that ordered buffers are currently only used in two situations: 1.) inode chunk allocation where previously logged buffers are not possible and 2.) extent swap which will be updated to handle ordered buffer failures in a separate patch. Signed-off-by: Brian Foster <bfoster@redhat.com> Reviewed-by: Darrick J. Wong <darrick.wong@oracle.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
2017-08-30 01:08:40 +08:00
bool xfs_trans_ordered_buf(xfs_trans_t *, struct xfs_buf *);
void xfs_trans_dquot_buf(xfs_trans_t *, struct xfs_buf *, uint);
void xfs_trans_inode_alloc_buf(xfs_trans_t *, struct xfs_buf *);
void xfs_trans_ichgtime(struct xfs_trans *, struct xfs_inode *, int);
void xfs_trans_ijoin(struct xfs_trans *, struct xfs_inode *, uint);
void xfs_trans_log_buf(struct xfs_trans *, struct xfs_buf *, uint,
uint);
void xfs_trans_dirty_buf(struct xfs_trans *, struct xfs_buf *);
bool xfs_trans_buf_is_dirty(struct xfs_buf *bp);
void xfs_trans_log_inode(xfs_trans_t *, struct xfs_inode *, uint);
int xfs_trans_commit(struct xfs_trans *);
int xfs_trans_roll(struct xfs_trans **);
int xfs_trans_roll_inode(struct xfs_trans **, struct xfs_inode *);
void xfs_trans_cancel(xfs_trans_t *);
[XFS] Move AIL pushing into it's own thread When many hundreds to thousands of threads all try to do simultaneous transactions and the log is in a tail-pushing situation (i.e. full), we can get multiple threads walking the AIL list and contending on the AIL lock. The AIL push is, in effect, a simple I/O dispatch algorithm complicated by the ordering constraints placed on it by the transaction subsystem. It really does not need multiple threads to push on it - even when only a single CPU is pushing the AIL, it can push the I/O out far faster that pretty much any disk subsystem can handle. So, to avoid contention problems stemming from multiple list walkers, move the list walk off into another thread and simply provide a "target" to push to. When a thread requires a push, it sets the target and wakes the push thread, then goes to sleep waiting for the required amount of space to become available in the log. This mechanism should also be a lot fairer under heavy load as the waiters will queue in arrival order, rather than queuing in "who completed a push first" order. Also, by moving the pushing to a separate thread we can do more effectively overload detection and prevention as we can keep context from loop iteration to loop iteration. That is, we can push only part of the list each loop and not have to loop back to the start of the list every time we run. This should also help by reducing the number of items we try to lock and/or push items that we cannot move. Note that this patch is not intended to solve the inefficiencies in the AIL structure and the associated issues with extremely large list contents. That needs to be addresses separately; parallel access would cause problems to any new structure as well, so I'm only aiming to isolate the structure from unbounded parallelism here. SGI-PV: 972759 SGI-Modid: xfs-linux-melb:xfs-kern:30371a Signed-off-by: David Chinner <dgc@sgi.com> Signed-off-by: Lachlan McIlroy <lachlan@sgi.com>
2008-02-05 09:13:32 +08:00
int xfs_trans_ail_init(struct xfs_mount *);
void xfs_trans_ail_destroy(struct xfs_mount *);
void xfs_trans_buf_set_type(struct xfs_trans *, struct xfs_buf *,
enum xfs_blft);
void xfs_trans_buf_copy_type(struct xfs_buf *dst_bp,
struct xfs_buf *src_bp);
extern struct kmem_cache *xfs_trans_cache;
xfs: periodically relog deferred intent items There's a subtle design flaw in the deferred log item code that can lead to pinning the log tail. Taking up the defer ops chain examples from the previous commit, we can get trapped in sequences like this: Caller hands us a transaction t0 with D0-D3 attached. The defer ops chain will look like the following if the transaction rolls succeed: t1: D0(t0), D1(t0), D2(t0), D3(t0) t2: d4(t1), d5(t1), D1(t0), D2(t0), D3(t0) t3: d5(t1), D1(t0), D2(t0), D3(t0) ... t9: d9(t7), D3(t0) t10: D3(t0) t11: d10(t10), d11(t10) t12: d11(t10) In transaction 9, we finish d9 and try to roll to t10 while holding onto an intent item for D3 that we logged in t0. The previous commit changed the order in which we place new defer ops in the defer ops processing chain to reduce the maximum chain length. Now make xfs_defer_finish_noroll capable of relogging the entire chain periodically so that we can always move the log tail forward. Most chains will never get relogged, except for operations that generate very long chains (large extents containing many blocks with different sharing levels) or are on filesystems with small logs and a lot of ongoing metadata updates. Callers are now required to ensure that the transaction reservation is large enough to handle logging done items and new intent items for the maximum possible chain length. Most callers are careful to keep the chain lengths low, so the overhead should be minimal. The decision to relog an intent item is made based on whether the intent was logged in a previous checkpoint, since there's no point in relogging an intent into the same checkpoint. Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com> Reviewed-by: Brian Foster <bfoster@redhat.com>
2020-09-28 07:18:13 +08:00
static inline struct xfs_log_item *
xfs_trans_item_relog(
struct xfs_log_item *lip,
struct xfs_trans *tp)
{
return lip->li_ops->iop_relog(lip, tp);
}
struct xfs_dquot;
int xfs_trans_alloc_inode(struct xfs_inode *ip, struct xfs_trans_res *resv,
unsigned int dblocks, unsigned int rblocks, bool force,
struct xfs_trans **tpp);
int xfs_trans_alloc_icreate(struct xfs_mount *mp, struct xfs_trans_res *resv,
struct xfs_dquot *udqp, struct xfs_dquot *gdqp,
struct xfs_dquot *pdqp, unsigned int dblocks,
struct xfs_trans **tpp);
int xfs_trans_alloc_ichange(struct xfs_inode *ip, struct xfs_dquot *udqp,
struct xfs_dquot *gdqp, struct xfs_dquot *pdqp, bool force,
struct xfs_trans **tpp);
int xfs_trans_alloc_dir(struct xfs_inode *dp, struct xfs_trans_res *resv,
struct xfs_inode *ip, unsigned int *dblocks,
struct xfs_trans **tpp, int *nospace_error);
static inline void
xfs_trans_set_context(
struct xfs_trans *tp)
{
ASSERT(current->journal_info == NULL);
tp->t_pflags = memalloc_nofs_save();
current->journal_info = tp;
}
static inline void
xfs_trans_clear_context(
struct xfs_trans *tp)
{
if (current->journal_info == tp) {
memalloc_nofs_restore(tp->t_pflags);
current->journal_info = NULL;
}
}
static inline void
xfs_trans_switch_context(
struct xfs_trans *old_tp,
struct xfs_trans *new_tp)
{
ASSERT(current->journal_info == old_tp);
new_tp->t_pflags = old_tp->t_pflags;
old_tp->t_pflags = 0;
current->journal_info = new_tp;
}
#endif /* __XFS_TRANS_H__ */