- 03 Apr, 2014 40 commits
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Rashika Kheria authored
Mark function as static in mmap.c because they are not used outside this file. This eliminates the following warning in mm/mmap.c: mm/mmap.c:407:6: warning: no previous prototype for `validate_mm' [-Wmissing-prototypes] Signed-off-by:
Rashika Kheria <rashika.kheria@gmail.com> Reviewed-by:
Josh Triplett <josh@joshtriplett.org> Reviewed-by:
Rik van Riel <riel@redhat.com> Acked-by:
David Rientjes <rientjes@google.com> Signed-off-by:
Andrew Morton <akpm@linux-foundation.org> Signed-off-by:
Linus Torvalds <torvalds@linux-foundation.org>
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Rashika Kheria authored
mark functions as static in memory.c because they are not used outside this file. This eliminates the following warnings in mm/memory.c: mm/memory.c:3530:5: warning: no previous prototype for `numa_migrate_prep' [-Wmissing-prototypes] mm/memory.c:3545:5: warning: no previous prototype for `do_numa_page' [-Wmissing-prototypes] Signed-off-by:
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Rashika Kheria authored
Mark function as static in compaction.c because it is not used outside this file. This eliminates the following warning from mm/compaction.c: mm/compaction.c:1190:9: warning: no previous prototype for `sysfs_compact_node' [-Wmissing-prototypes Signed-off-by:
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David Rientjes authored
Page migration will fail for memory that is pinned in memory with, for example, get_user_pages(). In this case, it is unnecessary to take zone->lru_lock or isolating the page and passing it to page migration which will ultimately fail. This is a racy check, the page can still change from under us, but in that case we'll just fail later when attempting to move the page. This avoids very expensive memory compaction when faulting transparent hugepages after pinning a lot of memory with a Mellanox driver. On a 128GB machine and pinning ~120GB of memory, before this patch we see the enormous disparity in the number of page migration failures because of the pinning (from /proc/vmstat): compact_pages_moved 8450 compact_pagemigrate_failed 15614415 0.05% of pages isolated are successfully migrated and explicitly triggering memory compaction takes 102 seconds. After the patch: compact_pages_moved 9197 compact_pagemigrate_failed 7 99.9% of pages isolated are now successfully migrated in this configuration and memory compaction takes less than one second. Signed-off-by:
Hugh Dickins <hughd@google.com> Acked-by:
Mel Gorman <mgorman@suse.de> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Rik van Riel <riel@redhat.com> Cc: Greg Thelen <gthelen@google.com> Signed-off-by:
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David Rientjes authored
Both prep_compound_huge_page() and prep_compound_gigantic_page() are only called at bootstrap and can be marked as __init. The __SetPageTail(page) in prep_compound_gigantic_page() happening before page->first_page is initialized is not concerning since this is bootstrap. Signed-off-by:
Michal Hocko <mhocko@suse.cz> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Reviewed-by:
Davidlohr Bueso <davidlohr@hp.com> Signed-off-by:
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Johannes Weiner authored
Previously, page cache radix tree nodes were freed after reclaim emptied out their page pointers. But now reclaim stores shadow entries in their place, which are only reclaimed when the inodes themselves are reclaimed. This is problematic for bigger files that are still in use after they have a significant amount of their cache reclaimed, without any of those pages actually refaulting. The shadow entries will just sit there and waste memory. In the worst case, the shadow entries will accumulate until the machine runs out of memory. To get this under control, the VM will track radix tree nodes exclusively containing shadow entries on a per-NUMA node list. Per-NUMA rather than global because we expect the radix tree nodes themselves to be allocated node-locally and we want to reduce cross-node references of otherwise independent cache workloads. A simple shrinker will then reclaim these nodes on memory pressure. A few things need to be stored in the radix tree node to implement the shadow node LRU and allow tree deletions coming from the list: 1. There is no index available that would describe the reverse path from the node up to the tree root, which is needed to perform a deletion. To solve this, encode in each node its offset inside the parent. This can be stored in the unused upper bits of the same member that stores the node's height at no extra space cost. 2. The number of shadow entries needs to be counted in addition to the regular entries, to quickly detect when the node is ready to go to the shadow node LRU list. The current entry count is an unsigned int but the maximum number of entries is 64, so a shadow counter can easily be stored in the unused upper bits. 3. Tree modification needs tree lock and tree root, which are located in the address space, so store an address_space backpointer in the node. The parent pointer of the node is in a union with the 2-word rcu_head, so the backpointer comes at no extra cost as well. 4. The node needs to be linked to an LRU list, which requires a list head inside the node. This does increase the size of the node, but it does not change the number of objects that fit into a slab page. [akpm@linux-foundation.org: export the right function] Signed-off-by:
Johannes Weiner <hannes@cmpxchg.org> Reviewed-by:
Minchan Kim <minchan@kernel.org> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Bob Liu <bob.liu@oracle.com> Cc: Christoph Hellwig <hch@infradead.org> Cc: Dave Chinner <david@fromorbit.com> Cc: Greg Thelen <gthelen@google.com> Cc: Hugh Dickins <hughd@google.com> Cc: Jan Kara <jack@suse.cz> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Luigi Semenzato <semenzato@google.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Metin Doslu <metin@citusdata.com> Cc: Michel Lespinasse <walken@google.com> Cc: Ozgun Erdogan <ozgun@citusdata.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Roman Gushchin <klamm@yandex-team.ru> Cc: Ryan Mallon <rmallon@gmail.com> Cc: Tejun Heo <tj@kernel.org> Cc: Vlastimil Babka <vbabka@suse.cz> Signed-off-by:
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Johannes Weiner authored
Make struct radix_tree_node part of the public interface and provide API functions to create, look up, and delete whole nodes. Refactor the existing insert, look up, delete functions on top of these new node primitives. This will allow the VM to track and garbage collect page cache radix tree nodes. [sasha.levin@oracle.com: return correct error code on insertion failure] Signed-off-by:
Sasha Levin <sasha.levin@oracle.com> Signed-off-by:
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Johannes Weiner authored
The VM maintains cached filesystem pages on two types of lists. One list holds the pages recently faulted into the cache, the other list holds pages that have been referenced repeatedly on that first list. The idea is to prefer reclaiming young pages over those that have shown to benefit from caching in the past. We call the recently usedbut ultimately was not significantly better than a FIFO policy and still thrashed cache based on eviction speed, rather than actual demand for cache. This patch solves one half of the problem by decoupling the ability to detect working set changes from the inactive list size. By maintaining a history of recently evicted file pages it can detect frequently used pages with an arbitrarily small inactive list size, and subsequently apply pressure on the active list based on actual demand for cache, not just overall eviction speed. Every zone maintains a counter that tracks inactive list aging speed. When a page is evicted, a snapshot of this counter is stored in the now-empty page cache radix tree slot. On refault, the minimum access distance of the page can be assessed, to evaluate whether the page should be part of the active list or not. This fixes the VM's blindness towards working set changes in excess of the inactive list. And it's the foundation to further improve the protection ability and reduce the minimum inactive list size of 50%. Signed-off-by:
Bob Liu <bob.liu@oracle.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Christoph Hellwig <hch@infradead.org> Cc: Dave Chinner <david@fromorbit.com> Cc: Greg Thelen <gthelen@google.com> Cc: Hugh Dickins <hughd@google.com> Cc: Jan Kara <jack@suse.cz> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Luigi Semenzato <semenzato@google.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Metin Doslu <metin@citusdata.com> Cc: Michel Lespinasse <walken@google.com> Cc: Ozgun Erdogan <ozgun@citusdata.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Roman Gushchin <klamm@yandex-team.ru> Cc: Ryan Mallon <rmallon@gmail.com> Cc: Tejun Heo <tj@kernel.org> Cc: Vlastimil Babka <vbabka@suse.cz> Signed-off-by:
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Johannes Weiner authored
Reclaim will be leaving shadow entries in the page cache radix tree upon evicting the real page. As those pages are found from the LRU, an iput() can lead to the inode being freed concurrently. At this point, reclaim must no longer install shadow pages because the inode freeing code needs to ensure the page tree is really empty. Add an address_space flag, AS_EXITING, that the inode freeing code sets under the tree lock before doing the final truncate. Reclaim will check for this flag before installing shadow pages. Signed-off-by:
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Johannes Weiner authored
shmem mappings already contain exceptional entries where swap slot information is remembered. To be able to store eviction information for regular page cache, prepare every site dealing with the radix trees directly to handle entries other than pages. The common lookup functions will filter out non-page entries and return NULL for page cache holes, just as before. But provide a raw version of the API which returns non-page entries as well, and switch shmem over to use it. Signed-off-by:
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Johannes Weiner authored
The radix tree hole searching code is only used for page cache, for example the readahead code trying to get a a picture of the area surrounding a fault. It sufficed to rely on the radix tree definition of holes, which is "empty tree slot". But this is about to change, though, as shadow page descriptors will be stored in the page cache after the actual pages get evicted from memory. Move the functions over to mm/filemap.c and make them native page cache operations, where they can later be adapted to handle the new definition of "page cache hole". Signed-off-by:
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Johannes Weiner authored
Page cache radix tree slots are usually stabilized by the page lock, but shmem's swap cookies have no such thing. Because the overall truncation loop is lockless, the swap entry is currently confirmed by a tree lookup and then deleted by another tree lookup under the same tree lock region. Use radix_tree_delete_item() instead, which does the verification and deletion with only one lookup. This also allows removing the delete-only special case from shmem_radix_tree_replace(). Signed-off-by:
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Johannes Weiner authored
Provide a function that does not just delete an entry at a given index, but also allows passing in an expected item. Delete only if that item is still located at the specified index. This is handy when lockless tree traversals want to delete entries as well because they don't have to do an second, locked lookup to verify the slot has not changed under them before deleting the entry. Signed-off-by:
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Johannes Weiner authored
This code used to have its own lru cache pagevec up until a0b8cab3 ("mm: remove lru parameter from __pagevec_lru_add and remove parts of pagevec API"). Now it's just add_to_page_cache() followed by lru_cache_add(), might as well use add_to_page_cache_lru() directly. Signed-off-by:
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Johannes Weiner authored
Summary: The VM maintains cached filesystem pages on two types of lists. One list holds the pages recently faulted into the cache, the other list holds pages that have been referenced repeatedly on that first list. The idea is to prefer reclaiming young pages over those that have shown to benefit from caching in the past. We call the recently used list "inactive list" and the frequently used list "active list". Currently, the VM aims for a 1:1 ratio between the lists, which is the "perfect" trade-off between the ability to *protect* frequently used pages and the ability to *detect* frequently used pages. This means that working set changes bigger than half of cache memory go undetected and thrash indefinitely, whereas working sets bigger than half of cache memory are unprotected against used-once streams that don't even need caching. This happens on file servers and media streaming servers, where the popular files and file sections change over time. Even though the individual files might be smaller than half of memory, concurrent access to many of them may still result in their inter-reference distance being greater than half of memory. It's also been reported as a problem on database workloads that switch back and forth between tables that are bigger than half of memory. In these cases the VM never recognizes the new working set and will for the remainder of the workload thrash disk data which could easily live in memory. Historically, every reclaim scan of the inactive list also took a smaller number of pages from the tail of the active list and moved them to the head of the inactive list. This model gave established working sets more gracetime in the face of temporary use-once streams, but ultimately was not significantly better than a FIFO policy and still thrashed cache based on eviction speed, rather than actual demand for cache. This series solves the problem by maintaining a history of pages evicted from the inactive list, enabling the VM to detect frequently used pages regardless of inactive list size and facilitate working set transitions. Tests: The reported database workload is easily demonstrated on a 8G machine with two filesets a 6G. This fio workload operates on one set first, then switches to the other. The VM should obviously always cache the set that the workload is currently using. This test is based on a problem encountered by Citus Data customers: http://citusdata.com/blog/72-linux-memory-manager-and-your-big-data unpatched: db1: READ: io=98304MB, aggrb=885559KB/s, minb=885559KB/s, maxb=885559KB/s, mint= 113672msec, maxt= 113672msec db2: READ: io=98304MB, aggrb= 66169KB/s, minb= 66169KB/s, maxb= 66169KB/s, mint=1521302msec, maxt=1521302msec sdb: ios=835750/4, merge=2/1, ticks=4659739/60016, in_queue=4719203, util=98.92% real 27m15.541s user 0m19.059s sys 0m51.459s patched: db1: READ: io=98304MB, aggrb=877783KB/s, minb=877783KB/s, maxb=877783KB/s, mint=114679msec, maxt=114679msec db2: READ: io=98304MB, aggrb=397449KB/s, minb=397449KB/s, maxb=397449KB/s, mint=253273msec, maxt=253273msec sdb: ios=170587/4, merge=2/1, ticks=954910/61123, in_queue=1015923, util=90.40% real 6m8.630s user 0m14.714s sys 0m31.233s As can be seen, the unpatched kernel simply never adapts to the workingset change and db2 is stuck indefinitely with secondary storage speed. The patched kernel needs 2-3 iterations over db2 before it replaces db1 and reaches full memory speed. Given the unbounded negative affect of the existing VM behavior, these patches should be considered correctness fixes rather than performance optimizations. Another test resembles a fileserver or streaming server workload, where data in excess of memory size is accessed at different frequencies. There is very hot data accessed at a high frequency. Machines should be fitted so that the hot set of such a workload can be fully cached or all bets are off. Then there is a very big (compared to available memory) set of data that is used-once or at a very low frequency; this is what drives the inactive list and does not really benefit from caching. Lastly, there is a big set of warm data in between that is accessed at medium frequencies and benefits from caching the pages between the first and last streamer of each burst. unpatched: hot: READ: io=128000MB, aggrb=160693KB/s, minb=160693KB/s, maxb=160693KB/s, mint=815665msec, maxt=815665msec warm: READ: io= 81920MB, aggrb=109853KB/s, minb= 27463KB/s, maxb= 29244KB/s, mint=717110msec, maxt=763617msec cold: READ: io= 30720MB, aggrb= 35245KB/s, minb= 35245KB/s, maxb= 35245KB/s, mint=892530msec, maxt=892530msec sdb: ios=797960/4, merge=11763/1, ticks=4307910/796, in_queue=4308380, util=100.00% patched: hot: READ: io=128000MB, aggrb=160678KB/s, minb=160678KB/s, maxb=160678KB/s, mint=815740msec, maxt=815740msec warm: READ: io= 81920MB, aggrb=147747KB/s, minb= 36936KB/s, maxb= 40960KB/s, mint=512000msec, maxt=567767msec cold: READ: io= 30720MB, aggrb= 40960KB/s, minb= 40960KB/s, maxb= 40960KB/s, mint=768000msec, maxt=768000msec sdb: ios=596514/4, merge=9341/1, ticks=2395362/997, in_queue=2396484, util=79.18% In both kernels, the hot set is propagated to the active list and then served from cache. In both kernels, the beginning of the warm set is propagated to the active list as well, but in the unpatched case the active list eventually takes up half of memory and no new pages from the warm set get activated, despite repeated access, and despite most of the active list soon being stale. The patched kernel on the other hand detects the thrashing and manages to keep this cache window rolling through the data set. This frees up enough IO bandwidth that the cold set is served at full speed as well and disk utilization even drops by 20%. For reference, this same test was performed with the traditional demotion mechanism, where deactivation is coupled to inactive list reclaim. However, this had the same outcome as the unpatched kernel: while the warm set does indeed get activated continuously, it is forced out of the active list by inactive list pressure, which is dictated primarily by the unrelated cold set. The warm set is evicted before subsequent streamers can benefit from it, even though there would be enough space available to cache the pages of interest. Costs: Page reclaim used to shrink the radix trees but now the tree nodes are reused for shadow entries, where the cost depends heavily on the page cache access patterns. However, with workloads that maintain spatial or temporal locality, the shadow entries are either refaulted quickly or reclaimed along with the inode object itself. Workloads that will experience a memory cost increase are those that don't really benefit from caching in the first place. A more predictable alternative would be a fixed-cost separate pool of shadow entries, but this would incur relatively higher memory cost for well-behaved workloads at the benefit of cornercases. It would also make the shadow entry lookup more costly compared to storing them directly in the cache structure. Future: To simplify the merging process, this patch set is implementing thrash detection on a global per-zone level only for now, but the design is such that it can be extended to memory cgroups as well. All we need to do is store the unique cgroup ID along the node and zone identifier inside the eviction cookie to identify the lruvec. Right now we have a fixed ratio (50:50) between inactive and active list but we already have complaints about working sets exceeding half of memory being pushed out of the cache by simple streaming in the background. Ultimately, we want to adjust this ratio and allow for a much smaller inactive list. These patches are an essential step in this direction because they decouple the VMs ability to detect working set changes from the inactive list size. This would allow us to base the inactive list size on the combined readahead window size for example and potentially protect a much bigger working set. It's also a big step towards activating pages with a reuse distance larger than memory, as long as they are the most frequently used pages in the workload. This will require knowing more about the access frequency of active pages than what we measure right now, so it's also deferred in this series. Another possibility of having thrashing information would be to revisit the idea of local reclaim in the form of zero-config memory control groups. Instead of having allocating tasks go straight to global reclaim, they could try to reclaim the pages in the memcg they are part of first as long as the group is not thrashing. This would allow a user to drop e.g. a back-up job in an otherwise unconfigured memcg and it would only inflate (and possibly do global reclaim) until it has enough memory to do proper readahead. But once it reaches that point and stops thrashing it would just recycle its own used-once pages without kicking out the cache of any other tasks in the system more than necessary. This patch (of 10): Fengguang Wu's build testing spotted problems with inc_zone_state() and dec_zone_state() on UP configurations in out-of-tree patches. inc_zone_state() is declared but not defined, dec_zone_state() is missing entirely. Just like with *_zone_page_state(), they can be defined like their preemption-unsafe counterparts on UP. [akpm@linux-foundation.org: make it build] Signed-off-by:
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Vladimir Davydov authored
The name `max_pass' is misleading, because this variable actually keeps the estimate number of freeable objects, not the maximal number of objects we can scan in this pass, which can be twice that. Rename it to reflect its actual meaning. Signed-off-by:
Vladimir Davydov <vdavydov@parallels.com> Acked-by:
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Davidlohr Bueso authored
The kernel can currently only handle a single hugetlb page fault at a time. This is due to a single mutex that serializes the entire path. This lock protects from spurious OOM errors under conditions of low availability of free hugepages. This problem is specific to hugepages, because it is normal to want to use every single hugepage in the system - with normal pages we simply assume there will always be a few spare pages which can be used temporarily until the race is resolved. Address this problem by using a table of mutexes, allowing a better chance of parallelization, where each hugepage is individually serialized. The hash key is selected depending on the mapping type. For shared ones it consists of the address space and file offset being faulted; while for private ones the mm and virtual address are used. The size of the table is selected based on a compromise of collisions and memory footprint of a series of database workloads. Large database workloads that make heavy use of hugepages can be particularly exposed to this issue, causing start-up times to be painfully slow. This patch reduces the startup time of a 10 Gb Oracle DB (with ~5000 faults) from 37.5 secs to 25.7 secs. Larger workloads will naturally benefit even more. NOTE: The only downside to this patch, detected by Joonsoo Kim, is that a small race is possible in private mappings: A child process (with its own mm, after cow) can instantiate a page that is already being handled by the parent in a cow fault. When low on pages, can trigger spurious OOMs. I have not been able to think of a efficient way of handling this... but do we really care about such a tiny window? We already maintain another theoretical race with normal pages. If not, one possible way to is to maintain the single hash for private mappings -- any workloads that *really* suffer from this scaling problem should already use shared mappings. [akpm@linux-foundation.org: remove stray + characters, go BUG if hugetlb_init() kmalloc fails] Signed-off-by:
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Joonsoo Kim authored
Util now, we get a resv_map by two ways according to each mapping type. This makes code dirty and unreadable. Unify it. [davidlohr@hp.com: code cleanups] Signed-off-by:
Joonsoo Kim <iamjoonsoo.kim@lge.com> Signed-off-by:
Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Reviewed-by:
Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Cc: David Gibson <david@gibson.dropbear.id.au> Signed-off-by:
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Joonsoo Kim authored
This is a preparation patch to unify the use of vma_resv_map() regardless of the map type. This patch prepares it by removing resv_map_put(), which only works for HPAGE_RESV_OWNER's resv_map, not for all resv_maps. [davidlohr@hp.com: update changelog] Signed-off-by:
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Davidlohr Bueso authored
There is a race condition if we map a same file on different processes. Region tracking is protected by mmap_sem and hugetlb_instantiation_mutex. When we do mmap, we don't grab a hugetlb_instantiation_mutex, but only mmap_sem (exclusively). This doesn't prevent other tasks from modifying the region structure, so it can be modified by two processes concurrently. To solve this, introduce a spinlock to resv_map and make region manipulation function grab it before they do actual work. [davidlohr@hp.com: updated changelog] Signed-off-by:
David Gibson <david@gibson.dropbear.id.au> Cc: David Gibson <david@gibson.dropbear.id.au> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Signed-off-by:
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Joonsoo Kim authored
To change a protection method for region tracking to find grained one, we pass the resv_map, instead of list_head, to region manipulation functions. This doesn't introduce any functional change, and it is just for preparing a next step. [davidlohr@hp.com: update changelog] Signed-off-by:
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Joonsoo Kim authored
Currently, to track reserved and allocated regions, we use two different ways, depending on the mapping. For MAP_SHARED, we use address_mapping's private_list and, while for MAP_PRIVATE, we use a resv_map. Now, we are preparing to change a coarse grained lock which protect a region structure to fine grained lock, and this difference hinder it. So, before changing it, unify region structure handling, consistently using a resv_map regardless of the kind of mapping. Signed-off-by:
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Mel Gorman authored
Since put_mems_allowed() is strictly optional, its a seqcount retry, we don't need to evaluate the function if the allocation was in fact successful, saving a smp_rmb some loads and comparisons on some relative fast-paths. Since the naming, get/put_mems_allowed() does suggest a mandatory pairing, rename the interface, as suggested by Mel, to resemble the seqcount interface. This gives us: read_mems_allowed_begin() and read_mems_allowed_retry(), where it is important to note that the return value of the latter call is inverted from its previous incarnation. Signed-off-by:
Peter Zijlstra <a.p.zijlstra@chello.nl> Signed-off-by:
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David Rientjes authored
The cached pageblock hint should be ignored when triggering compaction through /proc/sys/vm/compact_memory so all eligible memory is isolated. Manually invoking compaction is known to be expensive, there's no need to skip pageblocks based on heuristics (mainly for debugging). Signed-off-by:
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Vladimir Davydov authored
There is no need passing on a shrink_control struct from try_to_free_pages() and friends to do_try_to_free_pages() and then to shrink_zones(), because it is only used in shrink_zones() and the only field initialized on the top level is gfp_mask, which is always equal to scan_control.gfp_mask. So let's move shrink_control initialization to shrink_zones(). Signed-off-by:
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Vladimir Davydov authored
This reduces the indentation level of do_try_to_free_pages() and removes extra loop over all eligible zones counting the number of on-LRU pages. Signed-off-by:
Glauber Costa <glommer@gmail.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Michal Hocko <mhocko@suse.cz> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Rik van Riel <riel@redhat.com> Cc: Dave Chinner <dchinner@redhat.com> Signed-off-by:
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Vladimir Davydov authored
When direct reclaim is executed by a process bound to a set of NUMA nodes, we should scan only those nodes when possible, but currently we will scan kmem from all online nodes even if the kmem shrinker is NUMA aware. That said, binding a process to a particular NUMA node won't prevent it from shrinking inode/dentry caches from other nodes, which is not good. Fix this. Signed-off-by:
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Ben Zhang authored
I ran into a scenario where while one cpu was stuck and should have panic'd because of the NMI watchdog, it didn't. The reason was another cpu was spewing stack dumps on to the console. Upon investigation, I noticed that when writing to the console and also when dumping the stack, the watchdog is touched. This causes all the cpus to reset their NMI watchdog flags and the 'stuck' cpu just spins forever. This change causes the semantics of touch_nmi_watchdog to be changed slightly. Previously, I accidentally changed the semantics and we noticed there was a codepath in which touch_nmi_watchdog could be touched from a preemtible area. That caused a BUG() to happen when CONFIG_DEBUG_PREEMPT was enabled. I believe it was the acpi code. My attempt here re-introduces the change to have the touch_nmi_watchdog() code only touch the local cpu instead of all of the cpus. But instead of using __get_cpu_var(), I use the __raw_get_cpu_var() version. This avoids the preemption problem. However my reasoning wasn't because I was trying to be lazy. Instead I rationalized it as, well if preemption is enabled then interrupts should be enabled to and the NMI watchdog will have no reason to trigger. So it won't matter if the wrong cpu is touched because the percpu interrupt counters the NMI watchdog uses should still be incrementing. Don said: : I'm ok with this patch, though it does alter the behaviour of how : touch_nmi_watchdog works. For the most part I don't think most callers : need to touch all of the watchdogs (on each cpu). Perhaps a corner case : will pop up (the scheduler?? to mimic touch_all_softlockup_watchdogs() ). : : But this does address an issue where if a system is locked up and one cpu : is spewing out useful debug messages (or error messages), the hard lockup : will fail to go off. We have seen this on RHEL also. Signed-off-by:
Don Zickus <dzickus@redhat.com> Signed-off-by:
Ben Zhang <benzh@chromium.org> Signed-off-by:
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Dan Carpenter authored
We know that "ret > 0" is true here. These tests were left over from commit 02afc27f ('direct-io: Handle O_(D)SYNC AIO') and aren't needed any more. Signed-off-by:
Dan Carpenter <dan.carpenter@oracle.com> Signed-off-by:
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Gu Zheng authored
The return value of bio_get_nr_vecs() cannot be bigger than BIO_MAX_PAGES, so we can remove redundant the comparison between nr_pages and BIO_MAX_PAGES. Signed-off-by:
Gu Zheng <guz.fnst@cn.fujitsu.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Reviewed-by:
Jeff Moyer <jmoyer@redhat.com> Signed-off-by:
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Wengang Wang authored
This patch fixes the following crash: kernel BUG at fs/ocfs2/uptodate.c:530! Modules linked in: ocfs2(F) ocfs2_dlmfs ocfs2_stack_o2cb ocfs2_dlm ocfs2_nodemanager ocfs2_stackglue configfs bridge xen_pciback xen_netback xen_blkback xen_gntalloc xen_gntdev xen_evtchn xenfs xen_privcmd sunrpc 8021q garp stp llc bonding be2iscsi iscsi_boot_sysfs bnx2i cnic uio cxgb4i cxgb4 cxgb3i libcxgbi cxgb3 mdio ib_iser rdma_cm ib_cm iw_cm ib_sa ib_mad ib_core ib_addr ipv6 iscsi_tcp libiscsi_tcp libiscsi scsi_transport_iscsi iTCO_wdt iTCO_vendor_support dcdbas coretemp freq_table mperf microcode pcspkr serio_raw bnx2 lpc_ich mfd_core i5k_amb i5000_edac edac_core e1000e sg shpchp ext4(F) jbd2(F) mbcache(F) dm_round_robin(F) sr_mod(F) cdrom(F) usb_storage(F) sd_mod(F) crc_t10dif(F) pata_acpi(F) ata_generic(F) ata_piix(F) mptsas(F) mptscsih(F) mptbase(F) scsi_transport_sas(F) radeon(F) ttm(F) drm_kms_helper(F) drm(F) hwmon(F) i2c_algo_bit(F) i2c_core(F) dm_multipath(F) dm_mirror(F) dm_region_hash(F) dm_log(F) dm_mod(F) CPU 5 Pid: 21303, comm: xattr-test Tainted: GF W 3.8.13-30.el6uek.x86_64 #2 Dell Inc. PowerEdge 1950/0M788G RIP: ocfs2_set_new_buffer_uptodate+0x51/0x60 [ocfs2] Process xattr-test (pid: 21303, threadinfo ffff880017aca000, task ffff880016a2c480) Call Trace: ocfs2_init_xattr_bucket+0x8a/0x120 [ocfs2] ocfs2_cp_xattr_bucket+0xbb/0x1b0 [ocfs2] ocfs2_extend_xattr_bucket+0x20a/0x2f0 [ocfs2] ocfs2_add_new_xattr_bucket+0x23e/0x4b0 [ocfs2] ocfs2_xattr_set_entry_index_block+0x13c/0x3d0 [ocfs2] ocfs2_xattr_block_set+0xf9/0x220 [ocfs2] __ocfs2_xattr_set_handle+0x118/0x710 [ocfs2] ocfs2_xattr_set+0x691/0x880 [ocfs2] ocfs2_xattr_user_set+0x46/0x50 [ocfs2] generic_setxattr+0x96/0xa0 __vfs_setxattr_noperm+0x7b/0x170 vfs_setxattr+0xbc/0xc0 setxattr+0xde/0x230 sys_fsetxattr+0xc6/0xf0 system_call_fastpath+0x16/0x1b Code: 41 80 0c 24 01 48 89 df e8 7d f0 ff ff 4c 89 e6 48 89 df e8 a2 fe ff ff 48 89 df e8 3a f0 ff ff 48 8b 1c 24 4c 8b 64 24 08 c9 c3 <0f> 0b eb fe 90 90 90 90 90 90 90 90 90 90 90 55 48 89 e5 66 66 RIP ocfs2_set_new_buffer_uptodate+0x51/0x60 [ocfs2] It hit the BUG_ON() in ocfs2_set_new_buffer_uptodate(): void ocfs2_set_new_buffer_uptodate(struct ocfs2_caching_info *ci, struct buffer_head *bh) { /* This should definitely *not* exist in our cache */ if (ocfs2_buffer_cached(ci, bh)) printk(KERN_ERR "bh->b_blocknr: %lu @ %p\n", bh->b_blocknr, bh); BUG_ON(ocfs2_buffer_cached(ci, bh)); set_buffer_uptodate(bh); ocfs2_metadata_cache_io_lock(ci); ocfs2_set_buffer_uptodate(ci, bh); ocfs2_metadata_cache_io_unlock(ci); } The problem here is: We cached a block, but the buffer_head got reused. When we are to pick up this block again, a new buffer_head created with UPTODATE flag cleared. ocfs2_buffer_uptodate() returned false since no UPTODATE is set on the buffer_head. so we set this block to cache as a NEW block, then it failed at asserting block is not in cache. The fix is to add a new parameter indicating the bucket is a new allocated or not to ocfs2_init_xattr_bucket(). ocfs2_init_xattr_bucket() assert block not cached accordingly. Signed-off-by:Wengang Wang <wen.gang.wang@oracle.com> Cc: Joel Becker <jlbec@evilplan.org> Reviewed-by:
Mark Fasheh <mfasheh@suse.de> Cc: Joe Jin <joe.jin@oracle.com> Signed-off-by:
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jiangyiwen authored
The following case may lead to the same system inode ref in confusion. A thread B thread ocfs2_get_system_file_inode ->get_local_system_inode ->_ocfs2_get_system_file_inode because of *arr == NULL, ocfs2_get_system_file_inode ->get_local_system_inode ->_ocfs2_get_system_file_inode gets first ref thru _ocfs2_get_system_file_inode, gets second ref thru igrab and set *arr = inode at the moment, B thread also gets two refs, so lead to one more inode ref. So add mutex lock to avoid multi thread set two inode ref once at the same time. Signed-off-by:jiangyiwen <jiangyiwen@huawei.com> Reviewed-by:
Joseph Qi <joseph.qi@huawei.com> Cc: Joel Becker <jlbec@evilplan.org> Cc: Mark Fasheh <mfasheh@suse.com> Signed-off-by:
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jiangyiwen authored
In ocfs2_info_handle_freeinode() and ocfs2_test_inode_bit() func, after calls ocfs2_get_system_file_inode() to get inode ref, if calls ocfs2_info_scan_inode_alloc() or ocfs2_inode_lock() failed, we should iput inode alloc to avoid leaking the inode. Signed-off-by:
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Tariq Saeed authored
Orabug: 17330860 When accepting an incomming connection o2net_accept_one clones a child data socket from the parent listening socket. It then proceeds to setup the child with callback o2net_data_ready() and sk_user_data to NULL. If data arrives in this window, o2net_listen_data_ready will be called with some non-deterministic value in sk_user_data (not inherited). We panic when we page fault on sk_user_data -- in parent it is sock_def_readable(). The fix is to recognize that this is a data socket being set up by looking at the socket state and do nothing. Signed-off-by:
Tariq Saseed <tariq.x.saeed@oracle.com> Signed-off-by:
Srinivas Eeda <srinivas.eeda@oracle.com> Reviewed-by:
Mark Fasheh <mfasheh@suse.com> Cc: Joel Becker <jlbec@evilplan.org> Signed-off-by:
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Younger Liu authored
After updating alloc_dinode counts in ocfs2_alloc_dinode_update_counts(), if ocfs2_alloc_dinode_update_bitmap() failed, there is a rare case that some space may be lost. So, roll back alloc_dinode counts when ocfs2_block_group_set_bits() failed. [akpm@linux-foundation.org: coding-style fixes] Signed-off-by:
Younger Liu <younger.liucn@gmail.com> Reviewed-by:
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Wengang Wang authored
ocfs2_do_flock() calls ocfs2_file_lock() to get the cross-node clock and then call flock_lock_file_wait() to compete with local processes. In case flock_lock_file_wait() failed, say -ENOMEM, clean up work is not done. This patch adds the cleanup --drop the cross-node lock which was just granted. [akpm@linux-foundation.org: coding-style fixes] Signed-off-by:
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Darrick J. Wong authored
Ensure that ocfs2_update_inode_fsync_trans() is called any time we touch an inode in a given transaction. This is a follow-on to the previous patch to reduce lock contention and deadlocking during an fsync operation. Signed-off-by:
Darrick J. Wong <darrick.wong@oracle.com> Cc: Mark Fasheh <mfasheh@suse.de> Cc: Joel Becker <jlbec@evilplan.org> Cc: Wengang <wen.gang.wang@oracle.com> Cc: Greg Marsden <greg.marsden@oracle.com> Cc: Srinivas Eeda <srinivas.eeda@oracle.com> Signed-off-by:
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Tetsuo Handa authored
Commit 9548906b ('xattr: Constify ->name member of "struct xattr"') missed that ocfs2 is calling kfree(xattr->name). As a result, kernel panic occurs upon calling kfree(xattr->name) because xattr->name refers static constant names. This patch removes kfree(xattr->name) from ocfs2_mknod() and ocfs2_symlink(). Signed-off-by:
Tetsuo Handa <penguin-kernel@I-love.SAKURA.ne.jp> Reported-by:
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alex chen authored
Do not put bh when buffer_uptodate failed in ocfs2_write_block and ocfs2_write_super_or_backup, because it will put bh in b_end_io. Otherwise it will hit a warning "VFS: brelse: Trying to free free buffer". Signed-off-by:
Alex Chen <alex.chen@huawei.com> Reviewed-by:
Joel Becker <jlbec@evilplan.org> Cc: <stable@vger.kernel.org> Signed-off-by:
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Xue jiufei authored
When ocfs2_create_new_inode_locks() return error, inode open lock may not be obtainted for this inode. So other nodes can remove this file and free dinode when inode still remain in memory on this node, which is not correct and may trigger BUG. So __ocfs2_mknod_locked should return error when ocfs2_create_new_inode_locks() failed. Node_1 Node_2 create fileA, call ocfs2_mknod() -> ocfs2_get_init_inode(), allocate inodeA -> ocfs2_claim_new_inode(), claim dinode(dinodeA) -> call ocfs2_create_new_inode_locks(), create open lock failed, return error -> __ocfs2_mknod_locked return success unlink fileA try open lock succeed, and free dinodeA create another file, call ocfs2_mknod() -> ocfs2_get_init_inode(), allocate inodeB -> ocfs2_claim_new_inode(), as Node_2 had freed dinodeA, so claim dinodeA and update generation for dinodeA call __ocfs2_drop_dl_inodes()->ocfs2_delete_inode() to free inodeA, and finally triggers BUG on(inode->i_generation != le32_to_cpu(fe->i_generation)) in function ocfs2_inode_lock_update(). Signed-off-by:joyce.xue <xuejiufei@huawei.com> Cc: Mark Fasheh <mfasheh@suse.com> Cc: Joel Becker <jlbec@evilplan.org> Signed-off-by:
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