linux/mm/vmscan.c

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/*
* linux/mm/vmscan.c
*
* Copyright (C) 1991, 1992, 1993, 1994 Linus Torvalds
*
* Swap reorganised 29.12.95, Stephen Tweedie.
* kswapd added: 7.1.96 sct
* Removed kswapd_ctl limits, and swap out as many pages as needed
* to bring the system back to freepages.high: 2.4.97, Rik van Riel.
* Zone aware kswapd started 02/00, Kanoj Sarcar (kanoj@sgi.com).
* Multiqueue VM started 5.8.00, Rik van Riel.
*/
#include <linux/mm.h>
#include <linux/module.h>
#include <linux/slab.h>
#include <linux/kernel_stat.h>
#include <linux/swap.h>
#include <linux/pagemap.h>
#include <linux/init.h>
#include <linux/highmem.h>
#include <linux/file.h>
#include <linux/writeback.h>
#include <linux/blkdev.h>
#include <linux/buffer_head.h> /* for try_to_release_page(),
buffer_heads_over_limit */
#include <linux/mm_inline.h>
#include <linux/pagevec.h>
#include <linux/backing-dev.h>
#include <linux/rmap.h>
#include <linux/topology.h>
#include <linux/cpu.h>
#include <linux/cpuset.h>
#include <linux/notifier.h>
#include <linux/rwsem.h>
#include <asm/tlbflush.h>
#include <asm/div64.h>
#include <linux/swapops.h>
/* possible outcome of pageout() */
typedef enum {
/* failed to write page out, page is locked */
PAGE_KEEP,
/* move page to the active list, page is locked */
PAGE_ACTIVATE,
/* page has been sent to the disk successfully, page is unlocked */
PAGE_SUCCESS,
/* page is clean and locked */
PAGE_CLEAN,
} pageout_t;
struct scan_control {
/* Ask refill_inactive_zone, or shrink_cache to scan this many pages */
unsigned long nr_to_scan;
/* Incremented by the number of inactive pages that were scanned */
unsigned long nr_scanned;
/* Incremented by the number of pages reclaimed */
unsigned long nr_reclaimed;
unsigned long nr_mapped; /* From page_state */
/* How many pages shrink_cache() should reclaim */
int nr_to_reclaim;
/* Ask shrink_caches, or shrink_zone to scan at this priority */
unsigned int priority;
/* This context's GFP mask */
gfp_t gfp_mask;
int may_writepage;
[PATCH] VM: add may_swap flag to scan_control Here's the next round of these patches. These are totally different in an attempt to meet the "simpler" request after the last patches. For reference the earlier threads are: http://marc.theaimsgroup.com/?l=linux-kernel&m=110839604924587&w=2 http://marc.theaimsgroup.com/?l=linux-mm&m=111461480721249&w=2 This set of patches replaces my other vm- patches that are currently in -mm. So they're against 2.6.12-rc5-mm1 about half way through the -mm patchset. As I said already this patch is a lot simpler. The reclaim is turned on or off on a per-zone basis using a syscall. I haven't tested the x86 syscall, so it might be wrong. It uses the existing reclaim/pageout code with the small addition of a may_swap flag to scan_control (patch 1/4). I also added __GFP_NORECLAIM (patch 3/4) so that certain allocation types can be flagged to never cause reclaim. This was a deficiency that was in all of my earlier patch sets. Previously, doing a big buffered read would fill one zone with page cache and then start to reclaim from that same zone, leaving the other zones untouched. Adding some extra throttling on the reclaim was also required (patch 4/4). Without the machine would grind to a crawl when doing a "make -j" kernel build. Even with this patch the System Time is higher on average, but it seems tolerable. Here are some numbers for kernbench runs on a 2-node, 4cpu, 8Gig RAM Altix in the "make -j" run: wall user sys %cpu ctx sw. sleeps ---- ---- --- ---- ------ ------ No patch 1009 1384 847 258 298170 504402 w/patch, no reclaim 880 1376 667 288 254064 396745 w/patch & reclaim 1079 1385 926 252 291625 548873 These numbers are the average of 2 runs of 3 "make -j" runs done right after system boot. Run-to-run variability for "make -j" is huge, so these numbers aren't terribly useful except to seee that with reclaim the benchmark still finishes in a reasonable amount of time. I also looked at the NUMA hit/miss stats for the "make -j" runs and the reclaim doesn't make any difference when the machine is thrashing away. Doing a "make -j8" on a single node that is filled with page cache pages takes 700 seconds with reclaim turned on and 735 seconds without reclaim (due to remote memory accesses). The simple zone_reclaim syscall program is at http://www.bork.org/~mort/sgi/zone_reclaim.c This patch: This adds an extra switch to the scan_control struct. It simply lets the reclaim code know if its allowed to swap pages out. This was required for a simple per-zone reclaimer. Without this addition pages would be swapped out as soon as a zone ran out of memory and the early reclaim kicked in. Signed-off-by: Martin Hicks <mort@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-22 08:14:40 +08:00
/* Can pages be swapped as part of reclaim? */
int may_swap;
/* This context's SWAP_CLUSTER_MAX. If freeing memory for
* suspend, we effectively ignore SWAP_CLUSTER_MAX.
* In this context, it doesn't matter that we scan the
* whole list at once. */
int swap_cluster_max;
};
/*
* The list of shrinker callbacks used by to apply pressure to
* ageable caches.
*/
struct shrinker {
shrinker_t shrinker;
struct list_head list;
int seeks; /* seeks to recreate an obj */
long nr; /* objs pending delete */
};
#define lru_to_page(_head) (list_entry((_head)->prev, struct page, lru))
#ifdef ARCH_HAS_PREFETCH
#define prefetch_prev_lru_page(_page, _base, _field) \
do { \
if ((_page)->lru.prev != _base) { \
struct page *prev; \
\
prev = lru_to_page(&(_page->lru)); \
prefetch(&prev->_field); \
} \
} while (0)
#else
#define prefetch_prev_lru_page(_page, _base, _field) do { } while (0)
#endif
#ifdef ARCH_HAS_PREFETCHW
#define prefetchw_prev_lru_page(_page, _base, _field) \
do { \
if ((_page)->lru.prev != _base) { \
struct page *prev; \
\
prev = lru_to_page(&(_page->lru)); \
prefetchw(&prev->_field); \
} \
} while (0)
#else
#define prefetchw_prev_lru_page(_page, _base, _field) do { } while (0)
#endif
/*
* From 0 .. 100. Higher means more swappy.
*/
int vm_swappiness = 60;
static long total_memory;
static LIST_HEAD(shrinker_list);
static DECLARE_RWSEM(shrinker_rwsem);
/*
* Add a shrinker callback to be called from the vm
*/
struct shrinker *set_shrinker(int seeks, shrinker_t theshrinker)
{
struct shrinker *shrinker;
shrinker = kmalloc(sizeof(*shrinker), GFP_KERNEL);
if (shrinker) {
shrinker->shrinker = theshrinker;
shrinker->seeks = seeks;
shrinker->nr = 0;
down_write(&shrinker_rwsem);
list_add_tail(&shrinker->list, &shrinker_list);
up_write(&shrinker_rwsem);
}
return shrinker;
}
EXPORT_SYMBOL(set_shrinker);
/*
* Remove one
*/
void remove_shrinker(struct shrinker *shrinker)
{
down_write(&shrinker_rwsem);
list_del(&shrinker->list);
up_write(&shrinker_rwsem);
kfree(shrinker);
}
EXPORT_SYMBOL(remove_shrinker);
#define SHRINK_BATCH 128
/*
* Call the shrink functions to age shrinkable caches
*
* Here we assume it costs one seek to replace a lru page and that it also
* takes a seek to recreate a cache object. With this in mind we age equal
* percentages of the lru and ageable caches. This should balance the seeks
* generated by these structures.
*
* If the vm encounted mapped pages on the LRU it increase the pressure on
* slab to avoid swapping.
*
* We do weird things to avoid (scanned*seeks*entries) overflowing 32 bits.
*
* `lru_pages' represents the number of on-LRU pages in all the zones which
* are eligible for the caller's allocation attempt. It is used for balancing
* slab reclaim versus page reclaim.
*
* Returns the number of slab objects which we shrunk.
*/
static int shrink_slab(unsigned long scanned, gfp_t gfp_mask,
unsigned long lru_pages)
{
struct shrinker *shrinker;
int ret = 0;
if (scanned == 0)
scanned = SWAP_CLUSTER_MAX;
if (!down_read_trylock(&shrinker_rwsem))
return 1; /* Assume we'll be able to shrink next time */
list_for_each_entry(shrinker, &shrinker_list, list) {
unsigned long long delta;
unsigned long total_scan;
delta = (4 * scanned) / shrinker->seeks;
delta *= (*shrinker->shrinker)(0, gfp_mask);
do_div(delta, lru_pages + 1);
shrinker->nr += delta;
if (shrinker->nr < 0)
shrinker->nr = LONG_MAX; /* It wrapped! */
total_scan = shrinker->nr;
shrinker->nr = 0;
while (total_scan >= SHRINK_BATCH) {
long this_scan = SHRINK_BATCH;
int shrink_ret;
int nr_before;
nr_before = (*shrinker->shrinker)(0, gfp_mask);
shrink_ret = (*shrinker->shrinker)(this_scan, gfp_mask);
if (shrink_ret == -1)
break;
if (shrink_ret < nr_before)
ret += nr_before - shrink_ret;
mod_page_state(slabs_scanned, this_scan);
total_scan -= this_scan;
cond_resched();
}
shrinker->nr += total_scan;
}
up_read(&shrinker_rwsem);
return ret;
}
/* Called without lock on whether page is mapped, so answer is unstable */
static inline int page_mapping_inuse(struct page *page)
{
struct address_space *mapping;
/* Page is in somebody's page tables. */
if (page_mapped(page))
return 1;
/* Be more reluctant to reclaim swapcache than pagecache */
if (PageSwapCache(page))
return 1;
mapping = page_mapping(page);
if (!mapping)
return 0;
/* File is mmap'd by somebody? */
return mapping_mapped(mapping);
}
static inline int is_page_cache_freeable(struct page *page)
{
return page_count(page) - !!PagePrivate(page) == 2;
}
static int may_write_to_queue(struct backing_dev_info *bdi)
{
if (current_is_kswapd())
return 1;
if (current_is_pdflush()) /* This is unlikely, but why not... */
return 1;
if (!bdi_write_congested(bdi))
return 1;
if (bdi == current->backing_dev_info)
return 1;
return 0;
}
/*
* We detected a synchronous write error writing a page out. Probably
* -ENOSPC. We need to propagate that into the address_space for a subsequent
* fsync(), msync() or close().
*
* The tricky part is that after writepage we cannot touch the mapping: nothing
* prevents it from being freed up. But we have a ref on the page and once
* that page is locked, the mapping is pinned.
*
* We're allowed to run sleeping lock_page() here because we know the caller has
* __GFP_FS.
*/
static void handle_write_error(struct address_space *mapping,
struct page *page, int error)
{
lock_page(page);
if (page_mapping(page) == mapping) {
if (error == -ENOSPC)
set_bit(AS_ENOSPC, &mapping->flags);
else
set_bit(AS_EIO, &mapping->flags);
}
unlock_page(page);
}
/*
* pageout is called by shrink_list() for each dirty page. Calls ->writepage().
*/
static pageout_t pageout(struct page *page, struct address_space *mapping)
{
/*
* If the page is dirty, only perform writeback if that write
* will be non-blocking. To prevent this allocation from being
* stalled by pagecache activity. But note that there may be
* stalls if we need to run get_block(). We could test
* PagePrivate for that.
*
* If this process is currently in generic_file_write() against
* this page's queue, we can perform writeback even if that
* will block.
*
* If the page is swapcache, write it back even if that would
* block, for some throttling. This happens by accident, because
* swap_backing_dev_info is bust: it doesn't reflect the
* congestion state of the swapdevs. Easy to fix, if needed.
* See swapfile.c:page_queue_congested().
*/
if (!is_page_cache_freeable(page))
return PAGE_KEEP;
if (!mapping) {
/*
* Some data journaling orphaned pages can have
* page->mapping == NULL while being dirty with clean buffers.
*/
if (PagePrivate(page)) {
if (try_to_free_buffers(page)) {
ClearPageDirty(page);
printk("%s: orphaned page\n", __FUNCTION__);
return PAGE_CLEAN;
}
}
return PAGE_KEEP;
}
if (mapping->a_ops->writepage == NULL)
return PAGE_ACTIVATE;
if (!may_write_to_queue(mapping->backing_dev_info))
return PAGE_KEEP;
if (clear_page_dirty_for_io(page)) {
int res;
struct writeback_control wbc = {
.sync_mode = WB_SYNC_NONE,
.nr_to_write = SWAP_CLUSTER_MAX,
.nonblocking = 1,
.for_reclaim = 1,
};
SetPageReclaim(page);
res = mapping->a_ops->writepage(page, &wbc);
if (res < 0)
handle_write_error(mapping, page, res);
if (res == WRITEPAGE_ACTIVATE) {
ClearPageReclaim(page);
return PAGE_ACTIVATE;
}
if (!PageWriteback(page)) {
/* synchronous write or broken a_ops? */
ClearPageReclaim(page);
}
return PAGE_SUCCESS;
}
return PAGE_CLEAN;
}
/*
* shrink_list adds the number of reclaimed pages to sc->nr_reclaimed
*/
static int shrink_list(struct list_head *page_list, struct scan_control *sc)
{
LIST_HEAD(ret_pages);
struct pagevec freed_pvec;
int pgactivate = 0;
int reclaimed = 0;
cond_resched();
pagevec_init(&freed_pvec, 1);
while (!list_empty(page_list)) {
struct address_space *mapping;
struct page *page;
int may_enter_fs;
int referenced;
cond_resched();
page = lru_to_page(page_list);
list_del(&page->lru);
if (TestSetPageLocked(page))
goto keep;
BUG_ON(PageActive(page));
sc->nr_scanned++;
/* Double the slab pressure for mapped and swapcache pages */
if (page_mapped(page) || PageSwapCache(page))
sc->nr_scanned++;
if (PageWriteback(page))
goto keep_locked;
referenced = page_referenced(page, 1);
/* In active use or really unfreeable? Activate it. */
if (referenced && page_mapping_inuse(page))
goto activate_locked;
#ifdef CONFIG_SWAP
/*
* Anonymous process memory has backing store?
* Try to allocate it some swap space here.
*/
if (PageAnon(page) && !PageSwapCache(page)) {
if (!sc->may_swap)
goto keep_locked;
if (!add_to_swap(page))
goto activate_locked;
}
#endif /* CONFIG_SWAP */
mapping = page_mapping(page);
may_enter_fs = (sc->gfp_mask & __GFP_FS) ||
(PageSwapCache(page) && (sc->gfp_mask & __GFP_IO));
/*
* The page is mapped into the page tables of one or more
* processes. Try to unmap it here.
*/
if (page_mapped(page) && mapping) {
switch (try_to_unmap(page)) {
case SWAP_FAIL:
goto activate_locked;
case SWAP_AGAIN:
goto keep_locked;
case SWAP_SUCCESS:
; /* try to free the page below */
}
}
if (PageDirty(page)) {
if (referenced)
goto keep_locked;
if (!may_enter_fs)
goto keep_locked;
if (laptop_mode && !sc->may_writepage)
goto keep_locked;
/* Page is dirty, try to write it out here */
switch(pageout(page, mapping)) {
case PAGE_KEEP:
goto keep_locked;
case PAGE_ACTIVATE:
goto activate_locked;
case PAGE_SUCCESS:
if (PageWriteback(page) || PageDirty(page))
goto keep;
/*
* A synchronous write - probably a ramdisk. Go
* ahead and try to reclaim the page.
*/
if (TestSetPageLocked(page))
goto keep;
if (PageDirty(page) || PageWriteback(page))
goto keep_locked;
mapping = page_mapping(page);
case PAGE_CLEAN:
; /* try to free the page below */
}
}
/*
* If the page has buffers, try to free the buffer mappings
* associated with this page. If we succeed we try to free
* the page as well.
*
* We do this even if the page is PageDirty().
* try_to_release_page() does not perform I/O, but it is
* possible for a page to have PageDirty set, but it is actually
* clean (all its buffers are clean). This happens if the
* buffers were written out directly, with submit_bh(). ext3
* will do this, as well as the blockdev mapping.
* try_to_release_page() will discover that cleanness and will
* drop the buffers and mark the page clean - it can be freed.
*
* Rarely, pages can have buffers and no ->mapping. These are
* the pages which were not successfully invalidated in
* truncate_complete_page(). We try to drop those buffers here
* and if that worked, and the page is no longer mapped into
* process address space (page_count == 1) it can be freed.
* Otherwise, leave the page on the LRU so it is swappable.
*/
if (PagePrivate(page)) {
if (!try_to_release_page(page, sc->gfp_mask))
goto activate_locked;
if (!mapping && page_count(page) == 1)
goto free_it;
}
if (!mapping)
goto keep_locked; /* truncate got there first */
write_lock_irq(&mapping->tree_lock);
/*
* The non-racy check for busy page. It is critical to check
* PageDirty _after_ making sure that the page is freeable and
* not in use by anybody. (pagecache + us == 2)
*/
if (unlikely(page_count(page) != 2))
goto cannot_free;
smp_rmb();
if (unlikely(PageDirty(page)))
goto cannot_free;
#ifdef CONFIG_SWAP
if (PageSwapCache(page)) {
[PATCH] mm: split page table lock Christoph Lameter demonstrated very poor scalability on the SGI 512-way, with a many-threaded application which concurrently initializes different parts of a large anonymous area. This patch corrects that, by using a separate spinlock per page table page, to guard the page table entries in that page, instead of using the mm's single page_table_lock. (But even then, page_table_lock is still used to guard page table allocation, and anon_vma allocation.) In this implementation, the spinlock is tucked inside the struct page of the page table page: with a BUILD_BUG_ON in case it overflows - which it would in the case of 32-bit PA-RISC with spinlock debugging enabled. Splitting the lock is not quite for free: another cacheline access. Ideally, I suppose we would use split ptlock only for multi-threaded processes on multi-cpu machines; but deciding that dynamically would have its own costs. So for now enable it by config, at some number of cpus - since the Kconfig language doesn't support inequalities, let preprocessor compare that with NR_CPUS. But I don't think it's worth being user-configurable: for good testing of both split and unsplit configs, split now at 4 cpus, and perhaps change that to 8 later. There is a benefit even for singly threaded processes: kswapd can be attacking one part of the mm while another part is busy faulting. Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-10-30 09:16:40 +08:00
swp_entry_t swap = { .val = page_private(page) };
__delete_from_swap_cache(page);
write_unlock_irq(&mapping->tree_lock);
swap_free(swap);
__put_page(page); /* The pagecache ref */
goto free_it;
}
#endif /* CONFIG_SWAP */
__remove_from_page_cache(page);
write_unlock_irq(&mapping->tree_lock);
__put_page(page);
free_it:
unlock_page(page);
reclaimed++;
if (!pagevec_add(&freed_pvec, page))
__pagevec_release_nonlru(&freed_pvec);
continue;
cannot_free:
write_unlock_irq(&mapping->tree_lock);
goto keep_locked;
activate_locked:
SetPageActive(page);
pgactivate++;
keep_locked:
unlock_page(page);
keep:
list_add(&page->lru, &ret_pages);
BUG_ON(PageLRU(page));
}
list_splice(&ret_pages, page_list);
if (pagevec_count(&freed_pvec))
__pagevec_release_nonlru(&freed_pvec);
mod_page_state(pgactivate, pgactivate);
sc->nr_reclaimed += reclaimed;
return reclaimed;
}
/*
* zone->lru_lock is heavily contended. Some of the functions that
* shrink the lists perform better by taking out a batch of pages
* and working on them outside the LRU lock.
*
* For pagecache intensive workloads, this function is the hottest
* spot in the kernel (apart from copy_*_user functions).
*
* Appropriate locks must be held before calling this function.
*
* @nr_to_scan: The number of pages to look through on the list.
* @src: The LRU list to pull pages off.
* @dst: The temp list to put pages on to.
* @scanned: The number of pages that were scanned.
*
* returns how many pages were moved onto *@dst.
*/
static int isolate_lru_pages(int nr_to_scan, struct list_head *src,
struct list_head *dst, int *scanned)
{
int nr_taken = 0;
struct page *page;
int scan = 0;
while (scan++ < nr_to_scan && !list_empty(src)) {
page = lru_to_page(src);
prefetchw_prev_lru_page(page, src, flags);
if (!TestClearPageLRU(page))
BUG();
list_del(&page->lru);
if (get_page_testone(page)) {
/*
* It is being freed elsewhere
*/
__put_page(page);
SetPageLRU(page);
list_add(&page->lru, src);
continue;
} else {
list_add(&page->lru, dst);
nr_taken++;
}
}
*scanned = scan;
return nr_taken;
}
/*
* shrink_cache() adds the number of pages reclaimed to sc->nr_reclaimed
*/
static void shrink_cache(struct zone *zone, struct scan_control *sc)
{
LIST_HEAD(page_list);
struct pagevec pvec;
int max_scan = sc->nr_to_scan;
pagevec_init(&pvec, 1);
lru_add_drain();
spin_lock_irq(&zone->lru_lock);
while (max_scan > 0) {
struct page *page;
int nr_taken;
int nr_scan;
int nr_freed;
nr_taken = isolate_lru_pages(sc->swap_cluster_max,
&zone->inactive_list,
&page_list, &nr_scan);
zone->nr_inactive -= nr_taken;
zone->pages_scanned += nr_scan;
spin_unlock_irq(&zone->lru_lock);
if (nr_taken == 0)
goto done;
max_scan -= nr_scan;
if (current_is_kswapd())
mod_page_state_zone(zone, pgscan_kswapd, nr_scan);
else
mod_page_state_zone(zone, pgscan_direct, nr_scan);
nr_freed = shrink_list(&page_list, sc);
if (current_is_kswapd())
mod_page_state(kswapd_steal, nr_freed);
mod_page_state_zone(zone, pgsteal, nr_freed);
sc->nr_to_reclaim -= nr_freed;
spin_lock_irq(&zone->lru_lock);
/*
* Put back any unfreeable pages.
*/
while (!list_empty(&page_list)) {
page = lru_to_page(&page_list);
if (TestSetPageLRU(page))
BUG();
list_del(&page->lru);
if (PageActive(page))
add_page_to_active_list(zone, page);
else
add_page_to_inactive_list(zone, page);
if (!pagevec_add(&pvec, page)) {
spin_unlock_irq(&zone->lru_lock);
__pagevec_release(&pvec);
spin_lock_irq(&zone->lru_lock);
}
}
}
spin_unlock_irq(&zone->lru_lock);
done:
pagevec_release(&pvec);
}
/*
* This moves pages from the active list to the inactive list.
*
* We move them the other way if the page is referenced by one or more
* processes, from rmap.
*
* If the pages are mostly unmapped, the processing is fast and it is
* appropriate to hold zone->lru_lock across the whole operation. But if
* the pages are mapped, the processing is slow (page_referenced()) so we
* should drop zone->lru_lock around each page. It's impossible to balance
* this, so instead we remove the pages from the LRU while processing them.
* It is safe to rely on PG_active against the non-LRU pages in here because
* nobody will play with that bit on a non-LRU page.
*
* The downside is that we have to touch page->_count against each page.
* But we had to alter page->flags anyway.
*/
static void
refill_inactive_zone(struct zone *zone, struct scan_control *sc)
{
int pgmoved;
int pgdeactivate = 0;
int pgscanned;
int nr_pages = sc->nr_to_scan;
LIST_HEAD(l_hold); /* The pages which were snipped off */
LIST_HEAD(l_inactive); /* Pages to go onto the inactive_list */
LIST_HEAD(l_active); /* Pages to go onto the active_list */
struct page *page;
struct pagevec pvec;
int reclaim_mapped = 0;
long mapped_ratio;
long distress;
long swap_tendency;
lru_add_drain();
spin_lock_irq(&zone->lru_lock);
pgmoved = isolate_lru_pages(nr_pages, &zone->active_list,
&l_hold, &pgscanned);
zone->pages_scanned += pgscanned;
zone->nr_active -= pgmoved;
spin_unlock_irq(&zone->lru_lock);
/*
* `distress' is a measure of how much trouble we're having reclaiming
* pages. 0 -> no problems. 100 -> great trouble.
*/
distress = 100 >> zone->prev_priority;
/*
* The point of this algorithm is to decide when to start reclaiming
* mapped memory instead of just pagecache. Work out how much memory
* is mapped.
*/
mapped_ratio = (sc->nr_mapped * 100) / total_memory;
/*
* Now decide how much we really want to unmap some pages. The mapped
* ratio is downgraded - just because there's a lot of mapped memory
* doesn't necessarily mean that page reclaim isn't succeeding.
*
* The distress ratio is important - we don't want to start going oom.
*
* A 100% value of vm_swappiness overrides this algorithm altogether.
*/
swap_tendency = mapped_ratio / 2 + distress + vm_swappiness;
/*
* Now use this metric to decide whether to start moving mapped memory
* onto the inactive list.
*/
if (swap_tendency >= 100)
reclaim_mapped = 1;
while (!list_empty(&l_hold)) {
cond_resched();
page = lru_to_page(&l_hold);
list_del(&page->lru);
if (page_mapped(page)) {
if (!reclaim_mapped ||
(total_swap_pages == 0 && PageAnon(page)) ||
page_referenced(page, 0)) {
list_add(&page->lru, &l_active);
continue;
}
}
list_add(&page->lru, &l_inactive);
}
pagevec_init(&pvec, 1);
pgmoved = 0;
spin_lock_irq(&zone->lru_lock);
while (!list_empty(&l_inactive)) {
page = lru_to_page(&l_inactive);
prefetchw_prev_lru_page(page, &l_inactive, flags);
if (TestSetPageLRU(page))
BUG();
if (!TestClearPageActive(page))
BUG();
list_move(&page->lru, &zone->inactive_list);
pgmoved++;
if (!pagevec_add(&pvec, page)) {
zone->nr_inactive += pgmoved;
spin_unlock_irq(&zone->lru_lock);
pgdeactivate += pgmoved;
pgmoved = 0;
if (buffer_heads_over_limit)
pagevec_strip(&pvec);
__pagevec_release(&pvec);
spin_lock_irq(&zone->lru_lock);
}
}
zone->nr_inactive += pgmoved;
pgdeactivate += pgmoved;
if (buffer_heads_over_limit) {
spin_unlock_irq(&zone->lru_lock);
pagevec_strip(&pvec);
spin_lock_irq(&zone->lru_lock);
}
pgmoved = 0;
while (!list_empty(&l_active)) {
page = lru_to_page(&l_active);
prefetchw_prev_lru_page(page, &l_active, flags);
if (TestSetPageLRU(page))
BUG();
BUG_ON(!PageActive(page));
list_move(&page->lru, &zone->active_list);
pgmoved++;
if (!pagevec_add(&pvec, page)) {
zone->nr_active += pgmoved;
pgmoved = 0;
spin_unlock_irq(&zone->lru_lock);
__pagevec_release(&pvec);
spin_lock_irq(&zone->lru_lock);
}
}
zone->nr_active += pgmoved;
spin_unlock_irq(&zone->lru_lock);
pagevec_release(&pvec);
mod_page_state_zone(zone, pgrefill, pgscanned);
mod_page_state(pgdeactivate, pgdeactivate);
}
/*
* This is a basic per-zone page freer. Used by both kswapd and direct reclaim.
*/
static void
shrink_zone(struct zone *zone, struct scan_control *sc)
{
unsigned long nr_active;
unsigned long nr_inactive;
atomic_inc(&zone->reclaim_in_progress);
/*
* Add one to `nr_to_scan' just to make sure that the kernel will
* slowly sift through the active list.
*/
zone->nr_scan_active += (zone->nr_active >> sc->priority) + 1;
nr_active = zone->nr_scan_active;
if (nr_active >= sc->swap_cluster_max)
zone->nr_scan_active = 0;
else
nr_active = 0;
zone->nr_scan_inactive += (zone->nr_inactive >> sc->priority) + 1;
nr_inactive = zone->nr_scan_inactive;
if (nr_inactive >= sc->swap_cluster_max)
zone->nr_scan_inactive = 0;
else
nr_inactive = 0;
sc->nr_to_reclaim = sc->swap_cluster_max;
while (nr_active || nr_inactive) {
if (nr_active) {
sc->nr_to_scan = min(nr_active,
(unsigned long)sc->swap_cluster_max);
nr_active -= sc->nr_to_scan;
refill_inactive_zone(zone, sc);
}
if (nr_inactive) {
sc->nr_to_scan = min(nr_inactive,
(unsigned long)sc->swap_cluster_max);
nr_inactive -= sc->nr_to_scan;
shrink_cache(zone, sc);
if (sc->nr_to_reclaim <= 0)
break;
}
}
throttle_vm_writeout();
atomic_dec(&zone->reclaim_in_progress);
}
/*
* This is the direct reclaim path, for page-allocating processes. We only
* try to reclaim pages from zones which will satisfy the caller's allocation
* request.
*
* We reclaim from a zone even if that zone is over pages_high. Because:
* a) The caller may be trying to free *extra* pages to satisfy a higher-order
* allocation or
* b) The zones may be over pages_high but they must go *over* pages_high to
* satisfy the `incremental min' zone defense algorithm.
*
* Returns the number of reclaimed pages.
*
* If a zone is deemed to be full of pinned pages then just give it a light
* scan then give up on it.
*/
static void
shrink_caches(struct zone **zones, struct scan_control *sc)
{
int i;
for (i = 0; zones[i] != NULL; i++) {
struct zone *zone = zones[i];
if (zone->present_pages == 0)
continue;
[PATCH] cpusets: formalize intermediate GFP_KERNEL containment This patch makes use of the previously underutilized cpuset flag 'mem_exclusive' to provide what amounts to another layer of memory placement resolution. With this patch, there are now the following four layers of memory placement available: 1) The whole system (interrupt and GFP_ATOMIC allocations can use this), 2) The nearest enclosing mem_exclusive cpuset (GFP_KERNEL allocations can use), 3) The current tasks cpuset (GFP_USER allocations constrained to here), and 4) Specific node placement, using mbind and set_mempolicy. These nest - each layer is a subset (same or within) of the previous. Layer (2) above is new, with this patch. The call used to check whether a zone (its node, actually) is in a cpuset (in its mems_allowed, actually) is extended to take a gfp_mask argument, and its logic is extended, in the case that __GFP_HARDWALL is not set in the flag bits, to look up the cpuset hierarchy for the nearest enclosing mem_exclusive cpuset, to determine if placement is allowed. The definition of GFP_USER, which used to be identical to GFP_KERNEL, is changed to also set the __GFP_HARDWALL bit, in the previous cpuset_gfp_hardwall_flag patch. GFP_ATOMIC and GFP_KERNEL allocations will stay within the current tasks cpuset, so long as any node therein is not too tight on memory, but will escape to the larger layer, if need be. The intended use is to allow something like a batch manager to handle several jobs, each job in its own cpuset, but using common kernel memory for caches and such. Swapper and oom_kill activity is also constrained to Layer (2). A task in or below one mem_exclusive cpuset should not cause swapping on nodes in another non-overlapping mem_exclusive cpuset, nor provoke oom_killing of a task in another such cpuset. Heavy use of kernel memory for i/o caching and such by one job should not impact the memory available to jobs in other non-overlapping mem_exclusive cpusets. This patch enables providing hardwall, inescapable cpusets for memory allocations of each job, while sharing kernel memory allocations between several jobs, in an enclosing mem_exclusive cpuset. Like Dinakar's patch earlier to enable administering sched domains using the cpu_exclusive flag, this patch also provides a useful meaning to a cpuset flag that had previously done nothing much useful other than restrict what cpuset configurations were allowed. Signed-off-by: Paul Jackson <pj@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-09-07 06:18:12 +08:00
if (!cpuset_zone_allowed(zone, __GFP_HARDWALL))
continue;
zone->temp_priority = sc->priority;
if (zone->prev_priority > sc->priority)
zone->prev_priority = sc->priority;
if (zone->all_unreclaimable && sc->priority != DEF_PRIORITY)
continue; /* Let kswapd poll it */
shrink_zone(zone, sc);
}
}
/*
* This is the main entry point to direct page reclaim.
*
* If a full scan of the inactive list fails to free enough memory then we
* are "out of memory" and something needs to be killed.
*
* If the caller is !__GFP_FS then the probability of a failure is reasonably
* high - the zone may be full of dirty or under-writeback pages, which this
* caller can't do much about. We kick pdflush and take explicit naps in the
* hope that some of these pages can be written. But if the allocating task
* holds filesystem locks which prevent writeout this might not work, and the
* allocation attempt will fail.
*/
int try_to_free_pages(struct zone **zones, gfp_t gfp_mask)
{
int priority;
int ret = 0;
int total_scanned = 0, total_reclaimed = 0;
struct reclaim_state *reclaim_state = current->reclaim_state;
struct scan_control sc;
unsigned long lru_pages = 0;
int i;
sc.gfp_mask = gfp_mask;
sc.may_writepage = 0;
[PATCH] VM: add may_swap flag to scan_control Here's the next round of these patches. These are totally different in an attempt to meet the "simpler" request after the last patches. For reference the earlier threads are: http://marc.theaimsgroup.com/?l=linux-kernel&m=110839604924587&w=2 http://marc.theaimsgroup.com/?l=linux-mm&m=111461480721249&w=2 This set of patches replaces my other vm- patches that are currently in -mm. So they're against 2.6.12-rc5-mm1 about half way through the -mm patchset. As I said already this patch is a lot simpler. The reclaim is turned on or off on a per-zone basis using a syscall. I haven't tested the x86 syscall, so it might be wrong. It uses the existing reclaim/pageout code with the small addition of a may_swap flag to scan_control (patch 1/4). I also added __GFP_NORECLAIM (patch 3/4) so that certain allocation types can be flagged to never cause reclaim. This was a deficiency that was in all of my earlier patch sets. Previously, doing a big buffered read would fill one zone with page cache and then start to reclaim from that same zone, leaving the other zones untouched. Adding some extra throttling on the reclaim was also required (patch 4/4). Without the machine would grind to a crawl when doing a "make -j" kernel build. Even with this patch the System Time is higher on average, but it seems tolerable. Here are some numbers for kernbench runs on a 2-node, 4cpu, 8Gig RAM Altix in the "make -j" run: wall user sys %cpu ctx sw. sleeps ---- ---- --- ---- ------ ------ No patch 1009 1384 847 258 298170 504402 w/patch, no reclaim 880 1376 667 288 254064 396745 w/patch & reclaim 1079 1385 926 252 291625 548873 These numbers are the average of 2 runs of 3 "make -j" runs done right after system boot. Run-to-run variability for "make -j" is huge, so these numbers aren't terribly useful except to seee that with reclaim the benchmark still finishes in a reasonable amount of time. I also looked at the NUMA hit/miss stats for the "make -j" runs and the reclaim doesn't make any difference when the machine is thrashing away. Doing a "make -j8" on a single node that is filled with page cache pages takes 700 seconds with reclaim turned on and 735 seconds without reclaim (due to remote memory accesses). The simple zone_reclaim syscall program is at http://www.bork.org/~mort/sgi/zone_reclaim.c This patch: This adds an extra switch to the scan_control struct. It simply lets the reclaim code know if its allowed to swap pages out. This was required for a simple per-zone reclaimer. Without this addition pages would be swapped out as soon as a zone ran out of memory and the early reclaim kicked in. Signed-off-by: Martin Hicks <mort@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-22 08:14:40 +08:00
sc.may_swap = 1;
inc_page_state(allocstall);
for (i = 0; zones[i] != NULL; i++) {
struct zone *zone = zones[i];
[PATCH] cpusets: formalize intermediate GFP_KERNEL containment This patch makes use of the previously underutilized cpuset flag 'mem_exclusive' to provide what amounts to another layer of memory placement resolution. With this patch, there are now the following four layers of memory placement available: 1) The whole system (interrupt and GFP_ATOMIC allocations can use this), 2) The nearest enclosing mem_exclusive cpuset (GFP_KERNEL allocations can use), 3) The current tasks cpuset (GFP_USER allocations constrained to here), and 4) Specific node placement, using mbind and set_mempolicy. These nest - each layer is a subset (same or within) of the previous. Layer (2) above is new, with this patch. The call used to check whether a zone (its node, actually) is in a cpuset (in its mems_allowed, actually) is extended to take a gfp_mask argument, and its logic is extended, in the case that __GFP_HARDWALL is not set in the flag bits, to look up the cpuset hierarchy for the nearest enclosing mem_exclusive cpuset, to determine if placement is allowed. The definition of GFP_USER, which used to be identical to GFP_KERNEL, is changed to also set the __GFP_HARDWALL bit, in the previous cpuset_gfp_hardwall_flag patch. GFP_ATOMIC and GFP_KERNEL allocations will stay within the current tasks cpuset, so long as any node therein is not too tight on memory, but will escape to the larger layer, if need be. The intended use is to allow something like a batch manager to handle several jobs, each job in its own cpuset, but using common kernel memory for caches and such. Swapper and oom_kill activity is also constrained to Layer (2). A task in or below one mem_exclusive cpuset should not cause swapping on nodes in another non-overlapping mem_exclusive cpuset, nor provoke oom_killing of a task in another such cpuset. Heavy use of kernel memory for i/o caching and such by one job should not impact the memory available to jobs in other non-overlapping mem_exclusive cpusets. This patch enables providing hardwall, inescapable cpusets for memory allocations of each job, while sharing kernel memory allocations between several jobs, in an enclosing mem_exclusive cpuset. Like Dinakar's patch earlier to enable administering sched domains using the cpu_exclusive flag, this patch also provides a useful meaning to a cpuset flag that had previously done nothing much useful other than restrict what cpuset configurations were allowed. Signed-off-by: Paul Jackson <pj@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-09-07 06:18:12 +08:00
if (!cpuset_zone_allowed(zone, __GFP_HARDWALL))
continue;
zone->temp_priority = DEF_PRIORITY;
lru_pages += zone->nr_active + zone->nr_inactive;
}
for (priority = DEF_PRIORITY; priority >= 0; priority--) {
sc.nr_mapped = read_page_state(nr_mapped);
sc.nr_scanned = 0;
sc.nr_reclaimed = 0;
sc.priority = priority;
sc.swap_cluster_max = SWAP_CLUSTER_MAX;
if (!priority)
disable_swap_token();
shrink_caches(zones, &sc);
shrink_slab(sc.nr_scanned, gfp_mask, lru_pages);
if (reclaim_state) {
sc.nr_reclaimed += reclaim_state->reclaimed_slab;
reclaim_state->reclaimed_slab = 0;
}
total_scanned += sc.nr_scanned;
total_reclaimed += sc.nr_reclaimed;
if (total_reclaimed >= sc.swap_cluster_max) {
ret = 1;
goto out;
}
/*
* Try to write back as many pages as we just scanned. This
* tends to cause slow streaming writers to write data to the
* disk smoothly, at the dirtying rate, which is nice. But
* that's undesirable in laptop mode, where we *want* lumpy
* writeout. So in laptop mode, write out the whole world.
*/
if (total_scanned > sc.swap_cluster_max + sc.swap_cluster_max/2) {
wakeup_pdflush(laptop_mode ? 0 : total_scanned);
sc.may_writepage = 1;
}
/* Take a nap, wait for some writeback to complete */
if (sc.nr_scanned && priority < DEF_PRIORITY - 2)
blk_congestion_wait(WRITE, HZ/10);
}
out:
for (i = 0; zones[i] != 0; i++) {
struct zone *zone = zones[i];
[PATCH] cpusets: formalize intermediate GFP_KERNEL containment This patch makes use of the previously underutilized cpuset flag 'mem_exclusive' to provide what amounts to another layer of memory placement resolution. With this patch, there are now the following four layers of memory placement available: 1) The whole system (interrupt and GFP_ATOMIC allocations can use this), 2) The nearest enclosing mem_exclusive cpuset (GFP_KERNEL allocations can use), 3) The current tasks cpuset (GFP_USER allocations constrained to here), and 4) Specific node placement, using mbind and set_mempolicy. These nest - each layer is a subset (same or within) of the previous. Layer (2) above is new, with this patch. The call used to check whether a zone (its node, actually) is in a cpuset (in its mems_allowed, actually) is extended to take a gfp_mask argument, and its logic is extended, in the case that __GFP_HARDWALL is not set in the flag bits, to look up the cpuset hierarchy for the nearest enclosing mem_exclusive cpuset, to determine if placement is allowed. The definition of GFP_USER, which used to be identical to GFP_KERNEL, is changed to also set the __GFP_HARDWALL bit, in the previous cpuset_gfp_hardwall_flag patch. GFP_ATOMIC and GFP_KERNEL allocations will stay within the current tasks cpuset, so long as any node therein is not too tight on memory, but will escape to the larger layer, if need be. The intended use is to allow something like a batch manager to handle several jobs, each job in its own cpuset, but using common kernel memory for caches and such. Swapper and oom_kill activity is also constrained to Layer (2). A task in or below one mem_exclusive cpuset should not cause swapping on nodes in another non-overlapping mem_exclusive cpuset, nor provoke oom_killing of a task in another such cpuset. Heavy use of kernel memory for i/o caching and such by one job should not impact the memory available to jobs in other non-overlapping mem_exclusive cpusets. This patch enables providing hardwall, inescapable cpusets for memory allocations of each job, while sharing kernel memory allocations between several jobs, in an enclosing mem_exclusive cpuset. Like Dinakar's patch earlier to enable administering sched domains using the cpu_exclusive flag, this patch also provides a useful meaning to a cpuset flag that had previously done nothing much useful other than restrict what cpuset configurations were allowed. Signed-off-by: Paul Jackson <pj@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-09-07 06:18:12 +08:00
if (!cpuset_zone_allowed(zone, __GFP_HARDWALL))
continue;
zone->prev_priority = zone->temp_priority;
}
return ret;
}
/*
* For kswapd, balance_pgdat() will work across all this node's zones until
* they are all at pages_high.
*
* If `nr_pages' is non-zero then it is the number of pages which are to be
* reclaimed, regardless of the zone occupancies. This is a software suspend
* special.
*
* Returns the number of pages which were actually freed.
*
* There is special handling here for zones which are full of pinned pages.
* This can happen if the pages are all mlocked, or if they are all used by
* device drivers (say, ZONE_DMA). Or if they are all in use by hugetlb.
* What we do is to detect the case where all pages in the zone have been
* scanned twice and there has been zero successful reclaim. Mark the zone as
* dead and from now on, only perform a short scan. Basically we're polling
* the zone for when the problem goes away.
*
* kswapd scans the zones in the highmem->normal->dma direction. It skips
* zones which have free_pages > pages_high, but once a zone is found to have
* free_pages <= pages_high, we scan that zone and the lower zones regardless
* of the number of free pages in the lower zones. This interoperates with
* the page allocator fallback scheme to ensure that aging of pages is balanced
* across the zones.
*/
static int balance_pgdat(pg_data_t *pgdat, int nr_pages, int order)
{
int to_free = nr_pages;
int all_zones_ok;
int priority;
int i;
int total_scanned, total_reclaimed;
struct reclaim_state *reclaim_state = current->reclaim_state;
struct scan_control sc;
loop_again:
total_scanned = 0;
total_reclaimed = 0;
sc.gfp_mask = GFP_KERNEL;
sc.may_writepage = 0;
[PATCH] VM: add may_swap flag to scan_control Here's the next round of these patches. These are totally different in an attempt to meet the "simpler" request after the last patches. For reference the earlier threads are: http://marc.theaimsgroup.com/?l=linux-kernel&m=110839604924587&w=2 http://marc.theaimsgroup.com/?l=linux-mm&m=111461480721249&w=2 This set of patches replaces my other vm- patches that are currently in -mm. So they're against 2.6.12-rc5-mm1 about half way through the -mm patchset. As I said already this patch is a lot simpler. The reclaim is turned on or off on a per-zone basis using a syscall. I haven't tested the x86 syscall, so it might be wrong. It uses the existing reclaim/pageout code with the small addition of a may_swap flag to scan_control (patch 1/4). I also added __GFP_NORECLAIM (patch 3/4) so that certain allocation types can be flagged to never cause reclaim. This was a deficiency that was in all of my earlier patch sets. Previously, doing a big buffered read would fill one zone with page cache and then start to reclaim from that same zone, leaving the other zones untouched. Adding some extra throttling on the reclaim was also required (patch 4/4). Without the machine would grind to a crawl when doing a "make -j" kernel build. Even with this patch the System Time is higher on average, but it seems tolerable. Here are some numbers for kernbench runs on a 2-node, 4cpu, 8Gig RAM Altix in the "make -j" run: wall user sys %cpu ctx sw. sleeps ---- ---- --- ---- ------ ------ No patch 1009 1384 847 258 298170 504402 w/patch, no reclaim 880 1376 667 288 254064 396745 w/patch & reclaim 1079 1385 926 252 291625 548873 These numbers are the average of 2 runs of 3 "make -j" runs done right after system boot. Run-to-run variability for "make -j" is huge, so these numbers aren't terribly useful except to seee that with reclaim the benchmark still finishes in a reasonable amount of time. I also looked at the NUMA hit/miss stats for the "make -j" runs and the reclaim doesn't make any difference when the machine is thrashing away. Doing a "make -j8" on a single node that is filled with page cache pages takes 700 seconds with reclaim turned on and 735 seconds without reclaim (due to remote memory accesses). The simple zone_reclaim syscall program is at http://www.bork.org/~mort/sgi/zone_reclaim.c This patch: This adds an extra switch to the scan_control struct. It simply lets the reclaim code know if its allowed to swap pages out. This was required for a simple per-zone reclaimer. Without this addition pages would be swapped out as soon as a zone ran out of memory and the early reclaim kicked in. Signed-off-by: Martin Hicks <mort@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-22 08:14:40 +08:00
sc.may_swap = 1;
sc.nr_mapped = read_page_state(nr_mapped);
inc_page_state(pageoutrun);
for (i = 0; i < pgdat->nr_zones; i++) {
struct zone *zone = pgdat->node_zones + i;
zone->temp_priority = DEF_PRIORITY;
}
for (priority = DEF_PRIORITY; priority >= 0; priority--) {
int end_zone = 0; /* Inclusive. 0 = ZONE_DMA */
unsigned long lru_pages = 0;
/* The swap token gets in the way of swapout... */
if (!priority)
disable_swap_token();
all_zones_ok = 1;
if (nr_pages == 0) {
/*
* Scan in the highmem->dma direction for the highest
* zone which needs scanning
*/
for (i = pgdat->nr_zones - 1; i >= 0; i--) {
struct zone *zone = pgdat->node_zones + i;
if (zone->present_pages == 0)
continue;
if (zone->all_unreclaimable &&
priority != DEF_PRIORITY)
continue;
if (!zone_watermark_ok(zone, order,
zone->pages_high, 0, 0)) {
end_zone = i;
goto scan;
}
}
goto out;
} else {
end_zone = pgdat->nr_zones - 1;
}
scan:
for (i = 0; i <= end_zone; i++) {
struct zone *zone = pgdat->node_zones + i;
lru_pages += zone->nr_active + zone->nr_inactive;
}
/*
* Now scan the zone in the dma->highmem direction, stopping
* at the last zone which needs scanning.
*
* We do this because the page allocator works in the opposite
* direction. This prevents the page allocator from allocating
* pages behind kswapd's direction of progress, which would
* cause too much scanning of the lower zones.
*/
for (i = 0; i <= end_zone; i++) {
struct zone *zone = pgdat->node_zones + i;
int nr_slab;
if (zone->present_pages == 0)
continue;
if (zone->all_unreclaimable && priority != DEF_PRIORITY)
continue;
if (nr_pages == 0) { /* Not software suspend */
if (!zone_watermark_ok(zone, order,
zone->pages_high, end_zone, 0))
all_zones_ok = 0;
}
zone->temp_priority = priority;
if (zone->prev_priority > priority)
zone->prev_priority = priority;
sc.nr_scanned = 0;
sc.nr_reclaimed = 0;
sc.priority = priority;
sc.swap_cluster_max = nr_pages? nr_pages : SWAP_CLUSTER_MAX;
atomic_inc(&zone->reclaim_in_progress);
shrink_zone(zone, &sc);
atomic_dec(&zone->reclaim_in_progress);
reclaim_state->reclaimed_slab = 0;
nr_slab = shrink_slab(sc.nr_scanned, GFP_KERNEL,
lru_pages);
sc.nr_reclaimed += reclaim_state->reclaimed_slab;
total_reclaimed += sc.nr_reclaimed;
total_scanned += sc.nr_scanned;
if (zone->all_unreclaimable)
continue;
if (nr_slab == 0 && zone->pages_scanned >=
(zone->nr_active + zone->nr_inactive) * 4)
zone->all_unreclaimable = 1;
/*
* If we've done a decent amount of scanning and
* the reclaim ratio is low, start doing writepage
* even in laptop mode
*/
if (total_scanned > SWAP_CLUSTER_MAX * 2 &&
total_scanned > total_reclaimed+total_reclaimed/2)
sc.may_writepage = 1;
}
if (nr_pages && to_free > total_reclaimed)
continue; /* swsusp: need to do more work */
if (all_zones_ok)
break; /* kswapd: all done */
/*
* OK, kswapd is getting into trouble. Take a nap, then take
* another pass across the zones.
*/
if (total_scanned && priority < DEF_PRIORITY - 2)
blk_congestion_wait(WRITE, HZ/10);
/*
* We do this so kswapd doesn't build up large priorities for
* example when it is freeing in parallel with allocators. It
* matches the direct reclaim path behaviour in terms of impact
* on zone->*_priority.
*/
if ((total_reclaimed >= SWAP_CLUSTER_MAX) && (!nr_pages))
break;
}
out:
for (i = 0; i < pgdat->nr_zones; i++) {
struct zone *zone = pgdat->node_zones + i;
zone->prev_priority = zone->temp_priority;
}
if (!all_zones_ok) {
cond_resched();
goto loop_again;
}
return total_reclaimed;
}
/*
* The background pageout daemon, started as a kernel thread
* from the init process.
*
* This basically trickles out pages so that we have _some_
* free memory available even if there is no other activity
* that frees anything up. This is needed for things like routing
* etc, where we otherwise might have all activity going on in
* asynchronous contexts that cannot page things out.
*
* If there are applications that are active memory-allocators
* (most normal use), this basically shouldn't matter.
*/
static int kswapd(void *p)
{
unsigned long order;
pg_data_t *pgdat = (pg_data_t*)p;
struct task_struct *tsk = current;
DEFINE_WAIT(wait);
struct reclaim_state reclaim_state = {
.reclaimed_slab = 0,
};
cpumask_t cpumask;
daemonize("kswapd%d", pgdat->node_id);
cpumask = node_to_cpumask(pgdat->node_id);
if (!cpus_empty(cpumask))
set_cpus_allowed(tsk, cpumask);
current->reclaim_state = &reclaim_state;
/*
* Tell the memory management that we're a "memory allocator",
* and that if we need more memory we should get access to it
* regardless (see "__alloc_pages()"). "kswapd" should
* never get caught in the normal page freeing logic.
*
* (Kswapd normally doesn't need memory anyway, but sometimes
* you need a small amount of memory in order to be able to
* page out something else, and this flag essentially protects
* us from recursively trying to free more memory as we're
* trying to free the first piece of memory in the first place).
*/
tsk->flags |= PF_MEMALLOC|PF_KSWAPD;
order = 0;
for ( ; ; ) {
unsigned long new_order;
try_to_freeze();
prepare_to_wait(&pgdat->kswapd_wait, &wait, TASK_INTERRUPTIBLE);
new_order = pgdat->kswapd_max_order;
pgdat->kswapd_max_order = 0;
if (order < new_order) {
/*
* Don't sleep if someone wants a larger 'order'
* allocation
*/
order = new_order;
} else {
schedule();
order = pgdat->kswapd_max_order;
}
finish_wait(&pgdat->kswapd_wait, &wait);
balance_pgdat(pgdat, 0, order);
}
return 0;
}
/*
* A zone is low on free memory, so wake its kswapd task to service it.
*/
void wakeup_kswapd(struct zone *zone, int order)
{
pg_data_t *pgdat;
if (zone->present_pages == 0)
return;
pgdat = zone->zone_pgdat;
if (zone_watermark_ok(zone, order, zone->pages_low, 0, 0))
return;
if (pgdat->kswapd_max_order < order)
pgdat->kswapd_max_order = order;
[PATCH] cpusets: formalize intermediate GFP_KERNEL containment This patch makes use of the previously underutilized cpuset flag 'mem_exclusive' to provide what amounts to another layer of memory placement resolution. With this patch, there are now the following four layers of memory placement available: 1) The whole system (interrupt and GFP_ATOMIC allocations can use this), 2) The nearest enclosing mem_exclusive cpuset (GFP_KERNEL allocations can use), 3) The current tasks cpuset (GFP_USER allocations constrained to here), and 4) Specific node placement, using mbind and set_mempolicy. These nest - each layer is a subset (same or within) of the previous. Layer (2) above is new, with this patch. The call used to check whether a zone (its node, actually) is in a cpuset (in its mems_allowed, actually) is extended to take a gfp_mask argument, and its logic is extended, in the case that __GFP_HARDWALL is not set in the flag bits, to look up the cpuset hierarchy for the nearest enclosing mem_exclusive cpuset, to determine if placement is allowed. The definition of GFP_USER, which used to be identical to GFP_KERNEL, is changed to also set the __GFP_HARDWALL bit, in the previous cpuset_gfp_hardwall_flag patch. GFP_ATOMIC and GFP_KERNEL allocations will stay within the current tasks cpuset, so long as any node therein is not too tight on memory, but will escape to the larger layer, if need be. The intended use is to allow something like a batch manager to handle several jobs, each job in its own cpuset, but using common kernel memory for caches and such. Swapper and oom_kill activity is also constrained to Layer (2). A task in or below one mem_exclusive cpuset should not cause swapping on nodes in another non-overlapping mem_exclusive cpuset, nor provoke oom_killing of a task in another such cpuset. Heavy use of kernel memory for i/o caching and such by one job should not impact the memory available to jobs in other non-overlapping mem_exclusive cpusets. This patch enables providing hardwall, inescapable cpusets for memory allocations of each job, while sharing kernel memory allocations between several jobs, in an enclosing mem_exclusive cpuset. Like Dinakar's patch earlier to enable administering sched domains using the cpu_exclusive flag, this patch also provides a useful meaning to a cpuset flag that had previously done nothing much useful other than restrict what cpuset configurations were allowed. Signed-off-by: Paul Jackson <pj@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-09-07 06:18:12 +08:00
if (!cpuset_zone_allowed(zone, __GFP_HARDWALL))
return;
if (!waitqueue_active(&pgdat->kswapd_wait))
return;
wake_up_interruptible(&pgdat->kswapd_wait);
}
#ifdef CONFIG_PM
/*
* Try to free `nr_pages' of memory, system-wide. Returns the number of freed
* pages.
*/
int shrink_all_memory(int nr_pages)
{
pg_data_t *pgdat;
int nr_to_free = nr_pages;
int ret = 0;
struct reclaim_state reclaim_state = {
.reclaimed_slab = 0,
};
current->reclaim_state = &reclaim_state;
for_each_pgdat(pgdat) {
int freed;
freed = balance_pgdat(pgdat, nr_to_free, 0);
ret += freed;
nr_to_free -= freed;
if (nr_to_free <= 0)
break;
}
current->reclaim_state = NULL;
return ret;
}
#endif
#ifdef CONFIG_HOTPLUG_CPU
/* It's optimal to keep kswapds on the same CPUs as their memory, but
not required for correctness. So if the last cpu in a node goes
away, we get changed to run anywhere: as the first one comes back,
restore their cpu bindings. */
static int __devinit cpu_callback(struct notifier_block *nfb,
unsigned long action,
void *hcpu)
{
pg_data_t *pgdat;
cpumask_t mask;
if (action == CPU_ONLINE) {
for_each_pgdat(pgdat) {
mask = node_to_cpumask(pgdat->node_id);
if (any_online_cpu(mask) != NR_CPUS)
/* One of our CPUs online: restore mask */
set_cpus_allowed(pgdat->kswapd, mask);
}
}
return NOTIFY_OK;
}
#endif /* CONFIG_HOTPLUG_CPU */
static int __init kswapd_init(void)
{
pg_data_t *pgdat;
swap_setup();
for_each_pgdat(pgdat)
pgdat->kswapd
= find_task_by_pid(kernel_thread(kswapd, pgdat, CLONE_KERNEL));
total_memory = nr_free_pagecache_pages();
hotcpu_notifier(cpu_callback, 0);
return 0;
}
module_init(kswapd_init)
[PATCH] VM: early zone reclaim This is the core of the (much simplified) early reclaim. The goal of this patch is to reclaim some easily-freed pages from a zone before falling back onto another zone. One of the major uses of this is NUMA machines. With the default allocator behavior the allocator would look for memory in another zone, which might be off-node, before trying to reclaim from the current zone. This adds a zone tuneable to enable early zone reclaim. It is selected on a per-zone basis and is turned on/off via syscall. Adding some extra throttling on the reclaim was also required (patch 4/4). Without the machine would grind to a crawl when doing a "make -j" kernel build. Even with this patch the System Time is higher on average, but it seems tolerable. Here are some numbers for kernbench runs on a 2-node, 4cpu, 8Gig RAM Altix in the "make -j" run: wall user sys %cpu ctx sw. sleeps ---- ---- --- ---- ------ ------ No patch 1009 1384 847 258 298170 504402 w/patch, no reclaim 880 1376 667 288 254064 396745 w/patch & reclaim 1079 1385 926 252 291625 548873 These numbers are the average of 2 runs of 3 "make -j" runs done right after system boot. Run-to-run variability for "make -j" is huge, so these numbers aren't terribly useful except to seee that with reclaim the benchmark still finishes in a reasonable amount of time. I also looked at the NUMA hit/miss stats for the "make -j" runs and the reclaim doesn't make any difference when the machine is thrashing away. Doing a "make -j8" on a single node that is filled with page cache pages takes 700 seconds with reclaim turned on and 735 seconds without reclaim (due to remote memory accesses). The simple zone_reclaim syscall program is at http://www.bork.org/~mort/sgi/zone_reclaim.c Signed-off-by: Martin Hicks <mort@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-22 08:14:41 +08:00
/*
* Try to free up some pages from this zone through reclaim.
*/
int zone_reclaim(struct zone *zone, gfp_t gfp_mask, unsigned int order)
[PATCH] VM: early zone reclaim This is the core of the (much simplified) early reclaim. The goal of this patch is to reclaim some easily-freed pages from a zone before falling back onto another zone. One of the major uses of this is NUMA machines. With the default allocator behavior the allocator would look for memory in another zone, which might be off-node, before trying to reclaim from the current zone. This adds a zone tuneable to enable early zone reclaim. It is selected on a per-zone basis and is turned on/off via syscall. Adding some extra throttling on the reclaim was also required (patch 4/4). Without the machine would grind to a crawl when doing a "make -j" kernel build. Even with this patch the System Time is higher on average, but it seems tolerable. Here are some numbers for kernbench runs on a 2-node, 4cpu, 8Gig RAM Altix in the "make -j" run: wall user sys %cpu ctx sw. sleeps ---- ---- --- ---- ------ ------ No patch 1009 1384 847 258 298170 504402 w/patch, no reclaim 880 1376 667 288 254064 396745 w/patch & reclaim 1079 1385 926 252 291625 548873 These numbers are the average of 2 runs of 3 "make -j" runs done right after system boot. Run-to-run variability for "make -j" is huge, so these numbers aren't terribly useful except to seee that with reclaim the benchmark still finishes in a reasonable amount of time. I also looked at the NUMA hit/miss stats for the "make -j" runs and the reclaim doesn't make any difference when the machine is thrashing away. Doing a "make -j8" on a single node that is filled with page cache pages takes 700 seconds with reclaim turned on and 735 seconds without reclaim (due to remote memory accesses). The simple zone_reclaim syscall program is at http://www.bork.org/~mort/sgi/zone_reclaim.c Signed-off-by: Martin Hicks <mort@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-22 08:14:41 +08:00
{
struct scan_control sc;
int nr_pages = 1 << order;
int total_reclaimed = 0;
/* The reclaim may sleep, so don't do it if sleep isn't allowed */
if (!(gfp_mask & __GFP_WAIT))
return 0;
if (zone->all_unreclaimable)
return 0;
sc.gfp_mask = gfp_mask;
sc.may_writepage = 0;
sc.may_swap = 0;
sc.nr_mapped = read_page_state(nr_mapped);
sc.nr_scanned = 0;
sc.nr_reclaimed = 0;
/* scan at the highest priority */
sc.priority = 0;
disable_swap_token();
[PATCH] VM: early zone reclaim This is the core of the (much simplified) early reclaim. The goal of this patch is to reclaim some easily-freed pages from a zone before falling back onto another zone. One of the major uses of this is NUMA machines. With the default allocator behavior the allocator would look for memory in another zone, which might be off-node, before trying to reclaim from the current zone. This adds a zone tuneable to enable early zone reclaim. It is selected on a per-zone basis and is turned on/off via syscall. Adding some extra throttling on the reclaim was also required (patch 4/4). Without the machine would grind to a crawl when doing a "make -j" kernel build. Even with this patch the System Time is higher on average, but it seems tolerable. Here are some numbers for kernbench runs on a 2-node, 4cpu, 8Gig RAM Altix in the "make -j" run: wall user sys %cpu ctx sw. sleeps ---- ---- --- ---- ------ ------ No patch 1009 1384 847 258 298170 504402 w/patch, no reclaim 880 1376 667 288 254064 396745 w/patch & reclaim 1079 1385 926 252 291625 548873 These numbers are the average of 2 runs of 3 "make -j" runs done right after system boot. Run-to-run variability for "make -j" is huge, so these numbers aren't terribly useful except to seee that with reclaim the benchmark still finishes in a reasonable amount of time. I also looked at the NUMA hit/miss stats for the "make -j" runs and the reclaim doesn't make any difference when the machine is thrashing away. Doing a "make -j8" on a single node that is filled with page cache pages takes 700 seconds with reclaim turned on and 735 seconds without reclaim (due to remote memory accesses). The simple zone_reclaim syscall program is at http://www.bork.org/~mort/sgi/zone_reclaim.c Signed-off-by: Martin Hicks <mort@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-22 08:14:41 +08:00
if (nr_pages > SWAP_CLUSTER_MAX)
sc.swap_cluster_max = nr_pages;
else
sc.swap_cluster_max = SWAP_CLUSTER_MAX;
/* Don't reclaim the zone if there are other reclaimers active */
if (atomic_read(&zone->reclaim_in_progress) > 0)
goto out;
[PATCH] VM: early zone reclaim This is the core of the (much simplified) early reclaim. The goal of this patch is to reclaim some easily-freed pages from a zone before falling back onto another zone. One of the major uses of this is NUMA machines. With the default allocator behavior the allocator would look for memory in another zone, which might be off-node, before trying to reclaim from the current zone. This adds a zone tuneable to enable early zone reclaim. It is selected on a per-zone basis and is turned on/off via syscall. Adding some extra throttling on the reclaim was also required (patch 4/4). Without the machine would grind to a crawl when doing a "make -j" kernel build. Even with this patch the System Time is higher on average, but it seems tolerable. Here are some numbers for kernbench runs on a 2-node, 4cpu, 8Gig RAM Altix in the "make -j" run: wall user sys %cpu ctx sw. sleeps ---- ---- --- ---- ------ ------ No patch 1009 1384 847 258 298170 504402 w/patch, no reclaim 880 1376 667 288 254064 396745 w/patch & reclaim 1079 1385 926 252 291625 548873 These numbers are the average of 2 runs of 3 "make -j" runs done right after system boot. Run-to-run variability for "make -j" is huge, so these numbers aren't terribly useful except to seee that with reclaim the benchmark still finishes in a reasonable amount of time. I also looked at the NUMA hit/miss stats for the "make -j" runs and the reclaim doesn't make any difference when the machine is thrashing away. Doing a "make -j8" on a single node that is filled with page cache pages takes 700 seconds with reclaim turned on and 735 seconds without reclaim (due to remote memory accesses). The simple zone_reclaim syscall program is at http://www.bork.org/~mort/sgi/zone_reclaim.c Signed-off-by: Martin Hicks <mort@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-22 08:14:41 +08:00
shrink_zone(zone, &sc);
total_reclaimed = sc.nr_reclaimed;
out:
[PATCH] VM: early zone reclaim This is the core of the (much simplified) early reclaim. The goal of this patch is to reclaim some easily-freed pages from a zone before falling back onto another zone. One of the major uses of this is NUMA machines. With the default allocator behavior the allocator would look for memory in another zone, which might be off-node, before trying to reclaim from the current zone. This adds a zone tuneable to enable early zone reclaim. It is selected on a per-zone basis and is turned on/off via syscall. Adding some extra throttling on the reclaim was also required (patch 4/4). Without the machine would grind to a crawl when doing a "make -j" kernel build. Even with this patch the System Time is higher on average, but it seems tolerable. Here are some numbers for kernbench runs on a 2-node, 4cpu, 8Gig RAM Altix in the "make -j" run: wall user sys %cpu ctx sw. sleeps ---- ---- --- ---- ------ ------ No patch 1009 1384 847 258 298170 504402 w/patch, no reclaim 880 1376 667 288 254064 396745 w/patch & reclaim 1079 1385 926 252 291625 548873 These numbers are the average of 2 runs of 3 "make -j" runs done right after system boot. Run-to-run variability for "make -j" is huge, so these numbers aren't terribly useful except to seee that with reclaim the benchmark still finishes in a reasonable amount of time. I also looked at the NUMA hit/miss stats for the "make -j" runs and the reclaim doesn't make any difference when the machine is thrashing away. Doing a "make -j8" on a single node that is filled with page cache pages takes 700 seconds with reclaim turned on and 735 seconds without reclaim (due to remote memory accesses). The simple zone_reclaim syscall program is at http://www.bork.org/~mort/sgi/zone_reclaim.c Signed-off-by: Martin Hicks <mort@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-22 08:14:41 +08:00
return total_reclaimed;
}
asmlinkage long sys_set_zone_reclaim(unsigned int node, unsigned int zone,
unsigned int state)
{
struct zone *z;
int i;
if (!capable(CAP_SYS_ADMIN))
return -EACCES;
[PATCH] VM: early zone reclaim This is the core of the (much simplified) early reclaim. The goal of this patch is to reclaim some easily-freed pages from a zone before falling back onto another zone. One of the major uses of this is NUMA machines. With the default allocator behavior the allocator would look for memory in another zone, which might be off-node, before trying to reclaim from the current zone. This adds a zone tuneable to enable early zone reclaim. It is selected on a per-zone basis and is turned on/off via syscall. Adding some extra throttling on the reclaim was also required (patch 4/4). Without the machine would grind to a crawl when doing a "make -j" kernel build. Even with this patch the System Time is higher on average, but it seems tolerable. Here are some numbers for kernbench runs on a 2-node, 4cpu, 8Gig RAM Altix in the "make -j" run: wall user sys %cpu ctx sw. sleeps ---- ---- --- ---- ------ ------ No patch 1009 1384 847 258 298170 504402 w/patch, no reclaim 880 1376 667 288 254064 396745 w/patch & reclaim 1079 1385 926 252 291625 548873 These numbers are the average of 2 runs of 3 "make -j" runs done right after system boot. Run-to-run variability for "make -j" is huge, so these numbers aren't terribly useful except to seee that with reclaim the benchmark still finishes in a reasonable amount of time. I also looked at the NUMA hit/miss stats for the "make -j" runs and the reclaim doesn't make any difference when the machine is thrashing away. Doing a "make -j8" on a single node that is filled with page cache pages takes 700 seconds with reclaim turned on and 735 seconds without reclaim (due to remote memory accesses). The simple zone_reclaim syscall program is at http://www.bork.org/~mort/sgi/zone_reclaim.c Signed-off-by: Martin Hicks <mort@sgi.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-22 08:14:41 +08:00
if (node >= MAX_NUMNODES || !node_online(node))
return -EINVAL;
/* This will break if we ever add more zones */
if (!(zone & (1<<ZONE_DMA|1<<ZONE_NORMAL|1<<ZONE_HIGHMEM)))
return -EINVAL;
for (i = 0; i < MAX_NR_ZONES; i++) {
if (!(zone & 1<<i))
continue;
z = &NODE_DATA(node)->node_zones[i];
if (state)
z->reclaim_pages = 1;
else
z->reclaim_pages = 0;
}
return 0;
}