![Vladimir Davydov](/assets/img/avatar_default.png)
Kmem accounting of memcg is unusable now, because it lacks slab shrinker support. That means when we hit the limit we will get ENOMEM w/o any chance to recover. What we should do then is to call shrink_slab, which would reclaim old inode/dentry caches from this cgroup. This is what this patch set is intended to do. Basically, it does two things. First, it introduces the notion of per-memcg slab shrinker. A shrinker that wants to reclaim objects per cgroup should mark itself as SHRINKER_MEMCG_AWARE. Then it will be passed the memory cgroup to scan from in shrink_control->memcg. For such shrinkers shrink_slab iterates over the whole cgroup subtree under the target cgroup and calls the shrinker for each kmem-active memory cgroup. Secondly, this patch set makes the list_lru structure per-memcg. It's done transparently to list_lru users - everything they have to do is to tell list_lru_init that they want memcg-aware list_lru. Then the list_lru will automatically distribute objects among per-memcg lists basing on which cgroup the object is accounted to. This way to make FS shrinkers (icache, dcache) memcg-aware we only need to make them use memcg-aware list_lru, and this is what this patch set does. As before, this patch set only enables per-memcg kmem reclaim when the pressure goes from memory.limit, not from memory.kmem.limit. Handling memory.kmem.limit is going to be tricky due to GFP_NOFS allocations, and it is still unclear whether we will have this knob in the unified hierarchy. This patch (of 9): NUMA aware slab shrinkers use the list_lru structure to distribute objects coming from different NUMA nodes to different lists. Whenever such a shrinker needs to count or scan objects from a particular node, it issues commands like this: count = list_lru_count_node(lru, sc->nid); freed = list_lru_walk_node(lru, sc->nid, isolate_func, isolate_arg, &sc->nr_to_scan); where sc is an instance of the shrink_control structure passed to it from vmscan. To simplify this, let's add special list_lru functions to be used by shrinkers, list_lru_shrink_count() and list_lru_shrink_walk(), which consolidate the nid and nr_to_scan arguments in the shrink_control structure. This will also allow us to avoid patching shrinkers that use list_lru when we make shrink_slab() per-memcg - all we will have to do is extend the shrink_control structure to include the target memcg and make list_lru_shrink_{count,walk} handle this appropriately. Signed-off-by: Vladimir Davydov <vdavydov@parallels.com> Suggested-by: Dave Chinner <david@fromorbit.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: Greg Thelen <gthelen@google.com> Cc: Glauber Costa <glommer@gmail.com> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Tejun Heo <tj@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
415 lines
14 KiB
C
415 lines
14 KiB
C
/*
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* Workingset detection
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*
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* Copyright (C) 2013 Red Hat, Inc., Johannes Weiner
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*/
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#include <linux/memcontrol.h>
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#include <linux/writeback.h>
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#include <linux/pagemap.h>
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#include <linux/atomic.h>
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#include <linux/module.h>
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#include <linux/swap.h>
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#include <linux/fs.h>
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#include <linux/mm.h>
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/*
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* Double CLOCK lists
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*
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* Per zone, two clock lists are maintained for file pages: the
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* inactive and the active list. Freshly faulted pages start out at
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* the head of the inactive list and page reclaim scans pages from the
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* tail. Pages that are accessed multiple times on the inactive list
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* are promoted to the active list, to protect them from reclaim,
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* whereas active pages are demoted to the inactive list when the
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* active list grows too big.
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*
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* fault ------------------------+
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* |
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* +--------------+ | +-------------+
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* reclaim <- | inactive | <-+-- demotion | active | <--+
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* +--------------+ +-------------+ |
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* | |
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* +-------------- promotion ------------------+
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*
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*
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* Access frequency and refault distance
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*
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* A workload is thrashing when its pages are frequently used but they
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* are evicted from the inactive list every time before another access
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* would have promoted them to the active list.
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*
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* In cases where the average access distance between thrashing pages
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* is bigger than the size of memory there is nothing that can be
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* done - the thrashing set could never fit into memory under any
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* circumstance.
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*
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* However, the average access distance could be bigger than the
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* inactive list, yet smaller than the size of memory. In this case,
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* the set could fit into memory if it weren't for the currently
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* active pages - which may be used more, hopefully less frequently:
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*
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* +-memory available to cache-+
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* | |
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* +-inactive------+-active----+
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* a b | c d e f g h i | J K L M N |
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* +---------------+-----------+
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*
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* It is prohibitively expensive to accurately track access frequency
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* of pages. But a reasonable approximation can be made to measure
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* thrashing on the inactive list, after which refaulting pages can be
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* activated optimistically to compete with the existing active pages.
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*
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* Approximating inactive page access frequency - Observations:
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*
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* 1. When a page is accessed for the first time, it is added to the
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* head of the inactive list, slides every existing inactive page
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* towards the tail by one slot, and pushes the current tail page
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* out of memory.
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*
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* 2. When a page is accessed for the second time, it is promoted to
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* the active list, shrinking the inactive list by one slot. This
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* also slides all inactive pages that were faulted into the cache
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* more recently than the activated page towards the tail of the
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* inactive list.
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*
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* Thus:
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*
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* 1. The sum of evictions and activations between any two points in
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* time indicate the minimum number of inactive pages accessed in
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* between.
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*
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* 2. Moving one inactive page N page slots towards the tail of the
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* list requires at least N inactive page accesses.
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*
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* Combining these:
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*
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* 1. When a page is finally evicted from memory, the number of
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* inactive pages accessed while the page was in cache is at least
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* the number of page slots on the inactive list.
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*
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* 2. In addition, measuring the sum of evictions and activations (E)
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* at the time of a page's eviction, and comparing it to another
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* reading (R) at the time the page faults back into memory tells
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* the minimum number of accesses while the page was not cached.
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* This is called the refault distance.
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*
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* Because the first access of the page was the fault and the second
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* access the refault, we combine the in-cache distance with the
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* out-of-cache distance to get the complete minimum access distance
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* of this page:
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*
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* NR_inactive + (R - E)
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*
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* And knowing the minimum access distance of a page, we can easily
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* tell if the page would be able to stay in cache assuming all page
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* slots in the cache were available:
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*
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* NR_inactive + (R - E) <= NR_inactive + NR_active
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*
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* which can be further simplified to
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*
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* (R - E) <= NR_active
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*
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* Put into words, the refault distance (out-of-cache) can be seen as
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* a deficit in inactive list space (in-cache). If the inactive list
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* had (R - E) more page slots, the page would not have been evicted
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* in between accesses, but activated instead. And on a full system,
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* the only thing eating into inactive list space is active pages.
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*
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*
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* Activating refaulting pages
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*
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* All that is known about the active list is that the pages have been
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* accessed more than once in the past. This means that at any given
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* time there is actually a good chance that pages on the active list
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* are no longer in active use.
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*
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* So when a refault distance of (R - E) is observed and there are at
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* least (R - E) active pages, the refaulting page is activated
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* optimistically in the hope that (R - E) active pages are actually
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* used less frequently than the refaulting page - or even not used at
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* all anymore.
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*
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* If this is wrong and demotion kicks in, the pages which are truly
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* used more frequently will be reactivated while the less frequently
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* used once will be evicted from memory.
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*
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* But if this is right, the stale pages will be pushed out of memory
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* and the used pages get to stay in cache.
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*
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*
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* Implementation
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*
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* For each zone's file LRU lists, a counter for inactive evictions
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* and activations is maintained (zone->inactive_age).
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*
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* On eviction, a snapshot of this counter (along with some bits to
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* identify the zone) is stored in the now empty page cache radix tree
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* slot of the evicted page. This is called a shadow entry.
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*
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* On cache misses for which there are shadow entries, an eligible
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* refault distance will immediately activate the refaulting page.
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*/
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static void *pack_shadow(unsigned long eviction, struct zone *zone)
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{
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eviction = (eviction << NODES_SHIFT) | zone_to_nid(zone);
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eviction = (eviction << ZONES_SHIFT) | zone_idx(zone);
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eviction = (eviction << RADIX_TREE_EXCEPTIONAL_SHIFT);
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return (void *)(eviction | RADIX_TREE_EXCEPTIONAL_ENTRY);
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}
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static void unpack_shadow(void *shadow,
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struct zone **zone,
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unsigned long *distance)
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{
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unsigned long entry = (unsigned long)shadow;
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unsigned long eviction;
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unsigned long refault;
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unsigned long mask;
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int zid, nid;
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entry >>= RADIX_TREE_EXCEPTIONAL_SHIFT;
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zid = entry & ((1UL << ZONES_SHIFT) - 1);
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entry >>= ZONES_SHIFT;
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nid = entry & ((1UL << NODES_SHIFT) - 1);
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entry >>= NODES_SHIFT;
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eviction = entry;
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*zone = NODE_DATA(nid)->node_zones + zid;
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refault = atomic_long_read(&(*zone)->inactive_age);
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mask = ~0UL >> (NODES_SHIFT + ZONES_SHIFT +
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RADIX_TREE_EXCEPTIONAL_SHIFT);
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/*
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* The unsigned subtraction here gives an accurate distance
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* across inactive_age overflows in most cases.
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*
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* There is a special case: usually, shadow entries have a
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* short lifetime and are either refaulted or reclaimed along
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* with the inode before they get too old. But it is not
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* impossible for the inactive_age to lap a shadow entry in
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* the field, which can then can result in a false small
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* refault distance, leading to a false activation should this
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* old entry actually refault again. However, earlier kernels
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* used to deactivate unconditionally with *every* reclaim
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* invocation for the longest time, so the occasional
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* inappropriate activation leading to pressure on the active
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* list is not a problem.
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*/
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*distance = (refault - eviction) & mask;
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}
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/**
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* workingset_eviction - note the eviction of a page from memory
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* @mapping: address space the page was backing
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* @page: the page being evicted
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*
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* Returns a shadow entry to be stored in @mapping->page_tree in place
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* of the evicted @page so that a later refault can be detected.
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*/
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void *workingset_eviction(struct address_space *mapping, struct page *page)
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{
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struct zone *zone = page_zone(page);
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unsigned long eviction;
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eviction = atomic_long_inc_return(&zone->inactive_age);
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return pack_shadow(eviction, zone);
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}
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/**
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* workingset_refault - evaluate the refault of a previously evicted page
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* @shadow: shadow entry of the evicted page
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*
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* Calculates and evaluates the refault distance of the previously
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* evicted page in the context of the zone it was allocated in.
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*
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* Returns %true if the page should be activated, %false otherwise.
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*/
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bool workingset_refault(void *shadow)
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{
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unsigned long refault_distance;
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struct zone *zone;
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unpack_shadow(shadow, &zone, &refault_distance);
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inc_zone_state(zone, WORKINGSET_REFAULT);
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if (refault_distance <= zone_page_state(zone, NR_ACTIVE_FILE)) {
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inc_zone_state(zone, WORKINGSET_ACTIVATE);
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return true;
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}
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return false;
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}
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/**
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* workingset_activation - note a page activation
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* @page: page that is being activated
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*/
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void workingset_activation(struct page *page)
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{
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atomic_long_inc(&page_zone(page)->inactive_age);
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}
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/*
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* Shadow entries reflect the share of the working set that does not
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* fit into memory, so their number depends on the access pattern of
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* the workload. In most cases, they will refault or get reclaimed
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* along with the inode, but a (malicious) workload that streams
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* through files with a total size several times that of available
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* memory, while preventing the inodes from being reclaimed, can
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* create excessive amounts of shadow nodes. To keep a lid on this,
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* track shadow nodes and reclaim them when they grow way past the
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* point where they would still be useful.
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*/
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struct list_lru workingset_shadow_nodes;
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static unsigned long count_shadow_nodes(struct shrinker *shrinker,
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struct shrink_control *sc)
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{
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unsigned long shadow_nodes;
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unsigned long max_nodes;
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unsigned long pages;
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/* list_lru lock nests inside IRQ-safe mapping->tree_lock */
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local_irq_disable();
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shadow_nodes = list_lru_shrink_count(&workingset_shadow_nodes, sc);
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local_irq_enable();
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pages = node_present_pages(sc->nid);
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/*
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* Active cache pages are limited to 50% of memory, and shadow
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* entries that represent a refault distance bigger than that
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* do not have any effect. Limit the number of shadow nodes
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* such that shadow entries do not exceed the number of active
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* cache pages, assuming a worst-case node population density
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* of 1/8th on average.
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*
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* On 64-bit with 7 radix_tree_nodes per page and 64 slots
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* each, this will reclaim shadow entries when they consume
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* ~2% of available memory:
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*
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* PAGE_SIZE / radix_tree_nodes / node_entries / PAGE_SIZE
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*/
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max_nodes = pages >> (1 + RADIX_TREE_MAP_SHIFT - 3);
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if (shadow_nodes <= max_nodes)
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return 0;
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return shadow_nodes - max_nodes;
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}
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static enum lru_status shadow_lru_isolate(struct list_head *item,
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spinlock_t *lru_lock,
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void *arg)
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{
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struct address_space *mapping;
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struct radix_tree_node *node;
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unsigned int i;
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int ret;
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/*
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* Page cache insertions and deletions synchroneously maintain
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* the shadow node LRU under the mapping->tree_lock and the
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* lru_lock. Because the page cache tree is emptied before
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* the inode can be destroyed, holding the lru_lock pins any
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* address_space that has radix tree nodes on the LRU.
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*
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* We can then safely transition to the mapping->tree_lock to
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* pin only the address_space of the particular node we want
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* to reclaim, take the node off-LRU, and drop the lru_lock.
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*/
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node = container_of(item, struct radix_tree_node, private_list);
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mapping = node->private_data;
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/* Coming from the list, invert the lock order */
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if (!spin_trylock(&mapping->tree_lock)) {
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spin_unlock(lru_lock);
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ret = LRU_RETRY;
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goto out;
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}
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list_del_init(item);
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spin_unlock(lru_lock);
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/*
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* The nodes should only contain one or more shadow entries,
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* no pages, so we expect to be able to remove them all and
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* delete and free the empty node afterwards.
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*/
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BUG_ON(!node->count);
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BUG_ON(node->count & RADIX_TREE_COUNT_MASK);
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for (i = 0; i < RADIX_TREE_MAP_SIZE; i++) {
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if (node->slots[i]) {
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BUG_ON(!radix_tree_exceptional_entry(node->slots[i]));
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node->slots[i] = NULL;
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BUG_ON(node->count < (1U << RADIX_TREE_COUNT_SHIFT));
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node->count -= 1U << RADIX_TREE_COUNT_SHIFT;
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BUG_ON(!mapping->nrshadows);
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mapping->nrshadows--;
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}
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}
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BUG_ON(node->count);
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inc_zone_state(page_zone(virt_to_page(node)), WORKINGSET_NODERECLAIM);
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if (!__radix_tree_delete_node(&mapping->page_tree, node))
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BUG();
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spin_unlock(&mapping->tree_lock);
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ret = LRU_REMOVED_RETRY;
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out:
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local_irq_enable();
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cond_resched();
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local_irq_disable();
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spin_lock(lru_lock);
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return ret;
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}
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static unsigned long scan_shadow_nodes(struct shrinker *shrinker,
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struct shrink_control *sc)
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{
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unsigned long ret;
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/* list_lru lock nests inside IRQ-safe mapping->tree_lock */
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local_irq_disable();
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ret = list_lru_shrink_walk(&workingset_shadow_nodes, sc,
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shadow_lru_isolate, NULL);
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local_irq_enable();
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return ret;
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}
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static struct shrinker workingset_shadow_shrinker = {
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.count_objects = count_shadow_nodes,
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.scan_objects = scan_shadow_nodes,
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.seeks = DEFAULT_SEEKS,
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.flags = SHRINKER_NUMA_AWARE,
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};
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/*
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* Our list_lru->lock is IRQ-safe as it nests inside the IRQ-safe
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* mapping->tree_lock.
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*/
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static struct lock_class_key shadow_nodes_key;
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static int __init workingset_init(void)
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{
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int ret;
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ret = list_lru_init_key(&workingset_shadow_nodes, &shadow_nodes_key);
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if (ret)
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goto err;
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ret = register_shrinker(&workingset_shadow_shrinker);
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if (ret)
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goto err_list_lru;
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return 0;
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err_list_lru:
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list_lru_destroy(&workingset_shadow_nodes);
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err:
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return ret;
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}
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module_init(workingset_init);
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