blob: 8c03afe1d67cb95d6aecb248e522640115e0ee0d [file] [log] [blame]
// SPDX-License-Identifier: GPL-2.0
* Workingset detection
* Copyright (C) 2013 Red Hat, Inc., Johannes Weiner
#include <linux/memcontrol.h>
#include <linux/mm_inline.h>
#include <linux/writeback.h>
#include <linux/shmem_fs.h>
#include <linux/pagemap.h>
#include <linux/atomic.h>
#include <linux/module.h>
#include <linux/swap.h>
#include <linux/dax.h>
#include <linux/fs.h>
#include <linux/mm.h>
* Double CLOCK lists
* Per node, 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.
* Refaulting inactive 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.
* That means if inactive cache is refaulting with a suitable refault
* distance, we assume the cache workingset is transitioning and put
* pressure on the current active list.
* 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.
* Refaulting active pages
* If on the other hand the refaulting pages have recently been
* deactivated, it means that the active list is no longer protecting
* actively used cache from reclaim. The cache is NOT transitioning to
* a different workingset; the existing workingset is thrashing in the
* space allocated to the page cache.
* Implementation
* For each node's LRU lists, a counter for inactive evictions and
* activations is maintained (node->nonresident_age).
* On eviction, a snapshot of this counter (along with some bits to
* identify the node) is stored in the now empty page cache
* 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.
* Eviction timestamps need to be able to cover the full range of
* actionable refaults. However, bits are tight in the xarray
* entry, and after storing the identifier for the lruvec there might
* not be enough left to represent every single actionable refault. In
* that case, we have to sacrifice granularity for distance, and group
* evictions into coarser buckets by shaving off lower timestamp bits.
static unsigned int bucket_order __read_mostly;
static void *pack_shadow(int memcgid, pg_data_t *pgdat, unsigned long eviction,
bool workingset)
eviction >>= bucket_order;
eviction &= EVICTION_MASK;
eviction = (eviction << MEM_CGROUP_ID_SHIFT) | memcgid;
eviction = (eviction << NODES_SHIFT) | pgdat->node_id;
eviction = (eviction << WORKINGSET_SHIFT) | workingset;
return xa_mk_value(eviction);
static void unpack_shadow(void *shadow, int *memcgidp, pg_data_t **pgdat,
unsigned long *evictionp, bool *workingsetp)
unsigned long entry = xa_to_value(shadow);
int memcgid, nid;
bool workingset;
workingset = entry & ((1UL << WORKINGSET_SHIFT) - 1);
nid = entry & ((1UL << NODES_SHIFT) - 1);
entry >>= NODES_SHIFT;
memcgid = entry & ((1UL << MEM_CGROUP_ID_SHIFT) - 1);
*memcgidp = memcgid;
*pgdat = NODE_DATA(nid);
*evictionp = entry << bucket_order;
*workingsetp = workingset;
* workingset_age_nonresident - age non-resident entries as LRU ages
* @lruvec: the lruvec that was aged
* @nr_pages: the number of pages to count
* As in-memory pages are aged, non-resident pages need to be aged as
* well, in order for the refault distances later on to be comparable
* to the in-memory dimensions. This function allows reclaim and LRU
* operations to drive the non-resident aging along in parallel.
void workingset_age_nonresident(struct lruvec *lruvec, unsigned long nr_pages)
* Reclaiming a cgroup means reclaiming all its children in a
* round-robin fashion. That means that each cgroup has an LRU
* order that is composed of the LRU orders of its child
* cgroups; and every page has an LRU position not just in the
* cgroup that owns it, but in all of that group's ancestors.
* So when the physical inactive list of a leaf cgroup ages,
* the virtual inactive lists of all its parents, including
* the root cgroup's, age as well.
do {
atomic_long_add(nr_pages, &lruvec->nonresident_age);
} while ((lruvec = parent_lruvec(lruvec)));
* workingset_eviction - note the eviction of a page from memory
* @target_memcg: the cgroup that is causing the reclaim
* @page: the page being evicted
* Return: a shadow entry to be stored in @page->mapping->i_pages in place
* of the evicted @page so that a later refault can be detected.
void *workingset_eviction(struct page *page, struct mem_cgroup *target_memcg)
struct pglist_data *pgdat = page_pgdat(page);
unsigned long eviction;
struct lruvec *lruvec;
int memcgid;
/* Page is fully exclusive and pins page's memory cgroup pointer */
VM_BUG_ON_PAGE(PageLRU(page), page);
VM_BUG_ON_PAGE(page_count(page), page);
VM_BUG_ON_PAGE(!PageLocked(page), page);
lruvec = mem_cgroup_lruvec(target_memcg, pgdat);
/* XXX: target_memcg can be NULL, go through lruvec */
memcgid = mem_cgroup_id(lruvec_memcg(lruvec));
eviction = atomic_long_read(&lruvec->nonresident_age);
workingset_age_nonresident(lruvec, thp_nr_pages(page));
return pack_shadow(memcgid, pgdat, eviction, PageWorkingset(page));
* workingset_refault - Evaluate the refault of a previously evicted folio.
* @folio: The freshly allocated replacement folio.
* @shadow: Shadow entry of the evicted folio.
* Calculates and evaluates the refault distance of the previously
* evicted folio in the context of the node and the memcg whose memory
* pressure caused the eviction.
void workingset_refault(struct folio *folio, void *shadow)
bool file = folio_is_file_lru(folio);
struct mem_cgroup *eviction_memcg;
struct lruvec *eviction_lruvec;
unsigned long refault_distance;
unsigned long workingset_size;
struct pglist_data *pgdat;
struct mem_cgroup *memcg;
unsigned long eviction;
struct lruvec *lruvec;
unsigned long refault;
bool workingset;
int memcgid;
long nr;
unpack_shadow(shadow, &memcgid, &pgdat, &eviction, &workingset);
* Look up the memcg associated with the stored ID. It might
* have been deleted since the folio's eviction.
* Note that in rare events the ID could have been recycled
* for a new cgroup that refaults a shared folio. This is
* impossible to tell from the available data. However, this
* should be a rare and limited disturbance, and activations
* are always speculative anyway. Ultimately, it's the aging
* algorithm's job to shake out the minimum access frequency
* for the active cache.
* XXX: On !CONFIG_MEMCG, this will always return NULL; it
* would be better if the root_mem_cgroup existed in all
* configurations instead.
eviction_memcg = mem_cgroup_from_id(memcgid);
if (!mem_cgroup_disabled() && !eviction_memcg)
goto out;
eviction_lruvec = mem_cgroup_lruvec(eviction_memcg, pgdat);
refault = atomic_long_read(&eviction_lruvec->nonresident_age);
* Calculate the refault distance
* The unsigned subtraction here gives an accurate distance
* across nonresident_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
* nonresident_age to lap a shadow entry in the field, which
* can then 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.
refault_distance = (refault - eviction) & EVICTION_MASK;
* The activation decision for this folio is made at the level
* where the eviction occurred, as that is where the LRU order
* during folio reclaim is being determined.
* However, the cgroup that will own the folio is the one that
* is actually experiencing the refault event.
nr = folio_nr_pages(folio);
memcg = folio_memcg(folio);
lruvec = mem_cgroup_lruvec(memcg, pgdat);
mod_lruvec_state(lruvec, WORKINGSET_REFAULT_BASE + file, nr);
* Compare the distance to the existing workingset size. We
* don't activate pages that couldn't stay resident even if
* all the memory was available to the workingset. Whether
* workingset competition needs to consider anon or not depends
* on having swap.
workingset_size = lruvec_page_state(eviction_lruvec, NR_ACTIVE_FILE);
if (!file) {
workingset_size += lruvec_page_state(eviction_lruvec,
if (mem_cgroup_get_nr_swap_pages(memcg) > 0) {
workingset_size += lruvec_page_state(eviction_lruvec,
if (file) {
workingset_size += lruvec_page_state(eviction_lruvec,
if (refault_distance > workingset_size)
goto out;
workingset_age_nonresident(lruvec, nr);
mod_lruvec_state(lruvec, WORKINGSET_ACTIVATE_BASE + file, nr);
/* Folio was active prior to eviction */
if (workingset) {
/* XXX: Move to lru_cache_add() when it supports new vs putback */
mod_lruvec_state(lruvec, WORKINGSET_RESTORE_BASE + file, nr);
* workingset_activation - note a page activation
* @folio: Folio that is being activated.
void workingset_activation(struct folio *folio)
struct mem_cgroup *memcg;
* Filter non-memcg pages here, e.g. unmap can call
* mark_page_accessed() on VDSO pages.
* XXX: See workingset_refault() - this should return
* root_mem_cgroup even for !CONFIG_MEMCG.
memcg = folio_memcg_rcu(folio);
if (!mem_cgroup_disabled() && !memcg)
goto out;
workingset_age_nonresident(folio_lruvec(folio), folio_nr_pages(folio));
* 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.
static struct list_lru shadow_nodes;
void workingset_update_node(struct xa_node *node)
* Track non-empty nodes that contain only shadow entries;
* unlink those that contain pages or are being freed.
* Avoid acquiring the list_lru lock when the nodes are
* already where they should be. The list_empty() test is safe
* as node->private_list is protected by the i_pages lock.
VM_WARN_ON_ONCE(!irqs_disabled()); /* For __inc_lruvec_page_state */
if (node->count && node->count == node->nr_values) {
if (list_empty(&node->private_list)) {
list_lru_add(&shadow_nodes, &node->private_list);
__inc_lruvec_kmem_state(node, WORKINGSET_NODES);
} else {
if (!list_empty(&node->private_list)) {
list_lru_del(&shadow_nodes, &node->private_list);
__dec_lruvec_kmem_state(node, WORKINGSET_NODES);
static unsigned long count_shadow_nodes(struct shrinker *shrinker,
struct shrink_control *sc)
unsigned long max_nodes;
unsigned long nodes;
unsigned long pages;
nodes = list_lru_shrink_count(&shadow_nodes, sc);
if (!nodes)
* Approximate a reasonable limit for the nodes
* containing shadow entries. We don't need to keep more
* shadow entries than possible pages on the active list,
* since refault distances bigger than that are dismissed.
* The size of the active list converges toward 100% of
* overall page cache as memory grows, with only a tiny
* inactive list. Assume the total cache size for that.
* Nodes might be sparsely populated, with only one shadow
* entry in the extreme case. Obviously, we cannot keep one
* node for every eligible shadow entry, so compromise on a
* worst-case density of 1/8th. Below that, not all eligible
* refaults can be detected anymore.
* On 64-bit with 7 xa_nodes per page and 64 slots
* each, this will reclaim shadow entries when they consume
* ~1.8% of available memory:
* PAGE_SIZE / xa_nodes / node_entries * 8 / PAGE_SIZE
if (sc->memcg) {
struct lruvec *lruvec;
int i;
lruvec = mem_cgroup_lruvec(sc->memcg, NODE_DATA(sc->nid));
for (pages = 0, i = 0; i < NR_LRU_LISTS; i++)
pages += lruvec_page_state_local(lruvec,
pages += lruvec_page_state_local(
pages += lruvec_page_state_local(
} else
pages = node_present_pages(sc->nid);
max_nodes = pages >> (XA_CHUNK_SHIFT - 3);
if (nodes <= max_nodes)
return 0;
return 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) __must_hold(lru_lock)
struct xa_node *node = container_of(item, struct xa_node, private_list);
struct address_space *mapping;
int ret;
* Page cache insertions and deletions synchronously maintain
* the shadow node LRU under the i_pages 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 nodes on the LRU.
* We can then safely transition to the i_pages 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.
mapping = container_of(node->array, struct address_space, i_pages);
/* Coming from the list, invert the lock order */
if (!xa_trylock(&mapping->i_pages)) {
ret = LRU_RETRY;
goto out;
if (!spin_trylock(&mapping->host->i_lock)) {
ret = LRU_RETRY;
goto out;
list_lru_isolate(lru, item);
__dec_lruvec_kmem_state(node, WORKINGSET_NODES);
* 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.
if (WARN_ON_ONCE(!node->nr_values))
goto out_invalid;
if (WARN_ON_ONCE(node->count != node->nr_values))
goto out_invalid;
xa_delete_node(node, workingset_update_node);
__inc_lruvec_kmem_state(node, WORKINGSET_NODERECLAIM);
if (mapping_shrinkable(mapping))
return ret;
static unsigned long scan_shadow_nodes(struct shrinker *shrinker,
struct shrink_control *sc)
/* list_lru lock nests inside the IRQ-safe i_pages lock */
return list_lru_shrink_walk_irq(&shadow_nodes, sc, shadow_lru_isolate,
static struct shrinker workingset_shadow_shrinker = {
.count_objects = count_shadow_nodes,
.scan_objects = scan_shadow_nodes,
.seeks = 0, /* ->count reports only fully expendable nodes */
* Our list_lru->lock is IRQ-safe as it nests inside the IRQ-safe
* i_pages lock.
static struct lock_class_key shadow_nodes_key;
static int __init workingset_init(void)
unsigned int timestamp_bits;
unsigned int max_order;
int ret;
* Calculate the eviction bucket size to cover the longest
* actionable refault distance, which is currently half of
* memory (totalram_pages/2). However, memory hotplug may add
* some more pages at runtime, so keep working with up to
* double the initial memory by using totalram_pages as-is.
timestamp_bits = BITS_PER_LONG - EVICTION_SHIFT;
max_order = fls_long(totalram_pages() - 1);
if (max_order > timestamp_bits)
bucket_order = max_order - timestamp_bits;
pr_info("workingset: timestamp_bits=%d max_order=%d bucket_order=%u\n",
timestamp_bits, max_order, bucket_order);
ret = prealloc_shrinker(&workingset_shadow_shrinker);
if (ret)
goto err;
ret = __list_lru_init(&shadow_nodes, true, &shadow_nodes_key,
if (ret)
goto err_list_lru;
return 0;
return ret;