blob: aa017133744b227bed7592ea6cc32f360c3e142c [file] [log] [blame]
* Workingset detection
* Copyright (C) 2013 Red Hat, Inc., Johannes Weiner
#include <linux/memcontrol.h>
#include <linux/writeback.h>
#include <linux/pagemap.h>
#include <linux/atomic.h>
#include <linux/module.h>
#include <linux/swap.h>
#include <linux/fs.h>
#include <linux/mm.h>
* Double CLOCK lists
* Per zone, two clock lists are maintained for file pages: the
* inactive and the active list. Freshly faulted pages start out at
* the head of the inactive list and page reclaim scans pages from the
* tail. Pages that are accessed multiple times on the inactive list
* are promoted to the active list, to protect them from reclaim,
* whereas active pages are demoted to the inactive list when the
* active list grows too big.
* fault ------------------------+
* |
* +--------------+ | +-------------+
* reclaim <- | inactive | <-+-- demotion | active | <--+
* +--------------+ +-------------+ |
* | |
* +-------------- promotion ------------------+
* Access frequency and refault distance
* A workload is thrashing when its pages are frequently used but they
* are evicted from the inactive list every time before another access
* would have promoted them to the active list.
* In cases where the average access distance between thrashing pages
* is bigger than the size of memory there is nothing that can be
* done - the thrashing set could never fit into memory under any
* circumstance.
* However, the average access distance could be bigger than the
* inactive list, yet smaller than the size of memory. In this case,
* the set could fit into memory if it weren't for the currently
* active pages - which may be used more, hopefully less frequently:
* +-memory available to cache-+
* | |
* +-inactive------+-active----+
* a b | c d e f g h i | J K L M N |
* +---------------+-----------+
* It is prohibitively expensive to accurately track access frequency
* of pages. But a reasonable approximation can be made to measure
* thrashing on the inactive list, after which refaulting pages can be
* activated optimistically to compete with the existing active pages.
* Approximating inactive page access frequency - Observations:
* 1. When a page is accessed for the first time, it is added to the
* head of the inactive list, slides every existing inactive page
* towards the tail by one slot, and pushes the current tail page
* out of memory.
* 2. When a page is accessed for the second time, it is promoted to
* the active list, shrinking the inactive list by one slot. This
* also slides all inactive pages that were faulted into the cache
* more recently than the activated page towards the tail of the
* inactive list.
* Thus:
* 1. The sum of evictions and activations between any two points in
* time indicate the minimum number of inactive pages accessed in
* between.
* 2. Moving one inactive page N page slots towards the tail of the
* list requires at least N inactive page accesses.
* Combining these:
* 1. When a page is finally evicted from memory, the number of
* inactive pages accessed while the page was in cache is at least
* the number of page slots on the inactive list.
* 2. In addition, measuring the sum of evictions and activations (E)
* at the time of a page's eviction, and comparing it to another
* reading (R) at the time the page faults back into memory tells
* the minimum number of accesses while the page was not cached.
* This is called the refault distance.
* Because the first access of the page was the fault and the second
* access the refault, we combine the in-cache distance with the
* out-of-cache distance to get the complete minimum access distance
* of this page:
* NR_inactive + (R - E)
* And knowing the minimum access distance of a page, we can easily
* tell if the page would be able to stay in cache assuming all page
* slots in the cache were available:
* NR_inactive + (R - E) <= NR_inactive + NR_active
* which can be further simplified to
* (R - E) <= NR_active
* Put into words, the refault distance (out-of-cache) can be seen as
* a deficit in inactive list space (in-cache). If the inactive list
* had (R - E) more page slots, the page would not have been evicted
* in between accesses, but activated instead. And on a full system,
* the only thing eating into inactive list space is active pages.
* Activating refaulting pages
* All that is known about the active list is that the pages have been
* accessed more than once in the past. This means that at any given
* time there is actually a good chance that pages on the active list
* are no longer in active use.
* So when a refault distance of (R - E) is observed and there are at
* least (R - E) active pages, the refaulting page is activated
* optimistically in the hope that (R - E) active pages are actually
* used less frequently than the refaulting page - or even not used at
* all anymore.
* If this is wrong and demotion kicks in, the pages which are truly
* used more frequently will be reactivated while the less frequently
* used once will be evicted from memory.
* But if this is right, the stale pages will be pushed out of memory
* and the used pages get to stay in cache.
* Implementation
* For each zone's file LRU lists, a counter for inactive evictions
* and activations is maintained (zone->inactive_age).
* On eviction, a snapshot of this counter (along with some bits to
* identify the zone) is stored in the now empty page cache radix tree
* slot of the evicted page. This is called a shadow entry.
* On cache misses for which there are shadow entries, an eligible
* refault distance will immediately activate the refaulting page.
static void *pack_shadow(unsigned long eviction, struct zone *zone)
eviction = (eviction << NODES_SHIFT) | zone_to_nid(zone);
eviction = (eviction << ZONES_SHIFT) | zone_idx(zone);
eviction = (eviction << RADIX_TREE_EXCEPTIONAL_SHIFT);
return (void *)(eviction | RADIX_TREE_EXCEPTIONAL_ENTRY);
static void unpack_shadow(void *shadow,
struct zone **zone,
unsigned long *distance)
unsigned long entry = (unsigned long)shadow;
unsigned long eviction;
unsigned long refault;
unsigned long mask;
int zid, nid;
zid = entry & ((1UL << ZONES_SHIFT) - 1);
entry >>= ZONES_SHIFT;
nid = entry & ((1UL << NODES_SHIFT) - 1);
entry >>= NODES_SHIFT;
eviction = entry;
*zone = NODE_DATA(nid)->node_zones + zid;
refault = atomic_long_read(&(*zone)->inactive_age);
* The unsigned subtraction here gives an accurate distance
* across inactive_age overflows in most cases.
* There is a special case: usually, shadow entries have a
* short lifetime and are either refaulted or reclaimed along
* with the inode before they get too old. But it is not
* impossible for the inactive_age to lap a shadow entry in
* the field, which can then can result in a false small
* refault distance, leading to a false activation should this
* old entry actually refault again. However, earlier kernels
* used to deactivate unconditionally with *every* reclaim
* invocation for the longest time, so the occasional
* inappropriate activation leading to pressure on the active
* list is not a problem.
*distance = (refault - eviction) & mask;
* workingset_eviction - note the eviction of a page from memory
* @mapping: address space the page was backing
* @page: the page being evicted
* Returns a shadow entry to be stored in @mapping->page_tree in place
* of the evicted @page so that a later refault can be detected.
void *workingset_eviction(struct address_space *mapping, struct page *page)
struct zone *zone = page_zone(page);
unsigned long eviction;
eviction = atomic_long_inc_return(&zone->inactive_age);
return pack_shadow(eviction, zone);
* workingset_refault - evaluate the refault of a previously evicted page
* @shadow: shadow entry of the evicted page
* Calculates and evaluates the refault distance of the previously
* evicted page in the context of the zone it was allocated in.
* Returns %true if the page should be activated, %false otherwise.
bool workingset_refault(void *shadow)
unsigned long refault_distance;
struct zone *zone;
unpack_shadow(shadow, &zone, &refault_distance);
inc_zone_state(zone, WORKINGSET_REFAULT);
if (refault_distance <= zone_page_state(zone, NR_ACTIVE_FILE)) {
inc_zone_state(zone, WORKINGSET_ACTIVATE);
return true;
return false;
* workingset_activation - note a page activation
* @page: page that is being activated
void workingset_activation(struct page *page)
* Shadow entries reflect the share of the working set that does not
* fit into memory, so their number depends on the access pattern of
* the workload. In most cases, they will refault or get reclaimed
* along with the inode, but a (malicious) workload that streams
* through files with a total size several times that of available
* memory, while preventing the inodes from being reclaimed, can
* create excessive amounts of shadow nodes. To keep a lid on this,
* track shadow nodes and reclaim them when they grow way past the
* point where they would still be useful.
struct list_lru workingset_shadow_nodes;
static unsigned long count_shadow_nodes(struct shrinker *shrinker,
struct shrink_control *sc)
unsigned long shadow_nodes;
unsigned long max_nodes;
unsigned long pages;
/* list_lru lock nests inside IRQ-safe mapping->tree_lock */
shadow_nodes = list_lru_shrink_count(&workingset_shadow_nodes, sc);
pages = node_present_pages(sc->nid);
* Active cache pages are limited to 50% of memory, and shadow
* entries that represent a refault distance bigger than that
* do not have any effect. Limit the number of shadow nodes
* such that shadow entries do not exceed the number of active
* cache pages, assuming a worst-case node population density
* of 1/8th on average.
* On 64-bit with 7 radix_tree_nodes per page and 64 slots
* each, this will reclaim shadow entries when they consume
* ~2% of available memory:
* PAGE_SIZE / radix_tree_nodes / node_entries / PAGE_SIZE
max_nodes = pages >> (1 + RADIX_TREE_MAP_SHIFT - 3);
if (shadow_nodes <= max_nodes)
return 0;
return shadow_nodes - max_nodes;
static enum lru_status shadow_lru_isolate(struct list_head *item,
struct list_lru_one *lru,
spinlock_t *lru_lock,
void *arg)
struct address_space *mapping;
struct radix_tree_node *node;
unsigned int i;
int ret;
* Page cache insertions and deletions synchroneously maintain
* the shadow node LRU under the mapping->tree_lock and the
* lru_lock. Because the page cache tree is emptied before
* the inode can be destroyed, holding the lru_lock pins any
* address_space that has radix tree nodes on the LRU.
* We can then safely transition to the mapping->tree_lock to
* pin only the address_space of the particular node we want
* to reclaim, take the node off-LRU, and drop the lru_lock.
node = container_of(item, struct radix_tree_node, private_list);
mapping = node->private_data;
/* Coming from the list, invert the lock order */
if (!spin_trylock(&mapping->tree_lock)) {
ret = LRU_RETRY;
goto out;
list_lru_isolate(lru, item);
* The nodes should only contain one or more shadow entries,
* no pages, so we expect to be able to remove them all and
* delete and free the empty node afterwards.
for (i = 0; i < RADIX_TREE_MAP_SIZE; i++) {
if (node->slots[i]) {
node->slots[i] = NULL;
BUG_ON(node->count < (1U << RADIX_TREE_COUNT_SHIFT));
node->count -= 1U << RADIX_TREE_COUNT_SHIFT;
inc_zone_state(page_zone(virt_to_page(node)), WORKINGSET_NODERECLAIM);
if (!__radix_tree_delete_node(&mapping->page_tree, node))
return ret;
static unsigned long scan_shadow_nodes(struct shrinker *shrinker,
struct shrink_control *sc)
unsigned long ret;
/* list_lru lock nests inside IRQ-safe mapping->tree_lock */
ret = list_lru_shrink_walk(&workingset_shadow_nodes, sc,
shadow_lru_isolate, NULL);
return ret;
static struct shrinker workingset_shadow_shrinker = {
.count_objects = count_shadow_nodes,
.scan_objects = scan_shadow_nodes,
* Our list_lru->lock is IRQ-safe as it nests inside the IRQ-safe
* mapping->tree_lock.
static struct lock_class_key shadow_nodes_key;
static int __init workingset_init(void)
int ret;
ret = list_lru_init_key(&workingset_shadow_nodes, &shadow_nodes_key);
if (ret)
goto err;
ret = register_shrinker(&workingset_shadow_shrinker);
if (ret)
goto err_list_lru;
return 0;
return ret;