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2011-10-31vmscan: count pages into balanced for zone with good watermarkShaohua Li
It's possible a zone watermark is ok when entering the balance_pgdat() loop, while the zone is within the requested classzone_idx. Count pages from this zone into `balanced'. In this way, we can skip shrinking zones too much for high order allocation. Signed-off-by: Shaohua Li <shaohua.li@intel.com> Acked-by: Mel Gorman <mgorman@suse.de> Reviewed-by: Minchan Kim <minchan.kim@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: vmscan: immediately reclaim end-of-LRU dirty pages when writeback completesMel Gorman
When direct reclaim encounters a dirty page, it gets recycled around the LRU for another cycle. This patch marks the page PageReclaim similar to deactivate_page() so that the page gets reclaimed almost immediately after the page gets cleaned. This is to avoid reclaiming clean pages that are younger than a dirty page encountered at the end of the LRU that might have been something like a use-once page. Signed-off-by: Mel Gorman <mgorman@suse.de> Acked-by: Johannes Weiner <jweiner@redhat.com> Cc: Dave Chinner <david@fromorbit.com> Cc: Christoph Hellwig <hch@infradead.org> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Jan Kara <jack@suse.cz> Cc: Minchan Kim <minchan.kim@gmail.com> Cc: Rik van Riel <riel@redhat.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Alex Elder <aelder@sgi.com> Cc: Theodore Ts'o <tytso@mit.edu> Cc: Chris Mason <chris.mason@oracle.com> Cc: Dave Hansen <dave@linux.vnet.ibm.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: vmscan: throttle reclaim if encountering too many dirty pages under ↵Mel Gorman
writeback Workloads that are allocating frequently and writing files place a large number of dirty pages on the LRU. With use-once logic, it is possible for them to reach the end of the LRU quickly requiring the reclaimer to scan more to find clean pages. Ordinarily, processes that are dirtying memory will get throttled by dirty balancing but this is a global heuristic and does not take into account that LRUs are maintained on a per-zone basis. This can lead to a situation whereby reclaim is scanning heavily, skipping over a large number of pages under writeback and recycling them around the LRU consuming CPU. This patch checks how many of the number of pages isolated from the LRU were dirty and under writeback. If a percentage of them under writeback, the process will be throttled if a backing device or the zone is congested. Note that this applies whether it is anonymous or file-backed pages that are under writeback meaning that swapping is potentially throttled. This is intentional due to the fact if the swap device is congested, scanning more pages and dispatching more IO is not going to help matters. The percentage that must be in writeback depends on the priority. At default priority, all of them must be dirty. At DEF_PRIORITY-1, 50% of them must be, DEF_PRIORITY-2, 25% etc. i.e. as pressure increases the greater the likelihood the process will get throttled to allow the flusher threads to make some progress. Signed-off-by: Mel Gorman <mgorman@suse.de> Reviewed-by: Minchan Kim <minchan.kim@gmail.com> Acked-by: Johannes Weiner <jweiner@redhat.com> Cc: Dave Chinner <david@fromorbit.com> Cc: Christoph Hellwig <hch@infradead.org> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Jan Kara <jack@suse.cz> Cc: Rik van Riel <riel@redhat.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Alex Elder <aelder@sgi.com> Cc: Theodore Ts'o <tytso@mit.edu> Cc: Chris Mason <chris.mason@oracle.com> Cc: Dave Hansen <dave@linux.vnet.ibm.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: vmscan: do not writeback filesystem pages in kswapd except in high priorityMel Gorman
It is preferable that no dirty pages are dispatched for cleaning from the page reclaim path. At normal priorities, this patch prevents kswapd writing pages. However, page reclaim does have a requirement that pages be freed in a particular zone. If it is failing to make sufficient progress (reclaiming < SWAP_CLUSTER_MAX at any priority priority), the priority is raised to scan more pages. A priority of DEF_PRIORITY - 3 is considered to be the point where kswapd is getting into trouble reclaiming pages. If this priority is reached, kswapd will dispatch pages for writing. Signed-off-by: Mel Gorman <mgorman@suse.de> Reviewed-by: Minchan Kim <minchan.kim@gmail.com> Cc: Dave Chinner <david@fromorbit.com> Cc: Christoph Hellwig <hch@infradead.org> Cc: Johannes Weiner <jweiner@redhat.com> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Jan Kara <jack@suse.cz> Cc: Rik van Riel <riel@redhat.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Alex Elder <aelder@sgi.com> Cc: Theodore Ts'o <tytso@mit.edu> Cc: Chris Mason <chris.mason@oracle.com> Cc: Dave Hansen <dave@linux.vnet.ibm.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: vmscan: remove dead code related to lumpy reclaim waiting on pages under ↵Mel Gorman
writeback Lumpy reclaim worked with two passes - the first which queued pages for IO and the second which waited on writeback. As direct reclaim can no longer write pages there is some dead code. This patch removes it but direct reclaim will continue to wait on pages under writeback while in synchronous reclaim mode. Signed-off-by: Mel Gorman <mgorman@suse.de> Cc: Dave Chinner <david@fromorbit.com> Cc: Christoph Hellwig <hch@infradead.org> Cc: Johannes Weiner <jweiner@redhat.com> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Jan Kara <jack@suse.cz> Cc: Minchan Kim <minchan.kim@gmail.com> Cc: Rik van Riel <riel@redhat.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Alex Elder <aelder@sgi.com> Cc: Theodore Ts'o <tytso@mit.edu> Cc: Chris Mason <chris.mason@oracle.com> Cc: Dave Hansen <dave@linux.vnet.ibm.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: vmscan: do not writeback filesystem pages in direct reclaimMel Gorman
Testing from the XFS folk revealed that there is still too much I/O from the end of the LRU in kswapd. Previously it was considered acceptable by VM people for a small number of pages to be written back from reclaim with testing generally showing about 0.3% of pages reclaimed were written back (higher if memory was low). That writing back a small number of pages is ok has been heavily disputed for quite some time and Dave Chinner explained it well; It doesn't have to be a very high number to be a problem. IO is orders of magnitude slower than the CPU time it takes to flush a page, so the cost of making a bad flush decision is very high. And single page writeback from the LRU is almost always a bad flush decision. To complicate matters, filesystems respond very differently to requests from reclaim according to Christoph Hellwig; xfs tries to write it back if the requester is kswapd ext4 ignores the request if it's a delayed allocation btrfs ignores the request As a result, each filesystem has different performance characteristics when under memory pressure and there are many pages being dirtied. In some cases, the request is ignored entirely so the VM cannot depend on the IO being dispatched. The objective of this series is to reduce writing of filesystem-backed pages from reclaim, play nicely with writeback that is already in progress and throttle reclaim appropriately when writeback pages are encountered. The assumption is that the flushers will always write pages faster than if reclaim issues the IO. A secondary goal is to avoid the problem whereby direct reclaim splices two potentially deep call stacks together. There is a potential new problem as reclaim has less control over how long before a page in a particularly zone or container is cleaned and direct reclaimers depend on kswapd or flusher threads to do the necessary work. However, as filesystems sometimes ignore direct reclaim requests already, it is not expected to be a serious issue. Patch 1 disables writeback of filesystem pages from direct reclaim entirely. Anonymous pages are still written. Patch 2 removes dead code in lumpy reclaim as it is no longer able to synchronously write pages. This hurts lumpy reclaim but there is an expectation that compaction is used for hugepage allocations these days and lumpy reclaim's days are numbered. Patches 3-4 add warnings to XFS and ext4 if called from direct reclaim. With patch 1, this "never happens" and is intended to catch regressions in this logic in the future. Patch 5 disables writeback of filesystem pages from kswapd unless the priority is raised to the point where kswapd is considered to be in trouble. Patch 6 throttles reclaimers if too many dirty pages are being encountered and the zones or backing devices are congested. Patch 7 invalidates dirty pages found at the end of the LRU so they are reclaimed quickly after being written back rather than waiting for a reclaimer to find them I consider this series to be orthogonal to the writeback work but it is worth noting that the writeback work affects the viability of patch 8 in particular. I tested this on ext4 and xfs using fs_mark, a simple writeback test based on dd and a micro benchmark that does a streaming write to a large mapping (exercises use-once LRU logic) followed by streaming writes to a mix of anonymous and file-backed mappings. The command line for fs_mark when botted with 512M looked something like ./fs_mark -d /tmp/fsmark-2676 -D 100 -N 150 -n 150 -L 25 -t 1 -S0 -s 10485760 The number of files was adjusted depending on the amount of available memory so that the files created was about 3xRAM. For multiple threads, the -d switch is specified multiple times. The test machine is x86-64 with an older generation of AMD processor with 4 cores. The underlying storage was 4 disks configured as RAID-0 as this was the best configuration of storage I had available. Swap is on a separate disk. Dirty ratio was tuned to 40% instead of the default of 20%. Testing was run with and without monitors to both verify that the patches were operating as expected and that any performance gain was real and not due to interference from monitors. Here is a summary of results based on testing XFS. 512M1P-xfs Files/s mean 32.69 ( 0.00%) 34.44 ( 5.08%) 512M1P-xfs Elapsed Time fsmark 51.41 48.29 512M1P-xfs Elapsed Time simple-wb 114.09 108.61 512M1P-xfs Elapsed Time mmap-strm 113.46 109.34 512M1P-xfs Kswapd efficiency fsmark 62% 63% 512M1P-xfs Kswapd efficiency simple-wb 56% 61% 512M1P-xfs Kswapd efficiency mmap-strm 44% 42% 512M-xfs Files/s mean 30.78 ( 0.00%) 35.94 (14.36%) 512M-xfs Elapsed Time fsmark 56.08 48.90 512M-xfs Elapsed Time simple-wb 112.22 98.13 512M-xfs Elapsed Time mmap-strm 219.15 196.67 512M-xfs Kswapd efficiency fsmark 54% 56% 512M-xfs Kswapd efficiency simple-wb 54% 55% 512M-xfs Kswapd efficiency mmap-strm 45% 44% 512M-4X-xfs Files/s mean 30.31 ( 0.00%) 33.33 ( 9.06%) 512M-4X-xfs Elapsed Time fsmark 63.26 55.88 512M-4X-xfs Elapsed Time simple-wb 100.90 90.25 512M-4X-xfs Elapsed Time mmap-strm 261.73 255.38 512M-4X-xfs Kswapd efficiency fsmark 49% 50% 512M-4X-xfs Kswapd efficiency simple-wb 54% 56% 512M-4X-xfs Kswapd efficiency mmap-strm 37% 36% 512M-16X-xfs Files/s mean 60.89 ( 0.00%) 65.22 ( 6.64%) 512M-16X-xfs Elapsed Time fsmark 67.47 58.25 512M-16X-xfs Elapsed Time simple-wb 103.22 90.89 512M-16X-xfs Elapsed Time mmap-strm 237.09 198.82 512M-16X-xfs Kswapd efficiency fsmark 45% 46% 512M-16X-xfs Kswapd efficiency simple-wb 53% 55% 512M-16X-xfs Kswapd efficiency mmap-strm 33% 33% Up until 512-4X, the FSmark improvements were statistically significant. For the 4X and 16X tests the results were within standard deviations but just barely. The time to completion for all tests is improved which is an important result. In general, kswapd efficiency is not affected by skipping dirty pages. 1024M1P-xfs Files/s mean 39.09 ( 0.00%) 41.15 ( 5.01%) 1024M1P-xfs Elapsed Time fsmark 84.14 80.41 1024M1P-xfs Elapsed Time simple-wb 210.77 184.78 1024M1P-xfs Elapsed Time mmap-strm 162.00 160.34 1024M1P-xfs Kswapd efficiency fsmark 69% 75% 1024M1P-xfs Kswapd efficiency simple-wb 71% 77% 1024M1P-xfs Kswapd efficiency mmap-strm 43% 44% 1024M-xfs Files/s mean 35.45 ( 0.00%) 37.00 ( 4.19%) 1024M-xfs Elapsed Time fsmark 94.59 91.00 1024M-xfs Elapsed Time simple-wb 229.84 195.08 1024M-xfs Elapsed Time mmap-strm 405.38 440.29 1024M-xfs Kswapd efficiency fsmark 79% 71% 1024M-xfs Kswapd efficiency simple-wb 74% 74% 1024M-xfs Kswapd efficiency mmap-strm 39% 42% 1024M-4X-xfs Files/s mean 32.63 ( 0.00%) 35.05 ( 6.90%) 1024M-4X-xfs Elapsed Time fsmark 103.33 97.74 1024M-4X-xfs Elapsed Time simple-wb 204.48 178.57 1024M-4X-xfs Elapsed Time mmap-strm 528.38 511.88 1024M-4X-xfs Kswapd efficiency fsmark 81% 70% 1024M-4X-xfs Kswapd efficiency simple-wb 73% 72% 1024M-4X-xfs Kswapd efficiency mmap-strm 39% 38% 1024M-16X-xfs Files/s mean 42.65 ( 0.00%) 42.97 ( 0.74%) 1024M-16X-xfs Elapsed Time fsmark 103.11 99.11 1024M-16X-xfs Elapsed Time simple-wb 200.83 178.24 1024M-16X-xfs Elapsed Time mmap-strm 397.35 459.82 1024M-16X-xfs Kswapd efficiency fsmark 84% 69% 1024M-16X-xfs Kswapd efficiency simple-wb 74% 73% 1024M-16X-xfs Kswapd efficiency mmap-strm 39% 40% All FSMark tests up to 16X had statistically significant improvements. For the most part, tests are completing faster with the exception of the streaming writes to a mixture of anonymous and file-backed mappings which were slower in two cases In the cases where the mmap-strm tests were slower, there was more swapping due to dirty pages being skipped. The number of additional pages swapped is almost identical to the fewer number of pages written from reclaim. In other words, roughly the same number of pages were reclaimed but swapping was slower. As the test is a bit unrealistic and stresses memory heavily, the small shift is acceptable. 4608M1P-xfs Files/s mean 29.75 ( 0.00%) 30.96 ( 3.91%) 4608M1P-xfs Elapsed Time fsmark 512.01 492.15 4608M1P-xfs Elapsed Time simple-wb 618.18 566.24 4608M1P-xfs Elapsed Time mmap-strm 488.05 465.07 4608M1P-xfs Kswapd efficiency fsmark 93% 86% 4608M1P-xfs Kswapd efficiency simple-wb 88% 84% 4608M1P-xfs Kswapd efficiency mmap-strm 46% 45% 4608M-xfs Files/s mean 27.60 ( 0.00%) 28.85 ( 4.33%) 4608M-xfs Elapsed Time fsmark 555.96 532.34 4608M-xfs Elapsed Time simple-wb 659.72 571.85 4608M-xfs Elapsed Time mmap-strm 1082.57 1146.38 4608M-xfs Kswapd efficiency fsmark 89% 91% 4608M-xfs Kswapd efficiency simple-wb 88% 82% 4608M-xfs Kswapd efficiency mmap-strm 48% 46% 4608M-4X-xfs Files/s mean 26.00 ( 0.00%) 27.47 ( 5.35%) 4608M-4X-xfs Elapsed Time fsmark 592.91 564.00 4608M-4X-xfs Elapsed Time simple-wb 616.65 575.07 4608M-4X-xfs Elapsed Time mmap-strm 1773.02 1631.53 4608M-4X-xfs Kswapd efficiency fsmark 90% 94% 4608M-4X-xfs Kswapd efficiency simple-wb 87% 82% 4608M-4X-xfs Kswapd efficiency mmap-strm 43% 43% 4608M-16X-xfs Files/s mean 26.07 ( 0.00%) 26.42 ( 1.32%) 4608M-16X-xfs Elapsed Time fsmark 602.69 585.78 4608M-16X-xfs Elapsed Time simple-wb 606.60 573.81 4608M-16X-xfs Elapsed Time mmap-strm 1549.75 1441.86 4608M-16X-xfs Kswapd efficiency fsmark 98% 98% 4608M-16X-xfs Kswapd efficiency simple-wb 88% 82% 4608M-16X-xfs Kswapd efficiency mmap-strm 44% 42% Unlike the other tests, the fsmark results are not statistically significant but the min and max times are both improved and for the most part, tests completed faster. There are other indications that this is an improvement as well. For example, in the vast majority of cases, there were fewer pages scanned by direct reclaim implying in many cases that stalls due to direct reclaim are reduced. KSwapd is scanning more due to skipping dirty pages which is unfortunate but the CPU usage is still acceptable In an earlier set of tests, I used blktrace and in almost all cases throughput throughout the entire test was higher. However, I ended up discarding those results as recording blktrace data was too heavy for my liking. On a laptop, I plugged in a USB stick and ran a similar tests of tests using it as backing storage. A desktop environment was running and for the entire duration of the tests, firefox and gnome terminal were launching and exiting to vaguely simulate a user. 1024M-xfs Files/s mean 0.41 ( 0.00%) 0.44 ( 6.82%) 1024M-xfs Elapsed Time fsmark 2053.52 1641.03 1024M-xfs Elapsed Time simple-wb 1229.53 768.05 1024M-xfs Elapsed Time mmap-strm 4126.44 4597.03 1024M-xfs Kswapd efficiency fsmark 84% 85% 1024M-xfs Kswapd efficiency simple-wb 92% 81% 1024M-xfs Kswapd efficiency mmap-strm 60% 51% 1024M-xfs Avg wait ms fsmark 5404.53 4473.87 1024M-xfs Avg wait ms simple-wb 2541.35 1453.54 1024M-xfs Avg wait ms mmap-strm 3400.25 3852.53 The mmap-strm results were hurt because firefox launching had a tendency to push the test out of memory. On the postive side, firefox launched marginally faster with the patches applied. Time to completion for many tests was faster but more importantly - the "Avg wait" time as measured by iostat was far lower implying the system would be more responsive. It was also the case that "Avg wait ms" on the root filesystem was lower. I tested it manually and while the system felt slightly more responsive while copying data to a USB stick, it was marginal enough that it could be my imagination. This patch: do not writeback filesystem pages in direct reclaim. When kswapd is failing to keep zones above the min watermark, a process will enter direct reclaim in the same manner kswapd does. If a dirty page is encountered during the scan, this page is written to backing storage using mapping->writepage. This causes two problems. First, it can result in very deep call stacks, particularly if the target storage or filesystem are complex. Some filesystems ignore write requests from direct reclaim as a result. The second is that a single-page flush is inefficient in terms of IO. While there is an expectation that the elevator will merge requests, this does not always happen. Quoting Christoph Hellwig; The elevator has a relatively small window it can operate on, and can never fix up a bad large scale writeback pattern. This patch prevents direct reclaim writing back filesystem pages by checking if current is kswapd. Anonymous pages are still written to swap as there is not the equivalent of a flusher thread for anonymous pages. If the dirty pages cannot be written back, they are placed back on the LRU lists. There is now a direct dependency on dirty page balancing to prevent too many pages in the system being dirtied which would prevent reclaim making forward progress. Signed-off-by: Mel Gorman <mgorman@suse.de> Reviewed-by: Minchan Kim <minchan.kim@gmail.com> Cc: Dave Chinner <david@fromorbit.com> Cc: Christoph Hellwig <hch@infradead.org> Cc: Johannes Weiner <jweiner@redhat.com> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Jan Kara <jack@suse.cz> Cc: Rik van Riel <riel@redhat.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Alex Elder <aelder@sgi.com> Cc: Theodore Ts'o <tytso@mit.edu> Cc: Chris Mason <chris.mason@oracle.com> Cc: Dave Hansen <dave@linux.vnet.ibm.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: vmscan: drop nr_force_scan[] from get_scan_countJohannes Weiner
The nr_force_scan[] tuple holds the effective scan numbers for anon and file pages in case the situation called for a forced scan and the regularly calculated scan numbers turned out zero. However, the effective scan number can always be assumed to be SWAP_CLUSTER_MAX right before the division into anon and file. The numerators and denominator are properly set up for all cases, be it force scan for just file, just anon, or both, to do the right thing. Signed-off-by: Johannes Weiner <jweiner@redhat.com> Reviewed-by: Minchan Kim <minchan.kim@gmail.com> Acked-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Reviewed-by: Michal Hocko <mhocko@suse.cz> Cc: Ying Han <yinghan@google.com> Cc: Balbir Singh <bsingharora@gmail.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> Acked-by: Mel Gorman <mel@csn.ul.ie> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: output a list of loaded modules when we hit bad_page()Dave Jones
When we get a bad_page bug report, it's useful to see what modules the user had loaded. Signed-off-by: Dave Jones <davej@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31oom: fix race while temporarily setting current's oom_score_adjDavid Rientjes
test_set_oom_score_adj() was introduced in 72788c385604 ("oom: replace PF_OOM_ORIGIN with toggling oom_score_adj") to temporarily elevate current's oom_score_adj for ksm and swapoff without requiring an additional per-process flag. Using that function to both set oom_score_adj to OOM_SCORE_ADJ_MAX and then reinstate the previous value is racy since it's possible that userspace can set the value to something else itself before the old value is reinstated. That results in userspace setting current's oom_score_adj to a different value and then the kernel immediately setting it back to its previous value without notification. To fix this, a new compare_swap_oom_score_adj() function is introduced with the same semantics as the compare and swap CAS instruction, or CMPXCHG on x86. It is used to reinstate the previous value of oom_score_adj if and only if the present value is the same as the old value. Signed-off-by: David Rientjes <rientjes@google.com> Cc: Oleg Nesterov <oleg@redhat.com> Cc: Ying Han <yinghan@google.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31oom: remove oom_disable_countDavid Rientjes
This removes mm->oom_disable_count entirely since it's unnecessary and currently buggy. The counter was intended to be per-process but it's currently decremented in the exit path for each thread that exits, causing it to underflow. The count was originally intended to prevent oom killing threads that share memory with threads that cannot be killed since it doesn't lead to future memory freeing. The counter could be fixed to represent all threads sharing the same mm, but it's better to remove the count since: - it is possible that the OOM_DISABLE thread sharing memory with the victim is waiting on that thread to exit and will actually cause future memory freeing, and - there is no guarantee that a thread is disabled from oom killing just because another thread sharing its mm is oom disabled. Signed-off-by: David Rientjes <rientjes@google.com> Reported-by: Oleg Nesterov <oleg@redhat.com> Reviewed-by: Oleg Nesterov <oleg@redhat.com> Cc: Ying Han <yinghan@google.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31oom: avoid killing kthreads if they assume the oom killed thread's mmDavid Rientjes
After selecting a task to kill, the oom killer iterates all processes and kills all other threads that share the same mm_struct in different thread groups. It would not otherwise be helpful to kill a thread if its memory would not be subsequently freed. A kernel thread, however, may assume a user thread's mm by using use_mm(). This is only temporary and should not result in sending a SIGKILL to that kthread. This patch ensures that only user threads and not kthreads are sent a SIGKILL if they share the same mm_struct as the oom killed task. Signed-off-by: David Rientjes <rientjes@google.com> Reviewed-by: Michal Hocko <mhocko@suse.cz> Reviewed-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31oom: thaw threads if oom killed thread is frozen before deferringDavid Rientjes
If a thread has been oom killed and is frozen, thaw it before returning to the page allocator. Otherwise, it can stay frozen indefinitely and no memory will be freed. Signed-off-by: David Rientjes <rientjes@google.com> Reported-by: Konstantin Khlebnikov <khlebnikov@openvz.org> Cc: Oleg Nesterov <oleg@redhat.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: "Rafael J. Wysocki" <rjw@sisk.pl> Acked-by: Michal Hocko <mhocko@suse.cz> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm/page-writeback.c: document bdi_min_ratioJohannes Weiner
Looks like someone got distracted after adding the comment characters. Signed-off-by: Johannes Weiner <jweiner@redhat.com> Acked-by: Peter Zijlstra <peterz@infradead.org> Cc: Wu Fengguang <fengguang.wu@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31vmscan: add block plug for page reclaimShaohua Li
per-task block plug can reduce block queue lock contention and increase request merge. Currently page reclaim doesn't support it. I originally thought page reclaim doesn't need it, because kswapd thread count is limited and file cache write is done at flusher mostly. When I test a workload with heavy swap in a 4-node machine, each CPU is doing direct page reclaim and swap. This causes block queue lock contention. In my test, without below patch, the CPU utilization is about 2% ~ 7%. With the patch, the CPU utilization is about 1% ~ 3%. Disk throughput isn't changed. This should improve normal kswapd write and file cache write too (increase request merge for example), but might not be so obvious as I explain above. Signed-off-by: Shaohua Li <shaohua.li@intel.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Minchan Kim <minchan.kim@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: migration: clean up unmap_and_move()Minchan Kim
unmap_and_move() is one a big messy function. Clean it up. Signed-off-by: Minchan Kim <minchan.kim@gmail.com> Reviewed-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Rik van Riel <riel@redhat.com> Cc: Michal Hocko <mhocko@suse.cz> Cc: Andrea Arcangeli <aarcange@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: zone_reclaim: make isolate_lru_page() filter-awareMinchan Kim
In __zone_reclaim case, we don't want to shrink mapped page. Nonetheless, we have isolated mapped page and re-add it into LRU's head. It's unnecessary CPU overhead and makes LRU churning. Of course, when we isolate the page, the page might be mapped but when we try to migrate the page, the page would be not mapped. So it could be migrated. But race is rare and although it happens, it's no big deal. Signed-off-by: Minchan Kim <minchan.kim@gmail.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Reviewed-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Reviewed-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Reviewed-by: Michal Hocko <mhocko@suse.cz> Cc: Mel Gorman <mgorman@suse.de> Cc: Rik van Riel <riel@redhat.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: compaction: make isolate_lru_page() filter-awareMinchan Kim
In async mode, compaction doesn't migrate dirty or writeback pages. So, it's meaningless to pick the page and re-add it to lru list. Of course, when we isolate the page in compaction, the page might be dirty or writeback but when we try to migrate the page, the page would be not dirty, writeback. So it could be migrated. But it's very unlikely as isolate and migration cycle is much faster than writeout. So, this patch helps cpu overhead and prevent unnecessary LRU churning. Signed-off-by: Minchan Kim <minchan.kim@gmail.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Reviewed-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Reviewed-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Acked-by: Mel Gorman <mgorman@suse.de> Acked-by: Rik van Riel <riel@redhat.com> Reviewed-by: Michal Hocko <mhocko@suse.cz> Cc: Andrea Arcangeli <aarcange@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: change isolate mode from #define to bitwise typeMinchan Kim
Change ISOLATE_XXX macro with bitwise isolate_mode_t type. Normally, macro isn't recommended as it's type-unsafe and making debugging harder as symbol cannot be passed throught to the debugger. Quote from Johannes " Hmm, it would probably be cleaner to fully convert the isolation mode into independent flags. INACTIVE, ACTIVE, BOTH is currently a tri-state among flags, which is a bit ugly." This patch moves isolate mode from swap.h to mmzone.h by memcontrol.h Signed-off-by: Minchan Kim <minchan.kim@gmail.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Rik van Riel <riel@redhat.com> Cc: Michal Hocko <mhocko@suse.cz> Cc: Andrea Arcangeli <aarcange@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: compaction: trivial clean up in acct_isolated()Minchan Kim
acct_isolated of compaction uses page_lru_base_type which returns only base type of LRU list so it never returns LRU_ACTIVE_ANON or LRU_ACTIVE_FILE. In addtion, cc->nr_[anon|file] is used in only acct_isolated so it doesn't have fields in conpact_control. This patch removes fields from compact_control and makes clear function of acct_issolated which counts the number of anon|file pages isolated. Signed-off-by: Minchan Kim <minchan.kim@gmail.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Reviewed-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Reviewed-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Acked-by: Mel Gorman <mgorman@suse.de> Acked-by: Rik van Riel <riel@redhat.com> Reviewed-by: Michal Hocko <mhocko@suse.cz> Cc: Andrea Arcangeli <aarcange@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31Cross Memory AttachChristopher Yeoh
The basic idea behind cross memory attach is to allow MPI programs doing intra-node communication to do a single copy of the message rather than a double copy of the message via shared memory. The following patch attempts to achieve this by allowing a destination process, given an address and size from a source process, to copy memory directly from the source process into its own address space via a system call. There is also a symmetrical ability to copy from the current process's address space into a destination process's address space. - Use of /proc/pid/mem has been considered, but there are issues with using it: - Does not allow for specifying iovecs for both src and dest, assuming preadv or pwritev was implemented either the area read from or written to would need to be contiguous. - Currently mem_read allows only processes who are currently ptrace'ing the target and are still able to ptrace the target to read from the target. This check could possibly be moved to the open call, but its not clear exactly what race this restriction is stopping (reason appears to have been lost) - Having to send the fd of /proc/self/mem via SCM_RIGHTS on unix domain socket is a bit ugly from a userspace point of view, especially when you may have hundreds if not (eventually) thousands of processes that all need to do this with each other - Doesn't allow for some future use of the interface we would like to consider adding in the future (see below) - Interestingly reading from /proc/pid/mem currently actually involves two copies! (But this could be fixed pretty easily) As mentioned previously use of vmsplice instead was considered, but has problems. Since you need the reader and writer working co-operatively if the pipe is not drained then you block. Which requires some wrapping to do non blocking on the send side or polling on the receive. In all to all communication it requires ordering otherwise you can deadlock. And in the example of many MPI tasks writing to one MPI task vmsplice serialises the copying. There are some cases of MPI collectives where even a single copy interface does not get us the performance gain we could. For example in an MPI_Reduce rather than copy the data from the source we would like to instead use it directly in a mathops (say the reduce is doing a sum) as this would save us doing a copy. We don't need to keep a copy of the data from the source. I haven't implemented this, but I think this interface could in the future do all this through the use of the flags - eg could specify the math operation and type and the kernel rather than just copying the data would apply the specified operation between the source and destination and store it in the destination. Although we don't have a "second user" of the interface (though I've had some nibbles from people who may be interested in using it for intra process messaging which is not MPI). This interface is something which hardware vendors are already doing for their custom drivers to implement fast local communication. And so in addition to this being useful for OpenMPI it would mean the driver maintainers don't have to fix things up when the mm changes. There was some discussion about how much faster a true zero copy would go. Here's a link back to the email with some testing I did on that: http://marc.info/?l=linux-mm&m=130105930902915&w=2 There is a basic man page for the proposed interface here: http://ozlabs.org/~cyeoh/cma/process_vm_readv.txt This has been implemented for x86 and powerpc, other architecture should mainly (I think) just need to add syscall numbers for the process_vm_readv and process_vm_writev. There are 32 bit compatibility versions for 64-bit kernels. For arch maintainers there are some simple tests to be able to quickly verify that the syscalls are working correctly here: http://ozlabs.org/~cyeoh/cma/cma-test-20110718.tgz Signed-off-by: Chris Yeoh <yeohc@au1.ibm.com> Cc: Ingo Molnar <mingo@elte.hu> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Arnd Bergmann <arnd@arndb.de> Cc: Paul Mackerras <paulus@samba.org> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: David Howells <dhowells@redhat.com> Cc: James Morris <jmorris@namei.org> Cc: <linux-man@vger.kernel.org> Cc: <linux-arch@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-31mm: fix implicit stat.h usage in dmapool.cPaul Gortmaker
The removal of the implicitly everywhere module.h and its child includes will reveal this implicit stat.h usage: mm/dmapool.c:108: error: ‘S_IRUGO’ undeclared here (not in a function) Signed-off-by: Paul Gortmaker <paul.gortmaker@windriver.com>
2011-10-31mm: Map most files to use export.h instead of module.hPaul Gortmaker
The files changed within are only using the EXPORT_SYMBOL macro variants. They are not using core modular infrastructure and hence don't need module.h but only the export.h header. Signed-off-by: Paul Gortmaker <paul.gortmaker@windriver.com>
2011-10-31mm: Add export.h for EXPORT_SYMBOL to active symbol exportersPaul Gortmaker
These files were getting <linux/module.h> via an implicit include path, but we want to crush those out of existence since they cost time during compiles of processing thousands of lines of headers for no reason. Give them the lightweight header that just contains the EXPORT_SYMBOL infrastructure. Signed-off-by: Paul Gortmaker <paul.gortmaker@windriver.com>
2011-10-31mm: delete various needless include <linux/module.h>Paul Gortmaker
There is nothing modular in these files, and no reason to drag in all the 357 headers that module.h brings with it, since it just slows down compiles. Signed-off-by: Paul Gortmaker <paul.gortmaker@windriver.com>
2011-10-31writeback: Add a 'reason' to wb_writeback_workCurt Wohlgemuth
This creates a new 'reason' field in a wb_writeback_work structure, which unambiguously identifies who initiates writeback activity. A 'wb_reason' enumeration has been added to writeback.h, to enumerate the possible reasons. The 'writeback_work_class' and tracepoint event class and 'writeback_queue_io' tracepoints are updated to include the symbolic 'reason' in all trace events. And the 'writeback_inodes_sbXXX' family of routines has had a wb_stats parameter added to them, so callers can specify why writeback is being started. Acked-by: Jan Kara <jack@suse.cz> Signed-off-by: Curt Wohlgemuth <curtw@google.com> Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-31writeback: trace event balance_dirty_pagesWu Fengguang
Useful for analyzing the dynamics of the throttling algorithms and debugging user reported problems. Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-31writeback: trace event bdi_dirty_ratelimitWu Fengguang
It helps understand how various throttle bandwidths are updated. Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-28Merge branch 'for-next' of ↵Linus Torvalds
git://git.kernel.org/pub/scm/linux/kernel/git/hch/vfs-queue * 'for-next' of git://git.kernel.org/pub/scm/linux/kernel/git/hch/vfs-queue: (21 commits) leases: fix write-open/read-lease race nfs: drop unnecessary locking in llseek ext4: replace cut'n'pasted llseek code with generic_file_llseek_size vfs: add generic_file_llseek_size vfs: do (nearly) lockless generic_file_llseek direct-io: merge direct_io_walker into __blockdev_direct_IO direct-io: inline the complete submission path direct-io: separate map_bh from dio direct-io: use a slab cache for struct dio direct-io: rearrange fields in dio/dio_submit to avoid holes direct-io: fix a wrong comment direct-io: separate fields only used in the submission path from struct dio vfs: fix spinning prevention in prune_icache_sb vfs: add a comment to inode_permission() vfs: pass all mask flags check_acl and posix_acl_permission vfs: add hex format for MAY_* flag values vfs: indicate that the permission functions take all the MAY_* flags compat: sync compat_stats with statfs. vfs: add "device" tag to /proc/self/mountstats cleanup: vfs: small comment fix for block_invalidatepage ... Fix up trivial conflict in fs/gfs2/file.c (llseek changes)
2011-10-28vfs: iov_iter: have iov_iter_advance decrement nr_segs appropriatelyJeff Layton
Currently, when you call iov_iter_advance, then the pointer to the iovec array can be incremented, but it does not decrement the nr_segs value in the iov_iter struct. The result is a iov_iter struct with a nr_segs value that goes beyond the end of the array. While I'm not aware of anything that's specifically broken by this, it seems odd and a bit dangerous not to decrement that value. If someone were to trust the nr_segs value to be correct, then they could end up walking off the end of the array. Changing this might also provide some micro-optimization when dealing with the last iovec in an array. Many of the other routines that deal with iov_iter have optimized codepaths when nr_segs == 1. Cc: Nick Piggin <npiggin@suse.de> Signed-off-by: Jeff Layton <jlayton@redhat.com> Signed-off-by: Christoph Hellwig <hch@lst.de>
2011-10-26Merge branches 'slab/next' and 'slub/partial' into slab/for-linusPekka Enberg
2011-10-25Merge branch 'for-linus' of ↵Linus Torvalds
git://git.kernel.org/pub/scm/linux/kernel/git/jikos/trivial * 'for-linus' of git://git.kernel.org/pub/scm/linux/kernel/git/jikos/trivial: (59 commits) MAINTAINERS: linux-m32r is moderated for non-subscribers linux@lists.openrisc.net is moderated for non-subscribers Drop default from "DM365 codec select" choice parisc: Kconfig: cleanup Kernel page size default Kconfig: remove redundant CONFIG_ prefix on two symbols cris: remove arch/cris/arch-v32/lib/nand_init.S microblaze: add missing CONFIG_ prefixes h8300: drop puzzling Kconfig dependencies MAINTAINERS: microblaze-uclinux@itee.uq.edu.au is moderated for non-subscribers tty: drop superfluous dependency in Kconfig ARM: mxc: fix Kconfig typo 'i.MX51' Fix file references in Kconfig files aic7xxx: fix Kconfig references to READMEs Fix file references in drivers/ide/ thinkpad_acpi: Fix printk typo 'bluestooth' bcmring: drop commented out line in Kconfig btmrvl_sdio: fix typo 'btmrvl_sdio_sd6888' doc: raw1394: Trivial typo fix CIFS: Don't free volume_info->UNC until we are entirely done with it. treewide: Correct spelling of successfully in comments ...
2011-10-25Merge branch 'next' of git://selinuxproject.org/~jmorris/linux-securityLinus Torvalds
* 'next' of git://selinuxproject.org/~jmorris/linux-security: (95 commits) TOMOYO: Fix incomplete read after seek. Smack: allow to access /smack/access as normal user TOMOYO: Fix unused kernel config option. Smack: fix: invalid length set for the result of /smack/access Smack: compilation fix Smack: fix for /smack/access output, use string instead of byte Smack: domain transition protections (v3) Smack: Provide information for UDS getsockopt(SO_PEERCRED) Smack: Clean up comments Smack: Repair processing of fcntl Smack: Rule list lookup performance Smack: check permissions from user space (v2) TOMOYO: Fix quota and garbage collector. TOMOYO: Remove redundant tasklist_lock. TOMOYO: Fix domain transition failure warning. TOMOYO: Remove tomoyo_policy_memory_lock spinlock. TOMOYO: Simplify garbage collector. TOMOYO: Fix make namespacecheck warnings. target: check hex2bin result encrypted-keys: check hex2bin result ...
2011-10-20block: initialize the bounce pool if high memory may be added laterDavid Vrabel
init_emergency_pool() does not create the page pool for bouncing block requests if the current count of high pages is zero. If high memory may be added later (either via memory hotplug or a balloon driver in a virtualized system) then a oops occurs if a request with a high page need bouncing because the pool does not exist. So, always create the pool if memory hotplug is enabled and change the test so it's valid even if all high pages are currently in the balloon (the balloon drivers adjust totalhigh_pages but not max_pfn). Signed-off-by: David Vrabel <david.vrabel@citrix.com> Signed-off-by: Jens Axboe <axboe@kernel.dk>
2011-10-19mm: fix race between mremap and removing migration entryHugh Dickins
I don't usually pay much attention to the stale "? " addresses in stack backtraces, but this lucky report from Pawel Sikora hints that mremap's move_ptes() has inadequate locking against page migration. 3.0 BUG_ON(!PageLocked(p)) in migration_entry_to_page(): kernel BUG at include/linux/swapops.h:105! RIP: 0010:[<ffffffff81127b76>] [<ffffffff81127b76>] migration_entry_wait+0x156/0x160 [<ffffffff811016a1>] handle_pte_fault+0xae1/0xaf0 [<ffffffff810feee2>] ? __pte_alloc+0x42/0x120 [<ffffffff8112c26b>] ? do_huge_pmd_anonymous_page+0xab/0x310 [<ffffffff81102a31>] handle_mm_fault+0x181/0x310 [<ffffffff81106097>] ? vma_adjust+0x537/0x570 [<ffffffff81424bed>] do_page_fault+0x11d/0x4e0 [<ffffffff81109a05>] ? do_mremap+0x2d5/0x570 [<ffffffff81421d5f>] page_fault+0x1f/0x30 mremap's down_write of mmap_sem, together with i_mmap_mutex or lock, and pagetable locks, were good enough before page migration (with its requirement that every migration entry be found) came in, and enough while migration always held mmap_sem; but not enough nowadays, when there's memory hotremove and compaction. The danger is that move_ptes() lets a migration entry dodge around behind remove_migration_pte()'s back, so it's in the old location when looking at the new, then in the new location when looking at the old. Either mremap's move_ptes() must additionally take anon_vma lock(), or migration's remove_migration_pte() must stop peeking for is_swap_entry() before it takes pagetable lock. Consensus chooses the latter: we prefer to add overhead to migration than to mremapping, which gets used by JVMs and by exec stack setup. Reported-and-tested-by: Paweł Sikora <pluto@agmk.net> Signed-off-by: Hugh Dickins <hughd@google.com> Acked-by: Andrea Arcangeli <aarcange@redhat.com> Acked-by: Mel Gorman <mgorman@suse.de> Cc: stable@vger.kernel.org Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-10-11writeback: fix ppc compile warnings on do_div(long long, unsigned long)Wu Fengguang
Fix powerpc compile warnings mm/page-writeback.c: In function 'bdi_position_ratio': mm/page-writeback.c:622:3: warning: comparison of distinct pointer types lacks a cast [enabled by default] page-writeback.c:635:4: warning: comparison of distinct pointer types lacks a cast [enabled by default] Also fix gcc "uninitialized var" warnings. Reported-by: Stephen Rothwell <sfr@canb.auug.org.au> Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-03writeback: dirty position control - bdi reserve areaWu Fengguang
Keep a minimal pool of dirty pages for each bdi, so that the disk IO queues won't underrun. Also gently increase a small bdi_thresh to avoid it stuck in 0 for some light dirtied bdi. It's particularly useful for JBOD and small memory system. It may result in (pos_ratio > 1) at the setpoint and push the dirty pages high. This is more or less intended because the bdi is in the danger of IO queue underflow. Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-03writeback: control dirty pause timeWu Fengguang
The dirty pause time shall ultimately be controlled by adjusting nr_dirtied_pause, since there is relationship pause = pages_dirtied / task_ratelimit Assuming pages_dirtied ~= nr_dirtied_pause task_ratelimit ~= dirty_ratelimit We get nr_dirtied_pause ~= dirty_ratelimit * desired_pause Here dirty_ratelimit is preferred over task_ratelimit because it's more stable. It's also important to limit possible large transitional errors: - bw is changing quickly - pages_dirtied << nr_dirtied_pause on entering dirty exceeded area - pages_dirtied >> nr_dirtied_pause on btrfs (to be improved by a separate fix, but still expect non-trivial errors) So we end up using the above formula inside clamp_val(). The best test case for this code is to run 100 "dd bs=4M" tasks on btrfs and check its pause time distribution. Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-03writeback: limit max dirty pause timeWu Fengguang
Apply two policies to scale down the max pause time for 1) small number of concurrent dirtiers 2) small memory system (comparing to storage bandwidth) MAX_PAUSE=200ms may only be suitable for high end servers with lots of concurrent dirtiers, where the large pause time can reduce much overheads. Otherwise, smaller pause time is desirable whenever possible, so as to get good responsiveness and smooth user experiences. It's actually required for good disk utilization in the case when all the dirty pages can be synced to disk within MAX_PAUSE=200ms. Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-03writeback: IO-less balance_dirty_pages()Wu Fengguang
As proposed by Chris, Dave and Jan, don't start foreground writeback IO inside balance_dirty_pages(). Instead, simply let it idle sleep for some time to throttle the dirtying task. In the mean while, kick off the per-bdi flusher thread to do background writeback IO. RATIONALS ========= - disk seeks on concurrent writeback of multiple inodes (Dave Chinner) If every thread doing writes and being throttled start foreground writeback, it leads to N IO submitters from at least N different inodes at the same time, end up with N different sets of IO being issued with potentially zero locality to each other, resulting in much lower elevator sort/merge efficiency and hence we seek the disk all over the place to service the different sets of IO. OTOH, if there is only one submission thread, it doesn't jump between inodes in the same way when congestion clears - it keeps writing to the same inode, resulting in large related chunks of sequential IOs being issued to the disk. This is more efficient than the above foreground writeback because the elevator works better and the disk seeks less. - lock contention and cache bouncing on concurrent IO submitters (Dave Chinner) With this patchset, the fs_mark benchmark on a 12-drive software RAID0 goes from CPU bound to IO bound, freeing "3-4 CPUs worth of spinlock contention". * "CPU usage has dropped by ~55%", "it certainly appears that most of the CPU time saving comes from the removal of contention on the inode_wb_list_lock" (IMHO at least 10% comes from the reduction of cacheline bouncing, because the new code is able to call much less frequently into balance_dirty_pages() and hence access the global page states) * the user space "App overhead" is reduced by 20%, by avoiding the cacheline pollution by the complex writeback code path * "for a ~5% throughput reduction", "the number of write IOs have dropped by ~25%", and the elapsed time reduced from 41:42.17 to 40:53.23. * On a simple test of 100 dd, it reduces the CPU %system time from 30% to 3%, and improves IO throughput from 38MB/s to 42MB/s. - IO size too small for fast arrays and too large for slow USB sticks The write_chunk used by current balance_dirty_pages() cannot be directly set to some large value (eg. 128MB) for better IO efficiency. Because it could lead to more than 1 second user perceivable stalls. Even the current 4MB write size may be too large for slow USB sticks. The fact that balance_dirty_pages() starts IO on itself couples the IO size to wait time, which makes it hard to do suitable IO size while keeping the wait time under control. Now it's possible to increase writeback chunk size proportional to the disk bandwidth. In a simple test of 50 dd's on XFS, 1-HDD, 3GB ram, the larger writeback size dramatically reduces the seek count to 1/10 (far beyond my expectation) and improves the write throughput by 24%. - long block time in balance_dirty_pages() hurts desktop responsiveness Many of us may have the experience: it often takes a couple of seconds or even long time to stop a heavy writing dd/cp/tar command with Ctrl-C or "kill -9". - IO pipeline broken by bumpy write() progress There are a broad class of "loop {read(buf); write(buf);}" applications whose read() pipeline will be under-utilized or even come to a stop if the write()s have long latencies _or_ don't progress in a constant rate. The current threshold based throttling inherently transfers the large low level IO completion fluctuations to bumpy application write()s, and further deteriorates with increasing number of dirtiers and/or bdi's. For example, when doing 50 dd's + 1 remote rsync to an XFS partition, the rsync progresses very bumpy in legacy kernel, and throughput is improved by 67% by this patchset. (plus the larger write chunk size, it will be 93% speedup). The new rate based throttling can support 1000+ dd's with excellent smoothness, low latency and low overheads. For the above reasons, it's much better to do IO-less and low latency pauses in balance_dirty_pages(). Jan Kara, Dave Chinner and me explored the scheme to let balance_dirty_pages() wait for enough writeback IO completions to safeguard the dirty limit. However it's found to have two problems: - in large NUMA systems, the per-cpu counters may have big accounting errors, leading to big throttle wait time and jitters. - NFS may kill large amount of unstable pages with one single COMMIT. Because NFS server serves COMMIT with expensive fsync() IOs, it is desirable to delay and reduce the number of COMMITs. So it's not likely to optimize away such kind of bursty IO completions, and the resulted large (and tiny) stall times in IO completion based throttling. So here is a pause time oriented approach, which tries to control the pause time in each balance_dirty_pages() invocations, by controlling the number of pages dirtied before calling balance_dirty_pages(), for smooth and efficient dirty throttling: - avoid useless (eg. zero pause time) balance_dirty_pages() calls - avoid too small pause time (less than 4ms, which burns CPU power) - avoid too large pause time (more than 200ms, which hurts responsiveness) - avoid big fluctuations of pause times It can control pause times at will. The default policy (in a followup patch) will be to do ~10ms pauses in 1-dd case, and increase to ~100ms in 1000-dd case. BEHAVIOR CHANGE =============== (1) dirty threshold Users will notice that the applications will get throttled once crossing the global (background + dirty)/2=15% threshold, and then balanced around 17.5%. Before patch, the behavior is to just throttle it at 20% dirtyable memory in 1-dd case. Since the task will be soft throttled earlier than before, it may be perceived by end users as performance "slow down" if his application happens to dirty more than 15% dirtyable memory. (2) smoothness/responsiveness Users will notice a more responsive system during heavy writeback. "killall dd" will take effect instantly. Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-03writeback: per task dirty rate limitWu Fengguang
Add two fields to task_struct. 1) account dirtied pages in the individual tasks, for accuracy 2) per-task balance_dirty_pages() call intervals, for flexibility The balance_dirty_pages() call interval (ie. nr_dirtied_pause) will scale near-sqrt to the safety gap between dirty pages and threshold. The main problem of per-task nr_dirtied is, if 1k+ tasks start dirtying pages at exactly the same time, each task will be assigned a large initial nr_dirtied_pause, so that the dirty threshold will be exceeded long before each task reached its nr_dirtied_pause and hence call balance_dirty_pages(). The solution is to watch for the number of pages dirtied on each CPU in between the calls into balance_dirty_pages(). If it exceeds ratelimit_pages (3% dirty threshold), force call balance_dirty_pages() for a chance to set bdi->dirty_exceeded. In normal situations, this safeguarding condition is not expected to trigger at all. On the sqrt in dirty_poll_interval(): It will serve as an initial guess when dirty pages are still in the freerun area. When dirty pages are floating inside the dirty control scope [freerun, limit], a followup patch will use some refined dirty poll interval to get the desired pause time. thresh-dirty (MB) sqrt 1 16 2 22 4 32 8 45 16 64 32 90 64 128 128 181 256 256 512 362 1024 512 The above table means, given 1MB (or 1GB) gap and the dd tasks polling balance_dirty_pages() on every 16 (or 512) pages, the dirty limit won't be exceeded as long as there are less than 16 (or 512) concurrent dd's. So sqrt naturally leads to less overheads and more safe concurrent tasks for large memory servers, which have large (thresh-freerun) gaps. peter: keep the per-CPU ratelimit for safeguarding the 1k+ tasks case CC: Peter Zijlstra <a.p.zijlstra@chello.nl> Reviewed-by: Andrea Righi <andrea@betterlinux.com> Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-03writeback: stabilize bdi->dirty_ratelimitWu Fengguang
There are some imperfections in balanced_dirty_ratelimit. 1) large fluctuations The dirty_rate used for computing balanced_dirty_ratelimit is merely averaged in the past 200ms (very small comparing to the 3s estimation period for write_bw), which makes rather dispersed distribution of balanced_dirty_ratelimit. It's pretty hard to average out the singular points by increasing the estimation period. Considering that the averaging technique will introduce very undesirable time lags, I give it up totally. (btw, the 3s write_bw averaging time lag is much more acceptable because its impact is one-way and therefore won't lead to oscillations.) The more practical way is filtering -- most singular balanced_dirty_ratelimit points can be filtered out by remembering some prev_balanced_rate and prev_prev_balanced_rate. However the more reliable way is to guard balanced_dirty_ratelimit with task_ratelimit. 2) due to truncates and fs redirties, the (write_bw <=> dirty_rate) match could become unbalanced, which may lead to large systematical errors in balanced_dirty_ratelimit. The truncates, due to its possibly bumpy nature, can hardly be compensated smoothly. So let's face it. When some over-estimated balanced_dirty_ratelimit brings dirty_ratelimit high, dirty pages will go higher than the setpoint. task_ratelimit will in turn become lower than dirty_ratelimit. So if we consider both balanced_dirty_ratelimit and task_ratelimit and update dirty_ratelimit only when they are on the same side of dirty_ratelimit, the systematical errors in balanced_dirty_ratelimit won't be able to bring dirty_ratelimit far away. The balanced_dirty_ratelimit estimation may also be inaccurate near @limit or @freerun, however is less an issue. 3) since we ultimately want to - keep the fluctuations of task ratelimit as small as possible - keep the dirty pages around the setpoint as long time as possible the update policy used for (2) also serves the above goals nicely: if for some reason the dirty pages are high (task_ratelimit < dirty_ratelimit), and dirty_ratelimit is low (dirty_ratelimit < balanced_dirty_ratelimit), there is no point to bring up dirty_ratelimit in a hurry only to hurt both the above two goals. So, we make use of task_ratelimit to limit the update of dirty_ratelimit in two ways: 1) avoid changing dirty rate when it's against the position control target (the adjusted rate will slow down the progress of dirty pages going back to setpoint). 2) limit the step size. task_ratelimit is changing values step by step, leaving a consistent trace comparing to the randomly jumping balanced_dirty_ratelimit. task_ratelimit also has the nice smaller errors in stable state and typically larger errors when there are big errors in rate. So it's a pretty good limiting factor for the step size of dirty_ratelimit. Note that bdi->dirty_ratelimit is always tracking balanced_dirty_ratelimit. task_ratelimit is merely used as a limiting factor. Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-03writeback: dirty rate controlWu Fengguang
It's all about bdi->dirty_ratelimit, which aims to be (write_bw / N) when there are N dd tasks. On write() syscall, use bdi->dirty_ratelimit ============================================ balance_dirty_pages(pages_dirtied) { task_ratelimit = bdi->dirty_ratelimit * bdi_position_ratio(); pause = pages_dirtied / task_ratelimit; sleep(pause); } On every 200ms, update bdi->dirty_ratelimit =========================================== bdi_update_dirty_ratelimit() { task_ratelimit = bdi->dirty_ratelimit * bdi_position_ratio(); balanced_dirty_ratelimit = task_ratelimit * write_bw / dirty_rate; bdi->dirty_ratelimit = balanced_dirty_ratelimit } Estimation of balanced bdi->dirty_ratelimit =========================================== balanced task_ratelimit ----------------------- balance_dirty_pages() needs to throttle tasks dirtying pages such that the total amount of dirty pages stays below the specified dirty limit in order to avoid memory deadlocks. Furthermore we desire fairness in that tasks get throttled proportionally to the amount of pages they dirty. IOW we want to throttle tasks such that we match the dirty rate to the writeout bandwidth, this yields a stable amount of dirty pages: dirty_rate == write_bw (1) The fairness requirement gives us: task_ratelimit = balanced_dirty_ratelimit == write_bw / N (2) where N is the number of dd tasks. We don't know N beforehand, but still can estimate balanced_dirty_ratelimit within 200ms. Start by throttling each dd task at rate task_ratelimit = task_ratelimit_0 (3) (any non-zero initial value is OK) After 200ms, we measured dirty_rate = # of pages dirtied by all dd's / 200ms write_bw = # of pages written to the disk / 200ms For the aggressive dd dirtiers, the equality holds dirty_rate == N * task_rate == N * task_ratelimit_0 (4) Or task_ratelimit_0 == dirty_rate / N (5) Now we conclude that the balanced task ratelimit can be estimated by write_bw balanced_dirty_ratelimit = task_ratelimit_0 * ---------- (6) dirty_rate Because with (4) and (5) we can get the desired equality (1): write_bw balanced_dirty_ratelimit == (dirty_rate / N) * ---------- dirty_rate == write_bw / N Then using the balanced task ratelimit we can compute task pause times like: task_pause = task->nr_dirtied / task_ratelimit task_ratelimit with position control ------------------------------------ However, while the above gives us means of matching the dirty rate to the writeout bandwidth, it at best provides us with a stable dirty page count (assuming a static system). In order to control the dirty page count such that it is high enough to provide performance, but does not exceed the specified limit we need another control. The dirty position control works by extending (2) to task_ratelimit = balanced_dirty_ratelimit * pos_ratio (7) where pos_ratio is a negative feedback function that subjects to 1) f(setpoint) = 1.0 2) df/dx < 0 That is, if the dirty pages are ABOVE the setpoint, we throttle each task a bit more HEAVY than balanced_dirty_ratelimit, so that the dirty pages are created less fast than they are cleaned, thus DROP to the setpoints (and the reverse). Based on (7) and the assumption that both dirty_ratelimit and pos_ratio remains CONSTANT for the past 200ms, we get task_ratelimit_0 = balanced_dirty_ratelimit * pos_ratio (8) Putting (8) into (6), we get the formula used in bdi_update_dirty_ratelimit(): write_bw balanced_dirty_ratelimit *= pos_ratio * ---------- (9) dirty_rate Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-03writeback: add bg_threshold parameter to __bdi_update_bandwidth()Wu Fengguang
No behavior change. Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-03writeback: dirty position controlWu Fengguang
bdi_position_ratio() provides a scale factor to bdi->dirty_ratelimit, so that the resulted task rate limit can drive the dirty pages back to the global/bdi setpoints. Old scheme is, | free run area | throttle area ----------------------------------------+----------------------------> thresh^ dirty pages New scheme is, ^ task rate limit | | * | * | * |[free run] * [smooth throttled] | * | * | * ..bdi->dirty_ratelimit..........* | . * | . * | . * | . * | . * +-------------------------------.-----------------------*------------> setpoint^ limit^ dirty pages The slope of the bdi control line should be 1) large enough to pull the dirty pages to setpoint reasonably fast 2) small enough to avoid big fluctuations in the resulted pos_ratio and hence task ratelimit Since the fluctuation range of the bdi dirty pages is typically observed to be within 1-second worth of data, the bdi control line's slope is selected to be a linear function of bdi write bandwidth, so that it can adapt to slow/fast storage devices well. Assume the bdi control line pos_ratio = 1.0 + k * (dirty - bdi_setpoint) where k is the negative slope. If targeting for 12.5% fluctuation range in pos_ratio when dirty pages are fluctuating in range [bdi_setpoint - write_bw/2, bdi_setpoint + write_bw/2], we get slope k = - 1 / (8 * write_bw) Let pos_ratio(x_intercept) = 0, we get the parameter used in code: x_intercept = bdi_setpoint + 8 * write_bw The global/bdi slopes are nicely complementing each other when the system has only one major bdi (indicated by bdi_thresh ~= thresh): 1) slope of global control line => scaling to the control scope size 2) slope of main bdi control line => scaling to the writeout bandwidth so that - in memory tight systems, (1) becomes strong enough to squeeze dirty pages inside the control scope - in large memory systems where the "gravity" of (1) for pulling the dirty pages to setpoint is too weak, (2) can back (1) up and drive dirty pages to bdi_setpoint ~= setpoint reasonably fast. Unfortunately in JBOD setups, the fluctuation range of bdi threshold is related to memory size due to the interferences between disks. In this case, the bdi slope will be weighted sum of write_bw and bdi_thresh. Given equations span = x_intercept - bdi_setpoint k = df/dx = - 1 / span and the extremum values span = bdi_thresh dx = bdi_thresh we get df = - dx / span = - 1.0 That means, when bdi_dirty deviates bdi_thresh up, pos_ratio and hence task ratelimit will fluctuate by -100%. peter: use 3rd order polynomial for the global control line CC: Peter Zijlstra <a.p.zijlstra@chello.nl> Acked-by: Jan Kara <jack@suse.cz> Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-10-03writeback: account per-bdi accumulated dirtied pagesWu Fengguang
Introduce the BDI_DIRTIED counter. It will be used for estimating the bdi's dirty bandwidth. CC: Jan Kara <jack@suse.cz> CC: Michael Rubin <mrubin@google.com> CC: Peter Zijlstra <a.p.zijlstra@chello.nl> Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-09-27slub: Discard slab page when node partial > minimum partial numberAlex Shi
Discarding slab should be done when node partial > min_partial. Otherwise, node partial slab may eat up all memory. Signed-off-by: Alex Shi <alex.shi@intel.com> Acked-by: Christoph Lameter <cl@linux.com> Signed-off-by: Pekka Enberg <penberg@kernel.org>
2011-09-27slub: correct comments error for per cpu partialAlex Shi
Correct comment errors, that mistake cpu partial objects number as pages number, may make reader misunderstand. Signed-off-by: Alex Shi <alex.shi@intel.com> Reviewed-by: Christoph Lameter <cl@linux.com> Signed-off-by: Pekka Enberg <penberg@kernel.org>
2011-09-27mm: restrict access to slab files under procfs and sysfsVasiliy Kulikov
Historically /proc/slabinfo and files under /sys/kernel/slab/* have world read permissions and are accessible to the world. slabinfo contains rather private information related both to the kernel and userspace tasks. Depending on the situation, it might reveal either private information per se or information useful to make another targeted attack. Some examples of what can be learned by reading/watching for /proc/slabinfo entries: 1) dentry (and different *inode*) number might reveal other processes fs activity. The number of dentry "active objects" doesn't strictly show file count opened/touched by a process, however, there is a good correlation between them. The patch "proc: force dcache drop on unauthorized access" relies on the privacy of dentry count. 2) different inode entries might reveal the same information as (1), but these are more fine granted counters. If a filesystem is mounted in a private mount point (or even a private namespace) and fs type differs from other mounted fs types, fs activity in this mount point/namespace is revealed. If there is a single ecryptfs mount point, the whole fs activity of a single user is revealed. Number of files in ecryptfs mount point is a private information per se. 3) fuse_* reveals number of files / fs activity of a user in a user private mount point. It is approx. the same severity as ecryptfs infoleak in (2). 4) sysfs_dir_cache similar to (2) reveals devices' addition/removal, which can be otherwise hidden by "chmod 0700 /sys/". With 0444 slabinfo the precise number of sysfs files is known to the world. 5) buffer_head might reveal some kernel activity. With other information leaks an attacker might identify what specific kernel routines generate buffer_head activity. 6) *kmalloc* infoleaks are very situational. Attacker should watch for the specific kmalloc size entry and filter the noise related to the unrelated kernel activity. If an attacker has relatively silent victim system, he might get rather precise counters. Additional information sources might significantly increase the slabinfo infoleak benefits. E.g. if an attacker knows that the processes activity on the system is very low (only core daemons like syslog and cron), he may run setxid binaries / trigger local daemon activity / trigger network services activity / await sporadic cron jobs activity / etc. and get rather precise counters for fs and network activity of these privileged tasks, which is unknown otherwise. Also hiding slabinfo and /sys/kernel/slab/* is a one step to complicate exploitation of kernel heap overflows (and possibly, other bugs). The related discussion: http://thread.gmane.org/gmane.linux.kernel/1108378 To keep compatibility with old permission model where non-root monitoring daemon could watch for kernel memleaks though slabinfo one should do: groupadd slabinfo usermod -a -G slabinfo $MONITOR_USER And add the following commands to init scripts (to mountall.conf in Ubuntu's upstart case): chmod g+r /proc/slabinfo /sys/kernel/slab/*/* chgrp slabinfo /proc/slabinfo /sys/kernel/slab/*/* Signed-off-by: Vasiliy Kulikov <segoon@openwall.com> Reviewed-by: Kees Cook <kees@ubuntu.com> Reviewed-by: Dave Hansen <dave@linux.vnet.ibm.com> Acked-by: Christoph Lameter <cl@gentwo.org> Acked-by: David Rientjes <rientjes@google.com> CC: Valdis.Kletnieks@vt.edu CC: Linus Torvalds <torvalds@linux-foundation.org> CC: Alan Cox <alan@linux.intel.com> Signed-off-by: Pekka Enberg <penberg@kernel.org>
2011-09-21Merge branch 'for-linus' of git://git.kernel.dk/linux-blockLinus Torvalds
* 'for-linus' of git://git.kernel.dk/linux-block: floppy: use del_timer_sync() in init cleanup blk-cgroup: be able to remove the record of unplugged device block: Don't check QUEUE_FLAG_SAME_COMP in __blk_complete_request mm: Add comment explaining task state setting in bdi_forker_thread() mm: Cleanup clearing of BDI_pending bit in bdi_forker_thread() block: simplify force plug flush code a little bit block: change force plug flush call order block: Fix queue_flag update when rq_affinity goes from 2 to 1 block: separate priority boosting from REQ_META block: remove READ_META and WRITE_META xen-blkback: fixed indentation and comments xen-blkback: Don't disconnect backend until state switched to XenbusStateClosed.
2011-09-19Merge branch 'slab/urgent' of git://github.com/penberg/linuxLinus Torvalds
* 'slab/urgent' of git://github.com/penberg/linux: slub: add slab with one free object to partial list tail