前面已经分析了slub算法的初始化、缓存区的创建、对象的分配、对象的回收，最后分析一下slub分配算法的slab销毁具体实现。
Slab销毁的入口函数为kmem_cache_destroy()，其实现：
【file:/mm/slab_common.c】
void kmem_cache_destroy(struct kmem_cache *s)
{
/* Destroy all the children caches if we aren't a memcg cache */
kmem_cache_destroy_memcg_children(s);
get_online_cpus();
mutex_lock(&slab_mutex);
s->refcount--;
if (!s->refcount) {
list_del(&s->list);
if (!__kmem_cache_shutdown(s)) {
memcg_unregister_cache(s);
mutex_unlock(&slab_mutex);
if (s->flags & SLAB_DESTROY_BY_RCU)
rcu_barrier();
memcg_free_cache_params(s);
kfree(s->name);
kmem_cache_free(kmem_cache, s);
} else {
list_add(&s->list, &slab_caches);
mutex_unlock(&slab_mutex);
printk(KERN_ERR "kmem_cache_destroy %s: Slab cache still has objects\n",
s->name);
dump_stack();
}
} else {
mutex_unlock(&slab_mutex);
}
put_online_cpus();
}
该函数中kmem_cache_destroy_memcg_children()删除memcg中相关联的子cache数据，而get_online_cpus()是对cpu_online_map的加锁，其与末尾的put_online_cpus()是配对使用的。接着的mutex_lock()用于获取slab_mutex互斥锁，该锁主要用于全局资源保护。然后对kmem_cache的引用计数refcount自减操作，如果自减后if (!s->refcount)为true，即引用计数为0，表示该缓冲区不存在slab别名挂靠的情况，那么其kmem_cache结构可以删除，否则表示有其他缓冲区别名挂靠，仍有依赖，那么将会解锁slab_mutex并put_online_cpus()释放cpu_online_map锁，然后退出。
if (!s->refcount)为true的分支中，先list_del()将该slab管理结构kmem_cache从slab_caches全局链表中摘除，然后__kmem_cache_shutdown()删除kmem_cache结构信息。如果__kmem_cache_shutdown()执行成功则将返回0，继而if (!__kmem_cache_shutdown(s))为true，将会通过memcg_unregister_cache()去注册memcg的cache，并且memcg_free_cache_params()释放创建时申请的memcg_params资源空间，而kfree()和kmem_cache_free()释放slub的名称空间以及slab空间。如果__kmem_cache_shutdown()执行失败，那么将会把slab重新挂回至slab_caches链表，同时记录日志信息。
由此slab销毁完毕。
kmem_cache_destroy()的核心函数是__kmem_cache_shutdown()，深入分析__kmem_cache_shutdown()的实现：
【file:/mm/slub.c】
int __kmem_cache_shutdown(struct kmem_cache *s)
{
int rc = kmem_cache_close(s);
if (!rc) {
/*
* We do the same lock strategy around sysfs_slab_add, see
* __kmem_cache_create. Because this is pretty much the last
* operation we do and the lock will be released shortly after
* that in slab_common.c, we could just move sysfs_slab_remove
* to a later point in common code. We should do that when we
* have a common sysfs framework for all allocators.
*/
mutex_unlock(&slab_mutex);
sysfs_slab_remove(s);
mutex_lock(&slab_mutex);
}
return rc;
}
该函数主要通过kmem_cache_close()删除slab的管理数据kmem_cache，如果执行成功，继而进入if分支对sysfs模块的slab做移除操作。
具体看一下kmem_cache_close()的实现：
【file:/mm/slub.c】
/*
* Release all resources used by a slab cache.
*/
static inline int kmem_cache_close(struct kmem_cache *s)
{
int node;
flush_all(s);
/* Attempt to free all objects */
for_each_node_state(node, N_NORMAL_MEMORY) {
struct kmem_cache_node *n = get_node(s, node);
free_partial(s, n);
if (n->nr_partial || slabs_node(s, node))
return 1;
}
free_percpu(s->cpu_slab);
free_kmem_cache_nodes(s);
return 0;
}
该函数通过flush_all()释放本地CPU的缓存区，即kmem_cache_cpu管理的缓存区空间；然后通过for_each_node_state()遍历各节点，转而get_node()获取节点下的kmem_cache_node管理结构，然后将其半满队列中的缓存区进行释放free_partial()；最后将kmem_cache的每CPU缓存管理kmem_cache_cpu通过free_percpu()归还给系统，同时通过free_kmem_cache_nodes()释放各内存节点node的缓存管理结构kmem_cache_node占用的空间释放。
最后分析一下较为复杂的flush_all()的实现：
【file:/mm/slub.c】
static void flush_all(struct kmem_cache *s)
{
on_each_cpu_cond(has_cpu_slab, flush_cpu_slab, s, 1, GFP_ATOMIC);
}
看似封装了on_each_cpu_cond()函数，实际上on_each_cpu_cond()并不执行任何与资源释放的操作，其主要是遍历各个CPU，然后执行作为入参传入的函数has_cpu_slab()，以判断各个处理器上的资源是否存在，如果存在，继而将会通过flush_cpu_slab()对该处理器上的资源进行释放处理。
照例，还是详细看一下on_each_cpu_cond()函数实现：
【file:/mm/slub.c】
/*
* on_each_cpu_cond(): Call a function on each processor for which
* the supplied function cond_func returns true, optionally waiting
* for all the required CPUs to finish. This may include the local
* processor.
* @cond_func: A callback function that is passed a cpu id and
* the the info parameter. The function is called
* with preemption disabled. The function should
* return a blooean value indicating whether to IPI
* the specified CPU.
* @func: The function to run on all applicable CPUs.
* This must be fast and non-blocking.
* @info: An arbitrary pointer to pass to both functions.
* @wait: If true, wait (atomically) until function has
* completed on other CPUs.
* @gfp_flags: GFP flags to use when allocating the cpumask
* used internally by the function.
*
* The function might sleep if the GFP flags indicates a non
* atomic allocation is allowed.
*
* Preemption is disabled to protect against CPUs going offline but not online.
* CPUs going online during the call will not be seen or sent an IPI.
*
* You must not call this function with disabled interrupts or
* from a hardware interrupt handler or from a bottom half handler.
*/
void on_each_cpu_cond(bool (*cond_func)(int cpu, void *info),
smp_call_func_t func, void *info, bool wait,
gfp_t gfp_flags)
{
cpumask_var_t cpus;
int cpu, ret;
might_sleep_if(gfp_flags & __GFP_WAIT);
if (likely(zalloc_cpumask_var(&cpus, (gfp_flags|__GFP_NOWARN)))) {
preempt_disable();
for_each_online_cpu(cpu)
if (cond_func(cpu, info))
cpumask_set_cpu(cpu, cpus);
on_each_cpu_mask(cpus, func, info, wait);
preempt_enable();
free_cpumask_var(cpus);
} else {
/*
* No free cpumask, bother. No matter, we'll
* just have to IPI them one by one.
*/
preempt_disable();
for_each_online_cpu(cpu)
if (cond_func(cpu, info)) {
ret = smp_call_function_single(cpu, func,
info, wait);
WARN_ON_ONCE(!ret);
}
preempt_enable();
}
}
该函数的入参cond_func是一个钩子函数，用于根据调用者传入的CPU信息参数来判断是否需要打断该CPU以执行入参func的操作；而入参info是作为cond_func和func处理函数的入参；至于入参wait则是一个bool类型，用以判断是否需要等待func在各CPU上执行完毕，如果为true将会等待；最后的gfp_flags入参是作为申请cpumask空间的标识。
了解完参数的意思，那么具体看一下其实现，首先might_sleep_if()判断是否需要休眠等待，继而通过zalloc_cpumask_var()申请cpumask的空间；申请到空间后，preempt_disable()禁止内核抢占后，将for_each_online_cpu()遍历各个CPU，根据cond_func()（即has_cpu_slab()）判断是否需要对该CPU进行打断处理，如果需要则cpumask_set_cpu()对该CPU进行标志；标志完后，根据前面的标志，通过on_each_cpu_mask()打断各个标志位对应的CPU去执行func()的操作（即flush_cpu_slab()）；完了将会恢复抢占，释放cpumask空间。至于zalloc_cpumask_var()申请不到空间，将会逐个处理器进行打断再进行处理，其最终功能和作用与申请到空间的情况都是一致的，具体实现就不分析了。
相应看一下作为on_each_cpu_cond()入参的钩子函数has_cpu_slab()的实现：
【file:/mm/slub.c】
static bool has_cpu_slab(int cpu, void *info)
{
struct kmem_cache *s = info;
struct kmem_cache_cpu *c = per_cpu_ptr(s->cpu_slab, cpu);
return c->page || c->partial;
}
可以看到该函数主要是用于判断本地CPU是否占有缓存区，如果有则返回true。也即意味着该CPU需要被打断去执行其本地的缓存区释放操作。
至于on_each_cpu_cond()另一钩子函数flush_cpu_slab()的实现：
【file:/mm/slub.c】
static void flush_cpu_slab(void *d)
{
struct kmem_cache *s = d;
__flush_cpu_slab(s, smp_processor_id());
}
该函数封装了__flush_cpu_slab()，实现为：
【file:/mm/slub.c】
/*
* Flush cpu slab.
*
* Called from IPI handler with interrupts disabled.
*/
static inline void __flush_cpu_slab(struct kmem_cache *s, int cpu)
{
struct kmem_cache_cpu *c = per_cpu_ptr(s->cpu_slab, cpu);
if (likely(c)) {
if (c->page)
flush_slab(s, c);
unfreeze_partials(s, c);
}
}
函数实现很简单，主要用于将本地CPU的缓存区进行释放。其首先获取本地CPU的kmem_cache_cpu管理结构，如果本地CPU存在缓存区的占用，将会通过flush_slab()去释放本地缓存区，继而通过unfreeze_partials()将本地CPU半满缓存列表进行释放。
而flush_slab()具体实现：
【file:/mm/slub.c】
static inline void flush_slab(struct kmem_cache *s, struct kmem_cache_cpu *c)
{
stat(s, CPUSLAB_FLUSH);
deactivate_slab(s, c->page, c->freelist);
c->tid = next_tid(c->tid);
c->page = NULL;
c->freelist = NULL;
}
其主要是通过deactivate_slab()去激活本地缓存区，也即是将缓存区进行释放操作。具体deactivate_slab()的实现：
【file:/mm/slub.c】
/*
* Remove the cpu slab
*/
static void deactivate_slab(struct kmem_cache *s, struct page *page,
void *freelist)
{
enum slab_modes { M_NONE, M_PARTIAL, M_FULL, M_FREE };
struct kmem_cache_node *n = get_node(s, page_to_nid(page));
int lock = 0;
enum slab_modes l = M_NONE, m = M_NONE;
void *nextfree;
int tail = DEACTIVATE_TO_HEAD;
struct page new;
struct page old;
if (page->freelist) {
stat(s, DEACTIVATE_REMOTE_FREES);
tail = DEACTIVATE_TO_TAIL;
}
/*
* Stage one: Free all available per cpu objects back
* to the page freelist while it is still frozen. Leave the
* last one.
*
* There is no need to take the list->lock because the page
* is still frozen.
*/
while (freelist && (nextfree = get_freepointer(s, freelist))) {
void *prior;
unsigned long counters;
do {
prior = page->freelist;
counters = page->counters;
set_freepointer(s, freelist, prior);
new.counters = counters;
new.inuse--;
VM_BUG_ON(!new.frozen);
} while (!__cmpxchg_double_slab(s, page,
prior, counters,
freelist, new.counters,
"drain percpu freelist"));
freelist = nextfree;
}
/*
* Stage two: Ensure that the page is unfrozen while the
* list presence reflects the actual number of objects
* during unfreeze.
*
* We setup the list membership and then perform a cmpxchg
* with the count. If there is a mismatch then the page
* is not unfrozen but the page is on the wrong list.
*
* Then we restart the process which may have to remove
* the page from the list that we just put it on again
* because the number of objects in the slab may have
* changed.
*/
redo:
old.freelist = page->freelist;
old.counters = page->counters;
VM_BUG_ON(!old.frozen);
/* Determine target state of the slab */
new.counters = old.counters;
if (freelist) {
new.inuse--;
set_freepointer(s, freelist, old.freelist);
new.freelist = freelist;
} else
new.freelist = old.freelist;
new.frozen = 0;
if (!new.inuse && n->nr_partial > s->min_partial)
m = M_FREE;
else if (new.freelist) {
m = M_PARTIAL;
if (!lock) {
lock = 1;
/*
* Taking the spinlock removes the possiblity
* that acquire_slab() will see a slab page that
* is frozen
*/
spin_lock(&n->list_lock);
}
} else {
m = M_FULL;
if (kmem_cache_debug(s) && !lock) {
lock = 1;
/*
* This also ensures that the scanning of full
* slabs from diagnostic functions will not see
* any frozen slabs.
*/
spin_lock(&n->list_lock);
}
}
if (l != m) {
if (l == M_PARTIAL)
remove_partial(n, page);
else if (l == M_FULL)
remove_full(s, n, page);
if (m == M_PARTIAL) {
add_partial(n, page, tail);
stat(s, tail);
} else if (m == M_FULL) {
stat(s, DEACTIVATE_FULL);
add_full(s, n, page);
}
}
l = m;
if (!__cmpxchg_double_slab(s, page,
old.freelist, old.counters,
new.freelist, new.counters,
"unfreezing slab"))
goto redo;
if (lock)
spin_unlock(&n->list_lock);
if (m == M_FREE) {
stat(s, DEACTIVATE_EMPTY);
discard_slab(s, page);
stat(s, FREE_SLAB);
}
}
if (page->freelist)判断slab的空闲链表freelist是否为空，如果为空，意味着该缓存区的对象已经全部分配到了CPU的kmem_cache_cpu中freelist链表中；如果不为空，那么表示该CPU的slab对象被其他CPU释放了，将会更新统计同时设置tail标识为DEACTIVATE_TO_TAIL。
接下来的while循环是去激活本地CPU的slab步骤一，其主要是通过while循环遍历CPU上的freelist链表get_freepointer()获取空闲对象，继而通过内部的do-while循环，借用__cmpxchg_double_slab()比较交换将对象以插入缓存区页面的freelist空闲链表头的方式归还回去。__cmpxchg_double_slab()前面已经介绍过了的原子操作，这里将不再赘述。不过有个点值得注意的是该步骤的释放操作，其并未将所有的对象都归还回去，这是由于nextfree = get_freepointer(s, freelist)该步骤取下一个空闲对象时得到空指针，那么将会退出while循环；也就意味着如果deactivate_slab()入参中freelist不为空，那么while循环退出时，其也必定不为空，其具体用意稍后再分析。简而言之该步骤其目的是，当页面还处于冻结状态，将会释放每CPU的所有可用的对象回到缓冲区的空闲列表中。
然后是步骤二，即redo标签以下的动作，其首先将缓存区的freelist以及counters信息存到临时old结构中以备后用，接着if (freelist)如果为true，将会把前面步骤一未被归还的那个对象归还到缓冲区中，同时更新new信息，此时new.freelist持有该缓存区的所有空闲对象。往下new.frozen = 0将临时缓存区状态设置为非冻结；然后if (!new.inuse && n->nr_partial > s->min_partial) 表示该slab缓存区中无对象被使用，且部分满slab个数大于最小值，意味着该缓存区需要被销毁，标识m为M_FREE；而else if (new.freelist)表示freelist不为空，仅使用了部分对象，则标识m为M_PARTIAL；至于最后的else分支，表示freelist为空，该缓存区所有对象均已被使用，m标识为M_FULL。再往下if (l != m)的比较是用于判断上一次的缓存区状态l与接下来的操作状态m是否一致，不一致则意味着需要发生变更，其将会先判断l的状态为M_PARTIAL或M_FULL，继而采取对应的remove_partial()或remove_full()链表摘除操作；继而根据m的状态，往半满链表中添加add_partial()还是往满载链表中添加add_full()，接着将l的状态更新为m。现在到了if (!__cmpxchg_double_slab())，这里是用于判断自redo到此，缓存区是否发生过对象操作变更，如果没发生过的话，将会把new暂存的空闲对象挂载到缓存区中以及更新counters，否则将会跳转回redo标签重新执行前面的操作。至此，顺利的话，缓存区已经去激活完毕了。
最后如果m的状态为M_FREE，则表示该缓存区不需要再使用了，将通过discard_slab()将其销毁。
至此，slub算法分析完毕。