Just how the locking-end behaves, when unlocking, go ahead and
obtain the proper data structure immediately after the previous
(asm-end) call exits and there are (probably) pending waiters.
This simplifies a bit some of the layering.
Signed-off-by: Davidlohr Bueso <[email protected]>
---
kernel/locking/mutex.c | 7 ++++---
1 file changed, 4 insertions(+), 3 deletions(-)
diff --git a/kernel/locking/mutex.c b/kernel/locking/mutex.c
index ae712b2..ad0e333 100644
--- a/kernel/locking/mutex.c
+++ b/kernel/locking/mutex.c
@@ -679,9 +679,8 @@ EXPORT_SYMBOL_GPL(__ww_mutex_lock_interruptible);
* Release the lock, slowpath:
*/
static inline void
-__mutex_unlock_common_slowpath(atomic_t *lock_count, int nested)
+__mutex_unlock_common_slowpath(struct mutex *lock, int nested)
{
- struct mutex *lock = container_of(lock_count, struct mutex, count);
unsigned long flags;
/*
@@ -716,7 +715,9 @@ __mutex_unlock_common_slowpath(atomic_t *lock_count, int nested)
__visible void
__mutex_unlock_slowpath(atomic_t *lock_count)
{
- __mutex_unlock_common_slowpath(lock_count, 1);
+ struct mutex *lock = container_of(lock_count, struct mutex, count);
+
+ __mutex_unlock_common_slowpath(lock, 1);
}
#ifndef CONFIG_DEBUG_LOCK_ALLOC
--
1.8.1.4
When we fail to acquire the mutex in the fastpath, we end up calling
__mutex_lock_common(). A *lot* goes on in this function. Move out the
optimistic spinning code into mutex_optimistic_spin() and simplify
the former a bit. Furthermore, this is similar to what we have in
rwsems. No logical changes.
Signed-off-by: Davidlohr Bueso <[email protected]>
---
Changes from v1:
- Removed duplicate need_resched() check.
- Added missing lock_acquired() in mutex_optimistic_spin().
kernel/locking/mutex.c | 396 ++++++++++++++++++++++++++-----------------------
1 file changed, 214 insertions(+), 182 deletions(-)
diff --git a/kernel/locking/mutex.c b/kernel/locking/mutex.c
index 93bec48..0d8b6ed 100644
--- a/kernel/locking/mutex.c
+++ b/kernel/locking/mutex.c
@@ -106,6 +106,92 @@ void __sched mutex_lock(struct mutex *lock)
EXPORT_SYMBOL(mutex_lock);
#endif
+static __always_inline void ww_mutex_lock_acquired(struct ww_mutex *ww,
+ struct ww_acquire_ctx *ww_ctx)
+{
+#ifdef CONFIG_DEBUG_MUTEXES
+ /*
+ * If this WARN_ON triggers, you used ww_mutex_lock to acquire,
+ * but released with a normal mutex_unlock in this call.
+ *
+ * This should never happen, always use ww_mutex_unlock.
+ */
+ DEBUG_LOCKS_WARN_ON(ww->ctx);
+
+ /*
+ * Not quite done after calling ww_acquire_done() ?
+ */
+ DEBUG_LOCKS_WARN_ON(ww_ctx->done_acquire);
+
+ if (ww_ctx->contending_lock) {
+ /*
+ * After -EDEADLK you tried to
+ * acquire a different ww_mutex? Bad!
+ */
+ DEBUG_LOCKS_WARN_ON(ww_ctx->contending_lock != ww);
+
+ /*
+ * You called ww_mutex_lock after receiving -EDEADLK,
+ * but 'forgot' to unlock everything else first?
+ */
+ DEBUG_LOCKS_WARN_ON(ww_ctx->acquired > 0);
+ ww_ctx->contending_lock = NULL;
+ }
+
+ /*
+ * Naughty, using a different class will lead to undefined behavior!
+ */
+ DEBUG_LOCKS_WARN_ON(ww_ctx->ww_class != ww->ww_class);
+#endif
+ ww_ctx->acquired++;
+}
+
+/*
+ * after acquiring lock with fastpath or when we lost out in contested
+ * slowpath, set ctx and wake up any waiters so they can recheck.
+ *
+ * This function is never called when CONFIG_DEBUG_LOCK_ALLOC is set,
+ * as the fastpath and opportunistic spinning are disabled in that case.
+ */
+static __always_inline void
+ww_mutex_set_context_fastpath(struct ww_mutex *lock,
+ struct ww_acquire_ctx *ctx)
+{
+ unsigned long flags;
+ struct mutex_waiter *cur;
+
+ ww_mutex_lock_acquired(lock, ctx);
+
+ lock->ctx = ctx;
+
+ /*
+ * The lock->ctx update should be visible on all cores before
+ * the atomic read is done, otherwise contended waiters might be
+ * missed. The contended waiters will either see ww_ctx == NULL
+ * and keep spinning, or it will acquire wait_lock, add itself
+ * to waiter list and sleep.
+ */
+ smp_mb(); /* ^^^ */
+
+ /*
+ * Check if lock is contended, if not there is nobody to wake up
+ */
+ if (likely(atomic_read(&lock->base.count) == 0))
+ return;
+
+ /*
+ * Uh oh, we raced in fastpath, wake up everyone in this case,
+ * so they can see the new lock->ctx.
+ */
+ spin_lock_mutex(&lock->base.wait_lock, flags);
+ list_for_each_entry(cur, &lock->base.wait_list, list) {
+ debug_mutex_wake_waiter(&lock->base, cur);
+ wake_up_process(cur->task);
+ }
+ spin_unlock_mutex(&lock->base.wait_lock, flags);
+}
+
+
#ifdef CONFIG_MUTEX_SPIN_ON_OWNER
/*
* In order to avoid a stampede of mutex spinners from acquiring the mutex
@@ -180,6 +266,129 @@ static inline int mutex_can_spin_on_owner(struct mutex *lock)
*/
return retval;
}
+
+/*
+ * Atomically try to take the lock when it is available
+ */
+static inline bool mutex_try_to_acquire(struct mutex *lock)
+{
+ return !mutex_is_locked(lock) &&
+ (atomic_cmpxchg(&lock->count, 1, 0) == 1);
+}
+
+/*
+ * Optimistic spinning.
+ *
+ * We try to spin for acquisition when we find that the lock owner
+ * is currently running on a (different) CPU and while we don't
+ * need to reschedule. The rationale is that if the lock owner is
+ * running, it is likely to release the lock soon.
+ *
+ * Since this needs the lock owner, and this mutex implementation
+ * doesn't track the owner atomically in the lock field, we need to
+ * track it non-atomically.
+ *
+ * We can't do this for DEBUG_MUTEXES because that relies on wait_lock
+ * to serialize everything.
+ *
+ * The mutex spinners are queued up using MCS lock so that only one
+ * spinner can compete for the mutex. However, if mutex spinning isn't
+ * going to happen, there is no point in going through the lock/unlock
+ * overhead.
+ *
+ * Returns true when the lock was taken, otherwise false, indicating
+ * that we need to jump to the slowpath and sleep.
+ */
+static bool mutex_optimistic_spin(struct mutex *lock,
+ struct ww_acquire_ctx *ww_ctx, const bool use_ww_ctx)
+{
+ struct task_struct *task = current;
+
+ if (!mutex_can_spin_on_owner(lock))
+ goto done;
+
+ if (!osq_lock(&lock->osq))
+ goto done;
+
+ while (true) {
+ struct task_struct *owner;
+
+ if (use_ww_ctx && ww_ctx->acquired > 0) {
+ struct ww_mutex *ww;
+
+ ww = container_of(lock, struct ww_mutex, base);
+ /*
+ * If ww->ctx is set the contents are undefined, only
+ * by acquiring wait_lock there is a guarantee that
+ * they are not invalid when reading.
+ *
+ * As such, when deadlock detection needs to be
+ * performed the optimistic spinning cannot be done.
+ */
+ if (ACCESS_ONCE(ww->ctx))
+ break;
+ }
+
+ /*
+ * If there's an owner, wait for it to either
+ * release the lock or go to sleep.
+ */
+ owner = ACCESS_ONCE(lock->owner);
+ if (owner && !mutex_spin_on_owner(lock, owner))
+ break;
+
+ /* Try to acquire the mutex if it is unlocked. */
+ if (mutex_try_to_acquire(lock)) {
+ lock_acquired(&lock->dep_map, ip);
+
+ if (use_ww_ctx) {
+ struct ww_mutex *ww;
+ ww = container_of(lock, struct ww_mutex, base);
+
+ ww_mutex_set_context_fastpath(ww, ww_ctx);
+ }
+
+ mutex_set_owner(lock);
+ osq_unlock(&lock->osq);
+ return true;
+ }
+
+ /*
+ * When there's no owner, we might have preempted between the
+ * owner acquiring the lock and setting the owner field. If
+ * we're an RT task that will live-lock because we won't let
+ * the owner complete.
+ */
+ if (!owner && (need_resched() || rt_task(task)))
+ break;
+
+ /*
+ * The cpu_relax() call is a compiler barrier which forces
+ * everything in this loop to be re-loaded. We don't need
+ * memory barriers as we'll eventually observe the right
+ * values at the cost of a few extra spins.
+ */
+ cpu_relax_lowlatency();
+ }
+
+ osq_unlock(&lock->osq);
+done:
+ /*
+ * If we fell out of the spin path because of need_resched(),
+ * reschedule now, before we try-lock the mutex. This avoids getting
+ * scheduled out right after we obtained the mutex.
+ */
+ if (need_resched())
+ schedule_preempt_disabled();
+
+ return false;
+}
+#else
+static bool mutex_optimistic_spin(struct mutex *lock,
+ struct ww_acquire_ctx *ww_ctx, const bool use_ww_ctx)
+{
+ return false;
+}
#endif
__visible __used noinline
@@ -277,91 +486,6 @@ __mutex_lock_check_stamp(struct mutex *lock, struct ww_acquire_ctx *ctx)
return 0;
}
-static __always_inline void ww_mutex_lock_acquired(struct ww_mutex *ww,
- struct ww_acquire_ctx *ww_ctx)
-{
-#ifdef CONFIG_DEBUG_MUTEXES
- /*
- * If this WARN_ON triggers, you used ww_mutex_lock to acquire,
- * but released with a normal mutex_unlock in this call.
- *
- * This should never happen, always use ww_mutex_unlock.
- */
- DEBUG_LOCKS_WARN_ON(ww->ctx);
-
- /*
- * Not quite done after calling ww_acquire_done() ?
- */
- DEBUG_LOCKS_WARN_ON(ww_ctx->done_acquire);
-
- if (ww_ctx->contending_lock) {
- /*
- * After -EDEADLK you tried to
- * acquire a different ww_mutex? Bad!
- */
- DEBUG_LOCKS_WARN_ON(ww_ctx->contending_lock != ww);
-
- /*
- * You called ww_mutex_lock after receiving -EDEADLK,
- * but 'forgot' to unlock everything else first?
- */
- DEBUG_LOCKS_WARN_ON(ww_ctx->acquired > 0);
- ww_ctx->contending_lock = NULL;
- }
-
- /*
- * Naughty, using a different class will lead to undefined behavior!
- */
- DEBUG_LOCKS_WARN_ON(ww_ctx->ww_class != ww->ww_class);
-#endif
- ww_ctx->acquired++;
-}
-
-/*
- * after acquiring lock with fastpath or when we lost out in contested
- * slowpath, set ctx and wake up any waiters so they can recheck.
- *
- * This function is never called when CONFIG_DEBUG_LOCK_ALLOC is set,
- * as the fastpath and opportunistic spinning are disabled in that case.
- */
-static __always_inline void
-ww_mutex_set_context_fastpath(struct ww_mutex *lock,
- struct ww_acquire_ctx *ctx)
-{
- unsigned long flags;
- struct mutex_waiter *cur;
-
- ww_mutex_lock_acquired(lock, ctx);
-
- lock->ctx = ctx;
-
- /*
- * The lock->ctx update should be visible on all cores before
- * the atomic read is done, otherwise contended waiters might be
- * missed. The contended waiters will either see ww_ctx == NULL
- * and keep spinning, or it will acquire wait_lock, add itself
- * to waiter list and sleep.
- */
- smp_mb(); /* ^^^ */
-
- /*
- * Check if lock is contended, if not there is nobody to wake up
- */
- if (likely(atomic_read(&lock->base.count) == 0))
- return;
-
- /*
- * Uh oh, we raced in fastpath, wake up everyone in this case,
- * so they can see the new lock->ctx.
- */
- spin_lock_mutex(&lock->base.wait_lock, flags);
- list_for_each_entry(cur, &lock->base.wait_list, list) {
- debug_mutex_wake_waiter(&lock->base, cur);
- wake_up_process(cur->task);
- }
- spin_unlock_mutex(&lock->base.wait_lock, flags);
-}
-
/*
* Lock a mutex (possibly interruptible), slowpath:
*/
@@ -378,104 +502,12 @@ __mutex_lock_common(struct mutex *lock, long state, unsigned int subclass,
preempt_disable();
mutex_acquire_nest(&lock->dep_map, subclass, 0, nest_lock, ip);
-#ifdef CONFIG_MUTEX_SPIN_ON_OWNER
- /*
- * Optimistic spinning.
- *
- * We try to spin for acquisition when we find that the lock owner
- * is currently running on a (different) CPU and while we don't
- * need to reschedule. The rationale is that if the lock owner is
- * running, it is likely to release the lock soon.
- *
- * Since this needs the lock owner, and this mutex implementation
- * doesn't track the owner atomically in the lock field, we need to
- * track it non-atomically.
- *
- * We can't do this for DEBUG_MUTEXES because that relies on wait_lock
- * to serialize everything.
- *
- * The mutex spinners are queued up using MCS lock so that only one
- * spinner can compete for the mutex. However, if mutex spinning isn't
- * going to happen, there is no point in going through the lock/unlock
- * overhead.
- */
- if (!mutex_can_spin_on_owner(lock))
- goto slowpath;
-
- if (!osq_lock(&lock->osq))
- goto slowpath;
-
- for (;;) {
- struct task_struct *owner;
-
- if (use_ww_ctx && ww_ctx->acquired > 0) {
- struct ww_mutex *ww;
-
- ww = container_of(lock, struct ww_mutex, base);
- /*
- * If ww->ctx is set the contents are undefined, only
- * by acquiring wait_lock there is a guarantee that
- * they are not invalid when reading.
- *
- * As such, when deadlock detection needs to be
- * performed the optimistic spinning cannot be done.
- */
- if (ACCESS_ONCE(ww->ctx))
- break;
- }
-
- /*
- * If there's an owner, wait for it to either
- * release the lock or go to sleep.
- */
- owner = ACCESS_ONCE(lock->owner);
- if (owner && !mutex_spin_on_owner(lock, owner))
- break;
-
- /* Try to acquire the mutex if it is unlocked. */
- if (!mutex_is_locked(lock) &&
- (atomic_cmpxchg(&lock->count, 1, 0) == 1)) {
- lock_acquired(&lock->dep_map, ip);
- if (use_ww_ctx) {
- struct ww_mutex *ww;
- ww = container_of(lock, struct ww_mutex, base);
-
- ww_mutex_set_context_fastpath(ww, ww_ctx);
- }
-
- mutex_set_owner(lock);
- osq_unlock(&lock->osq);
- preempt_enable();
- return 0;
- }
-
- /*
- * When there's no owner, we might have preempted between the
- * owner acquiring the lock and setting the owner field. If
- * we're an RT task that will live-lock because we won't let
- * the owner complete.
- */
- if (!owner && (need_resched() || rt_task(task)))
- break;
-
- /*
- * The cpu_relax() call is a compiler barrier which forces
- * everything in this loop to be re-loaded. We don't need
- * memory barriers as we'll eventually observe the right
- * values at the cost of a few extra spins.
- */
- cpu_relax_lowlatency();
+ if (mutex_optimistic_spin(lock, ww_ctx, use_ww_ctx)) {
+ /* got the lock, yay! */
+ preempt_enable();
+ return 0;
}
- osq_unlock(&lock->osq);
-slowpath:
- /*
- * If we fell out of the spin path because of need_resched(),
- * reschedule now, before we try-lock the mutex. This avoids getting
- * scheduled out right after we obtained the mutex.
- */
- if (need_resched())
- schedule_preempt_disabled();
-#endif
+
spin_lock_mutex(&lock->wait_lock, flags);
/*
--
1.8.1.4
4badad35 (locking/mutex: Disable optimistic spinning on some
architectures) added a ARCH_SUPPORTS_ATOMIC_RMW flag to
disable the mutex optimistic feature on specific archs.
Because CONFIG_MUTEX_SPIN_ON_OWNER only depended on DEBUG and
SMP, it was ok to have the ->owner field conditional a bit
flexible. However by adding a new variable to the matter,
we can waste space with the unused field, ie: CONFIG_SMP &&
(!CONFIG_MUTEX_SPIN_ON_OWNER && !CONFIG_DEBUG_MUTEX).
Acked-by: Jason Low <[email protected]>
Signed-off-by: Davidlohr Bueso <[email protected]>
---
include/linux/mutex.h | 2 +-
kernel/locking/mutex.h | 2 +-
2 files changed, 2 insertions(+), 2 deletions(-)
diff --git a/include/linux/mutex.h b/include/linux/mutex.h
index 8d5535c..e4c2941 100644
--- a/include/linux/mutex.h
+++ b/include/linux/mutex.h
@@ -52,7 +52,7 @@ struct mutex {
atomic_t count;
spinlock_t wait_lock;
struct list_head wait_list;
-#if defined(CONFIG_DEBUG_MUTEXES) || defined(CONFIG_SMP)
+#if defined(CONFIG_DEBUG_MUTEXES) || defined(CONFIG_MUTEX_SPIN_ON_OWNER)
struct task_struct *owner;
#endif
#ifdef CONFIG_MUTEX_SPIN_ON_OWNER
diff --git a/kernel/locking/mutex.h b/kernel/locking/mutex.h
index 4115fbf..5cda397 100644
--- a/kernel/locking/mutex.h
+++ b/kernel/locking/mutex.h
@@ -16,7 +16,7 @@
#define mutex_remove_waiter(lock, waiter, ti) \
__list_del((waiter)->list.prev, (waiter)->list.next)
-#ifdef CONFIG_SMP
+#ifdef CONFIG_MUTEX_SPIN_ON_OWNER
static inline void mutex_set_owner(struct mutex *lock)
{
lock->owner = current;
--
1.8.1.4
Specifically:
Documentation/locking/lockdep-design.txt
Documentation/locking/lockstat.txt
Documentation/locking/mutex-design.txt
Documentation/locking/rt-mutex-design.txt
Documentation/locking/rt-mutex.txt
Documentation/locking/spinlocks.txt
Documentation/locking/ww-mutex-design.txt
Signed-off-by: Davidlohr Bueso <[email protected]>
---
Documentation/00-INDEX | 2 +
Documentation/DocBook/kernel-locking.tmpl | 2 +-
Documentation/lockdep-design.txt | 286 -----------
Documentation/locking/lockdep-design.txt | 286 +++++++++++
Documentation/locking/lockstat.txt | 178 +++++++
Documentation/locking/mutex-design.txt | 157 ++++++
Documentation/locking/rt-mutex-design.txt | 781 ++++++++++++++++++++++++++++++
Documentation/locking/rt-mutex.txt | 79 +++
Documentation/locking/spinlocks.txt | 167 +++++++
Documentation/locking/ww-mutex-design.txt | 344 +++++++++++++
Documentation/lockstat.txt | 178 -------
Documentation/mutex-design.txt | 157 ------
Documentation/rt-mutex-design.txt | 781 ------------------------------
Documentation/rt-mutex.txt | 79 ---
Documentation/spinlocks.txt | 167 -------
Documentation/ww-mutex-design.txt | 344 -------------
MAINTAINERS | 4 +-
drivers/gpu/drm/drm_modeset_lock.c | 2 +-
include/linux/lockdep.h | 2 +-
include/linux/mutex.h | 2 +-
include/linux/rwsem.h | 2 +-
kernel/locking/mutex.c | 2 +-
kernel/locking/rtmutex.c | 2 +-
lib/Kconfig.debug | 4 +-
24 files changed, 2005 insertions(+), 2003 deletions(-)
delete mode 100644 Documentation/lockdep-design.txt
create mode 100644 Documentation/locking/lockdep-design.txt
create mode 100644 Documentation/locking/lockstat.txt
create mode 100644 Documentation/locking/mutex-design.txt
create mode 100644 Documentation/locking/rt-mutex-design.txt
create mode 100644 Documentation/locking/rt-mutex.txt
create mode 100644 Documentation/locking/spinlocks.txt
create mode 100644 Documentation/locking/ww-mutex-design.txt
delete mode 100644 Documentation/lockstat.txt
delete mode 100644 Documentation/mutex-design.txt
delete mode 100644 Documentation/rt-mutex-design.txt
delete mode 100644 Documentation/rt-mutex.txt
delete mode 100644 Documentation/spinlocks.txt
delete mode 100644 Documentation/ww-mutex-design.txt
diff --git a/Documentation/00-INDEX b/Documentation/00-INDEX
index 27e67a9..1750fce 100644
--- a/Documentation/00-INDEX
+++ b/Documentation/00-INDEX
@@ -287,6 +287,8 @@ local_ops.txt
- semantics and behavior of local atomic operations.
lockdep-design.txt
- documentation on the runtime locking correctness validator.
+locking/
+ - directory with info about kernel locking primitives
lockstat.txt
- info on collecting statistics on locks (and contention).
lockup-watchdogs.txt
diff --git a/Documentation/DocBook/kernel-locking.tmpl b/Documentation/DocBook/kernel-locking.tmpl
index e584ee1..c70fd1b 100644
--- a/Documentation/DocBook/kernel-locking.tmpl
+++ b/Documentation/DocBook/kernel-locking.tmpl
@@ -1972,7 +1972,7 @@ machines due to caching.
<itemizedlist>
<listitem>
<para>
- <filename>Documentation/spinlocks.txt</filename>:
+ <filename>Documentation/locking/spinlocks.txt</filename>:
Linus Torvalds' spinlocking tutorial in the kernel sources.
</para>
</listitem>
diff --git a/Documentation/lockdep-design.txt b/Documentation/lockdep-design.txt
deleted file mode 100644
index 5dbc99c..0000000
--- a/Documentation/lockdep-design.txt
+++ /dev/null
@@ -1,286 +0,0 @@
-Runtime locking correctness validator
-=====================================
-
-started by Ingo Molnar <[email protected]>
-additions by Arjan van de Ven <[email protected]>
-
-Lock-class
-----------
-
-The basic object the validator operates upon is a 'class' of locks.
-
-A class of locks is a group of locks that are logically the same with
-respect to locking rules, even if the locks may have multiple (possibly
-tens of thousands of) instantiations. For example a lock in the inode
-struct is one class, while each inode has its own instantiation of that
-lock class.
-
-The validator tracks the 'state' of lock-classes, and it tracks
-dependencies between different lock-classes. The validator maintains a
-rolling proof that the state and the dependencies are correct.
-
-Unlike an lock instantiation, the lock-class itself never goes away: when
-a lock-class is used for the first time after bootup it gets registered,
-and all subsequent uses of that lock-class will be attached to this
-lock-class.
-
-State
------
-
-The validator tracks lock-class usage history into 4n + 1 separate state bits:
-
-- 'ever held in STATE context'
-- 'ever held as readlock in STATE context'
-- 'ever held with STATE enabled'
-- 'ever held as readlock with STATE enabled'
-
-Where STATE can be either one of (kernel/lockdep_states.h)
- - hardirq
- - softirq
- - reclaim_fs
-
-- 'ever used' [ == !unused ]
-
-When locking rules are violated, these state bits are presented in the
-locking error messages, inside curlies. A contrived example:
-
- modprobe/2287 is trying to acquire lock:
- (&sio_locks[i].lock){-.-...}, at: [<c02867fd>] mutex_lock+0x21/0x24
-
- but task is already holding lock:
- (&sio_locks[i].lock){-.-...}, at: [<c02867fd>] mutex_lock+0x21/0x24
-
-
-The bit position indicates STATE, STATE-read, for each of the states listed
-above, and the character displayed in each indicates:
-
- '.' acquired while irqs disabled and not in irq context
- '-' acquired in irq context
- '+' acquired with irqs enabled
- '?' acquired in irq context with irqs enabled.
-
-Unused mutexes cannot be part of the cause of an error.
-
-
-Single-lock state rules:
-------------------------
-
-A softirq-unsafe lock-class is automatically hardirq-unsafe as well. The
-following states are exclusive, and only one of them is allowed to be
-set for any lock-class:
-
- <hardirq-safe> and <hardirq-unsafe>
- <softirq-safe> and <softirq-unsafe>
-
-The validator detects and reports lock usage that violate these
-single-lock state rules.
-
-Multi-lock dependency rules:
-----------------------------
-
-The same lock-class must not be acquired twice, because this could lead
-to lock recursion deadlocks.
-
-Furthermore, two locks may not be taken in different order:
-
- <L1> -> <L2>
- <L2> -> <L1>
-
-because this could lead to lock inversion deadlocks. (The validator
-finds such dependencies in arbitrary complexity, i.e. there can be any
-other locking sequence between the acquire-lock operations, the
-validator will still track all dependencies between locks.)
-
-Furthermore, the following usage based lock dependencies are not allowed
-between any two lock-classes:
-
- <hardirq-safe> -> <hardirq-unsafe>
- <softirq-safe> -> <softirq-unsafe>
-
-The first rule comes from the fact the a hardirq-safe lock could be
-taken by a hardirq context, interrupting a hardirq-unsafe lock - and
-thus could result in a lock inversion deadlock. Likewise, a softirq-safe
-lock could be taken by an softirq context, interrupting a softirq-unsafe
-lock.
-
-The above rules are enforced for any locking sequence that occurs in the
-kernel: when acquiring a new lock, the validator checks whether there is
-any rule violation between the new lock and any of the held locks.
-
-When a lock-class changes its state, the following aspects of the above
-dependency rules are enforced:
-
-- if a new hardirq-safe lock is discovered, we check whether it
- took any hardirq-unsafe lock in the past.
-
-- if a new softirq-safe lock is discovered, we check whether it took
- any softirq-unsafe lock in the past.
-
-- if a new hardirq-unsafe lock is discovered, we check whether any
- hardirq-safe lock took it in the past.
-
-- if a new softirq-unsafe lock is discovered, we check whether any
- softirq-safe lock took it in the past.
-
-(Again, we do these checks too on the basis that an interrupt context
-could interrupt _any_ of the irq-unsafe or hardirq-unsafe locks, which
-could lead to a lock inversion deadlock - even if that lock scenario did
-not trigger in practice yet.)
-
-Exception: Nested data dependencies leading to nested locking
--------------------------------------------------------------
-
-There are a few cases where the Linux kernel acquires more than one
-instance of the same lock-class. Such cases typically happen when there
-is some sort of hierarchy within objects of the same type. In these
-cases there is an inherent "natural" ordering between the two objects
-(defined by the properties of the hierarchy), and the kernel grabs the
-locks in this fixed order on each of the objects.
-
-An example of such an object hierarchy that results in "nested locking"
-is that of a "whole disk" block-dev object and a "partition" block-dev
-object; the partition is "part of" the whole device and as long as one
-always takes the whole disk lock as a higher lock than the partition
-lock, the lock ordering is fully correct. The validator does not
-automatically detect this natural ordering, as the locking rule behind
-the ordering is not static.
-
-In order to teach the validator about this correct usage model, new
-versions of the various locking primitives were added that allow you to
-specify a "nesting level". An example call, for the block device mutex,
-looks like this:
-
-enum bdev_bd_mutex_lock_class
-{
- BD_MUTEX_NORMAL,
- BD_MUTEX_WHOLE,
- BD_MUTEX_PARTITION
-};
-
- mutex_lock_nested(&bdev->bd_contains->bd_mutex, BD_MUTEX_PARTITION);
-
-In this case the locking is done on a bdev object that is known to be a
-partition.
-
-The validator treats a lock that is taken in such a nested fashion as a
-separate (sub)class for the purposes of validation.
-
-Note: When changing code to use the _nested() primitives, be careful and
-check really thoroughly that the hierarchy is correctly mapped; otherwise
-you can get false positives or false negatives.
-
-Proof of 100% correctness:
---------------------------
-
-The validator achieves perfect, mathematical 'closure' (proof of locking
-correctness) in the sense that for every simple, standalone single-task
-locking sequence that occurred at least once during the lifetime of the
-kernel, the validator proves it with a 100% certainty that no
-combination and timing of these locking sequences can cause any class of
-lock related deadlock. [*]
-
-I.e. complex multi-CPU and multi-task locking scenarios do not have to
-occur in practice to prove a deadlock: only the simple 'component'
-locking chains have to occur at least once (anytime, in any
-task/context) for the validator to be able to prove correctness. (For
-example, complex deadlocks that would normally need more than 3 CPUs and
-a very unlikely constellation of tasks, irq-contexts and timings to
-occur, can be detected on a plain, lightly loaded single-CPU system as
-well!)
-
-This radically decreases the complexity of locking related QA of the
-kernel: what has to be done during QA is to trigger as many "simple"
-single-task locking dependencies in the kernel as possible, at least
-once, to prove locking correctness - instead of having to trigger every
-possible combination of locking interaction between CPUs, combined with
-every possible hardirq and softirq nesting scenario (which is impossible
-to do in practice).
-
-[*] assuming that the validator itself is 100% correct, and no other
- part of the system corrupts the state of the validator in any way.
- We also assume that all NMI/SMM paths [which could interrupt
- even hardirq-disabled codepaths] are correct and do not interfere
- with the validator. We also assume that the 64-bit 'chain hash'
- value is unique for every lock-chain in the system. Also, lock
- recursion must not be higher than 20.
-
-Performance:
-------------
-
-The above rules require _massive_ amounts of runtime checking. If we did
-that for every lock taken and for every irqs-enable event, it would
-render the system practically unusably slow. The complexity of checking
-is O(N^2), so even with just a few hundred lock-classes we'd have to do
-tens of thousands of checks for every event.
-
-This problem is solved by checking any given 'locking scenario' (unique
-sequence of locks taken after each other) only once. A simple stack of
-held locks is maintained, and a lightweight 64-bit hash value is
-calculated, which hash is unique for every lock chain. The hash value,
-when the chain is validated for the first time, is then put into a hash
-table, which hash-table can be checked in a lockfree manner. If the
-locking chain occurs again later on, the hash table tells us that we
-dont have to validate the chain again.
-
-Troubleshooting:
-----------------
-
-The validator tracks a maximum of MAX_LOCKDEP_KEYS number of lock classes.
-Exceeding this number will trigger the following lockdep warning:
-
- (DEBUG_LOCKS_WARN_ON(id >= MAX_LOCKDEP_KEYS))
-
-By default, MAX_LOCKDEP_KEYS is currently set to 8191, and typical
-desktop systems have less than 1,000 lock classes, so this warning
-normally results from lock-class leakage or failure to properly
-initialize locks. These two problems are illustrated below:
-
-1. Repeated module loading and unloading while running the validator
- will result in lock-class leakage. The issue here is that each
- load of the module will create a new set of lock classes for
- that module's locks, but module unloading does not remove old
- classes (see below discussion of reuse of lock classes for why).
- Therefore, if that module is loaded and unloaded repeatedly,
- the number of lock classes will eventually reach the maximum.
-
-2. Using structures such as arrays that have large numbers of
- locks that are not explicitly initialized. For example,
- a hash table with 8192 buckets where each bucket has its own
- spinlock_t will consume 8192 lock classes -unless- each spinlock
- is explicitly initialized at runtime, for example, using the
- run-time spin_lock_init() as opposed to compile-time initializers
- such as __SPIN_LOCK_UNLOCKED(). Failure to properly initialize
- the per-bucket spinlocks would guarantee lock-class overflow.
- In contrast, a loop that called spin_lock_init() on each lock
- would place all 8192 locks into a single lock class.
-
- The moral of this story is that you should always explicitly
- initialize your locks.
-
-One might argue that the validator should be modified to allow
-lock classes to be reused. However, if you are tempted to make this
-argument, first review the code and think through the changes that would
-be required, keeping in mind that the lock classes to be removed are
-likely to be linked into the lock-dependency graph. This turns out to
-be harder to do than to say.
-
-Of course, if you do run out of lock classes, the next thing to do is
-to find the offending lock classes. First, the following command gives
-you the number of lock classes currently in use along with the maximum:
-
- grep "lock-classes" /proc/lockdep_stats
-
-This command produces the following output on a modest system:
-
- lock-classes: 748 [max: 8191]
-
-If the number allocated (748 above) increases continually over time,
-then there is likely a leak. The following command can be used to
-identify the leaking lock classes:
-
- grep "BD" /proc/lockdep
-
-Run the command and save the output, then compare against the output from
-a later run of this command to identify the leakers. This same output
-can also help you find situations where runtime lock initialization has
-been omitted.
diff --git a/Documentation/locking/lockdep-design.txt b/Documentation/locking/lockdep-design.txt
new file mode 100644
index 0000000..5dbc99c
--- /dev/null
+++ b/Documentation/locking/lockdep-design.txt
@@ -0,0 +1,286 @@
+Runtime locking correctness validator
+=====================================
+
+started by Ingo Molnar <[email protected]>
+additions by Arjan van de Ven <[email protected]>
+
+Lock-class
+----------
+
+The basic object the validator operates upon is a 'class' of locks.
+
+A class of locks is a group of locks that are logically the same with
+respect to locking rules, even if the locks may have multiple (possibly
+tens of thousands of) instantiations. For example a lock in the inode
+struct is one class, while each inode has its own instantiation of that
+lock class.
+
+The validator tracks the 'state' of lock-classes, and it tracks
+dependencies between different lock-classes. The validator maintains a
+rolling proof that the state and the dependencies are correct.
+
+Unlike an lock instantiation, the lock-class itself never goes away: when
+a lock-class is used for the first time after bootup it gets registered,
+and all subsequent uses of that lock-class will be attached to this
+lock-class.
+
+State
+-----
+
+The validator tracks lock-class usage history into 4n + 1 separate state bits:
+
+- 'ever held in STATE context'
+- 'ever held as readlock in STATE context'
+- 'ever held with STATE enabled'
+- 'ever held as readlock with STATE enabled'
+
+Where STATE can be either one of (kernel/lockdep_states.h)
+ - hardirq
+ - softirq
+ - reclaim_fs
+
+- 'ever used' [ == !unused ]
+
+When locking rules are violated, these state bits are presented in the
+locking error messages, inside curlies. A contrived example:
+
+ modprobe/2287 is trying to acquire lock:
+ (&sio_locks[i].lock){-.-...}, at: [<c02867fd>] mutex_lock+0x21/0x24
+
+ but task is already holding lock:
+ (&sio_locks[i].lock){-.-...}, at: [<c02867fd>] mutex_lock+0x21/0x24
+
+
+The bit position indicates STATE, STATE-read, for each of the states listed
+above, and the character displayed in each indicates:
+
+ '.' acquired while irqs disabled and not in irq context
+ '-' acquired in irq context
+ '+' acquired with irqs enabled
+ '?' acquired in irq context with irqs enabled.
+
+Unused mutexes cannot be part of the cause of an error.
+
+
+Single-lock state rules:
+------------------------
+
+A softirq-unsafe lock-class is automatically hardirq-unsafe as well. The
+following states are exclusive, and only one of them is allowed to be
+set for any lock-class:
+
+ <hardirq-safe> and <hardirq-unsafe>
+ <softirq-safe> and <softirq-unsafe>
+
+The validator detects and reports lock usage that violate these
+single-lock state rules.
+
+Multi-lock dependency rules:
+----------------------------
+
+The same lock-class must not be acquired twice, because this could lead
+to lock recursion deadlocks.
+
+Furthermore, two locks may not be taken in different order:
+
+ <L1> -> <L2>
+ <L2> -> <L1>
+
+because this could lead to lock inversion deadlocks. (The validator
+finds such dependencies in arbitrary complexity, i.e. there can be any
+other locking sequence between the acquire-lock operations, the
+validator will still track all dependencies between locks.)
+
+Furthermore, the following usage based lock dependencies are not allowed
+between any two lock-classes:
+
+ <hardirq-safe> -> <hardirq-unsafe>
+ <softirq-safe> -> <softirq-unsafe>
+
+The first rule comes from the fact the a hardirq-safe lock could be
+taken by a hardirq context, interrupting a hardirq-unsafe lock - and
+thus could result in a lock inversion deadlock. Likewise, a softirq-safe
+lock could be taken by an softirq context, interrupting a softirq-unsafe
+lock.
+
+The above rules are enforced for any locking sequence that occurs in the
+kernel: when acquiring a new lock, the validator checks whether there is
+any rule violation between the new lock and any of the held locks.
+
+When a lock-class changes its state, the following aspects of the above
+dependency rules are enforced:
+
+- if a new hardirq-safe lock is discovered, we check whether it
+ took any hardirq-unsafe lock in the past.
+
+- if a new softirq-safe lock is discovered, we check whether it took
+ any softirq-unsafe lock in the past.
+
+- if a new hardirq-unsafe lock is discovered, we check whether any
+ hardirq-safe lock took it in the past.
+
+- if a new softirq-unsafe lock is discovered, we check whether any
+ softirq-safe lock took it in the past.
+
+(Again, we do these checks too on the basis that an interrupt context
+could interrupt _any_ of the irq-unsafe or hardirq-unsafe locks, which
+could lead to a lock inversion deadlock - even if that lock scenario did
+not trigger in practice yet.)
+
+Exception: Nested data dependencies leading to nested locking
+-------------------------------------------------------------
+
+There are a few cases where the Linux kernel acquires more than one
+instance of the same lock-class. Such cases typically happen when there
+is some sort of hierarchy within objects of the same type. In these
+cases there is an inherent "natural" ordering between the two objects
+(defined by the properties of the hierarchy), and the kernel grabs the
+locks in this fixed order on each of the objects.
+
+An example of such an object hierarchy that results in "nested locking"
+is that of a "whole disk" block-dev object and a "partition" block-dev
+object; the partition is "part of" the whole device and as long as one
+always takes the whole disk lock as a higher lock than the partition
+lock, the lock ordering is fully correct. The validator does not
+automatically detect this natural ordering, as the locking rule behind
+the ordering is not static.
+
+In order to teach the validator about this correct usage model, new
+versions of the various locking primitives were added that allow you to
+specify a "nesting level". An example call, for the block device mutex,
+looks like this:
+
+enum bdev_bd_mutex_lock_class
+{
+ BD_MUTEX_NORMAL,
+ BD_MUTEX_WHOLE,
+ BD_MUTEX_PARTITION
+};
+
+ mutex_lock_nested(&bdev->bd_contains->bd_mutex, BD_MUTEX_PARTITION);
+
+In this case the locking is done on a bdev object that is known to be a
+partition.
+
+The validator treats a lock that is taken in such a nested fashion as a
+separate (sub)class for the purposes of validation.
+
+Note: When changing code to use the _nested() primitives, be careful and
+check really thoroughly that the hierarchy is correctly mapped; otherwise
+you can get false positives or false negatives.
+
+Proof of 100% correctness:
+--------------------------
+
+The validator achieves perfect, mathematical 'closure' (proof of locking
+correctness) in the sense that for every simple, standalone single-task
+locking sequence that occurred at least once during the lifetime of the
+kernel, the validator proves it with a 100% certainty that no
+combination and timing of these locking sequences can cause any class of
+lock related deadlock. [*]
+
+I.e. complex multi-CPU and multi-task locking scenarios do not have to
+occur in practice to prove a deadlock: only the simple 'component'
+locking chains have to occur at least once (anytime, in any
+task/context) for the validator to be able to prove correctness. (For
+example, complex deadlocks that would normally need more than 3 CPUs and
+a very unlikely constellation of tasks, irq-contexts and timings to
+occur, can be detected on a plain, lightly loaded single-CPU system as
+well!)
+
+This radically decreases the complexity of locking related QA of the
+kernel: what has to be done during QA is to trigger as many "simple"
+single-task locking dependencies in the kernel as possible, at least
+once, to prove locking correctness - instead of having to trigger every
+possible combination of locking interaction between CPUs, combined with
+every possible hardirq and softirq nesting scenario (which is impossible
+to do in practice).
+
+[*] assuming that the validator itself is 100% correct, and no other
+ part of the system corrupts the state of the validator in any way.
+ We also assume that all NMI/SMM paths [which could interrupt
+ even hardirq-disabled codepaths] are correct and do not interfere
+ with the validator. We also assume that the 64-bit 'chain hash'
+ value is unique for every lock-chain in the system. Also, lock
+ recursion must not be higher than 20.
+
+Performance:
+------------
+
+The above rules require _massive_ amounts of runtime checking. If we did
+that for every lock taken and for every irqs-enable event, it would
+render the system practically unusably slow. The complexity of checking
+is O(N^2), so even with just a few hundred lock-classes we'd have to do
+tens of thousands of checks for every event.
+
+This problem is solved by checking any given 'locking scenario' (unique
+sequence of locks taken after each other) only once. A simple stack of
+held locks is maintained, and a lightweight 64-bit hash value is
+calculated, which hash is unique for every lock chain. The hash value,
+when the chain is validated for the first time, is then put into a hash
+table, which hash-table can be checked in a lockfree manner. If the
+locking chain occurs again later on, the hash table tells us that we
+dont have to validate the chain again.
+
+Troubleshooting:
+----------------
+
+The validator tracks a maximum of MAX_LOCKDEP_KEYS number of lock classes.
+Exceeding this number will trigger the following lockdep warning:
+
+ (DEBUG_LOCKS_WARN_ON(id >= MAX_LOCKDEP_KEYS))
+
+By default, MAX_LOCKDEP_KEYS is currently set to 8191, and typical
+desktop systems have less than 1,000 lock classes, so this warning
+normally results from lock-class leakage or failure to properly
+initialize locks. These two problems are illustrated below:
+
+1. Repeated module loading and unloading while running the validator
+ will result in lock-class leakage. The issue here is that each
+ load of the module will create a new set of lock classes for
+ that module's locks, but module unloading does not remove old
+ classes (see below discussion of reuse of lock classes for why).
+ Therefore, if that module is loaded and unloaded repeatedly,
+ the number of lock classes will eventually reach the maximum.
+
+2. Using structures such as arrays that have large numbers of
+ locks that are not explicitly initialized. For example,
+ a hash table with 8192 buckets where each bucket has its own
+ spinlock_t will consume 8192 lock classes -unless- each spinlock
+ is explicitly initialized at runtime, for example, using the
+ run-time spin_lock_init() as opposed to compile-time initializers
+ such as __SPIN_LOCK_UNLOCKED(). Failure to properly initialize
+ the per-bucket spinlocks would guarantee lock-class overflow.
+ In contrast, a loop that called spin_lock_init() on each lock
+ would place all 8192 locks into a single lock class.
+
+ The moral of this story is that you should always explicitly
+ initialize your locks.
+
+One might argue that the validator should be modified to allow
+lock classes to be reused. However, if you are tempted to make this
+argument, first review the code and think through the changes that would
+be required, keeping in mind that the lock classes to be removed are
+likely to be linked into the lock-dependency graph. This turns out to
+be harder to do than to say.
+
+Of course, if you do run out of lock classes, the next thing to do is
+to find the offending lock classes. First, the following command gives
+you the number of lock classes currently in use along with the maximum:
+
+ grep "lock-classes" /proc/lockdep_stats
+
+This command produces the following output on a modest system:
+
+ lock-classes: 748 [max: 8191]
+
+If the number allocated (748 above) increases continually over time,
+then there is likely a leak. The following command can be used to
+identify the leaking lock classes:
+
+ grep "BD" /proc/lockdep
+
+Run the command and save the output, then compare against the output from
+a later run of this command to identify the leakers. This same output
+can also help you find situations where runtime lock initialization has
+been omitted.
diff --git a/Documentation/locking/lockstat.txt b/Documentation/locking/lockstat.txt
new file mode 100644
index 0000000..7428773
--- /dev/null
+++ b/Documentation/locking/lockstat.txt
@@ -0,0 +1,178 @@
+
+LOCK STATISTICS
+
+- WHAT
+
+As the name suggests, it provides statistics on locks.
+
+- WHY
+
+Because things like lock contention can severely impact performance.
+
+- HOW
+
+Lockdep already has hooks in the lock functions and maps lock instances to
+lock classes. We build on that (see Documentation/lokcing/lockdep-design.txt).
+The graph below shows the relation between the lock functions and the various
+hooks therein.
+
+ __acquire
+ |
+ lock _____
+ | \
+ | __contended
+ | |
+ | <wait>
+ | _______/
+ |/
+ |
+ __acquired
+ |
+ .
+ <hold>
+ .
+ |
+ __release
+ |
+ unlock
+
+lock, unlock - the regular lock functions
+__* - the hooks
+<> - states
+
+With these hooks we provide the following statistics:
+
+ con-bounces - number of lock contention that involved x-cpu data
+ contentions - number of lock acquisitions that had to wait
+ wait time min - shortest (non-0) time we ever had to wait for a lock
+ max - longest time we ever had to wait for a lock
+ total - total time we spend waiting on this lock
+ avg - average time spent waiting on this lock
+ acq-bounces - number of lock acquisitions that involved x-cpu data
+ acquisitions - number of times we took the lock
+ hold time min - shortest (non-0) time we ever held the lock
+ max - longest time we ever held the lock
+ total - total time this lock was held
+ avg - average time this lock was held
+
+These numbers are gathered per lock class, per read/write state (when
+applicable).
+
+It also tracks 4 contention points per class. A contention point is a call site
+that had to wait on lock acquisition.
+
+ - CONFIGURATION
+
+Lock statistics are enabled via CONFIG_LOCK_STAT.
+
+ - USAGE
+
+Enable collection of statistics:
+
+# echo 1 >/proc/sys/kernel/lock_stat
+
+Disable collection of statistics:
+
+# echo 0 >/proc/sys/kernel/lock_stat
+
+Look at the current lock statistics:
+
+( line numbers not part of actual output, done for clarity in the explanation
+ below )
+
+# less /proc/lock_stat
+
+01 lock_stat version 0.4
+02-----------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------
+03 class name con-bounces contentions waittime-min waittime-max waittime-total waittime-avg acq-bounces acquisitions holdtime-min holdtime-max holdtime-total holdtime-avg
+04-----------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------
+05
+06 &mm->mmap_sem-W: 46 84 0.26 939.10 16371.53 194.90 47291 2922365 0.16 2220301.69 17464026916.32 5975.99
+07 &mm->mmap_sem-R: 37 100 1.31 299502.61 325629.52 3256.30 212344 34316685 0.10 7744.91 95016910.20 2.77
+08 ---------------
+09 &mm->mmap_sem 1 [<ffffffff811502a7>] khugepaged_scan_mm_slot+0x57/0x280
+19 &mm->mmap_sem 96 [<ffffffff815351c4>] __do_page_fault+0x1d4/0x510
+11 &mm->mmap_sem 34 [<ffffffff81113d77>] vm_mmap_pgoff+0x87/0xd0
+12 &mm->mmap_sem 17 [<ffffffff81127e71>] vm_munmap+0x41/0x80
+13 ---------------
+14 &mm->mmap_sem 1 [<ffffffff81046fda>] dup_mmap+0x2a/0x3f0
+15 &mm->mmap_sem 60 [<ffffffff81129e29>] SyS_mprotect+0xe9/0x250
+16 &mm->mmap_sem 41 [<ffffffff815351c4>] __do_page_fault+0x1d4/0x510
+17 &mm->mmap_sem 68 [<ffffffff81113d77>] vm_mmap_pgoff+0x87/0xd0
+18
+19.............................................................................................................................................................................................................................
+20
+21 unix_table_lock: 110 112 0.21 49.24 163.91 1.46 21094 66312 0.12 624.42 31589.81 0.48
+22 ---------------
+23 unix_table_lock 45 [<ffffffff8150ad8e>] unix_create1+0x16e/0x1b0
+24 unix_table_lock 47 [<ffffffff8150b111>] unix_release_sock+0x31/0x250
+25 unix_table_lock 15 [<ffffffff8150ca37>] unix_find_other+0x117/0x230
+26 unix_table_lock 5 [<ffffffff8150a09f>] unix_autobind+0x11f/0x1b0
+27 ---------------
+28 unix_table_lock 39 [<ffffffff8150b111>] unix_release_sock+0x31/0x250
+29 unix_table_lock 49 [<ffffffff8150ad8e>] unix_create1+0x16e/0x1b0
+30 unix_table_lock 20 [<ffffffff8150ca37>] unix_find_other+0x117/0x230
+31 unix_table_lock 4 [<ffffffff8150a09f>] unix_autobind+0x11f/0x1b0
+
+
+This excerpt shows the first two lock class statistics. Line 01 shows the
+output version - each time the format changes this will be updated. Line 02-04
+show the header with column descriptions. Lines 05-18 and 20-31 show the actual
+statistics. These statistics come in two parts; the actual stats separated by a
+short separator (line 08, 13) from the contention points.
+
+The first lock (05-18) is a read/write lock, and shows two lines above the
+short separator. The contention points don't match the column descriptors,
+they have two: contentions and [<IP>] symbol. The second set of contention
+points are the points we're contending with.
+
+The integer part of the time values is in us.
+
+Dealing with nested locks, subclasses may appear:
+
+32...........................................................................................................................................................................................................................
+33
+34 &rq->lock: 13128 13128 0.43 190.53 103881.26 7.91 97454 3453404 0.00 401.11 13224683.11 3.82
+35 ---------
+36 &rq->lock 645 [<ffffffff8103bfc4>] task_rq_lock+0x43/0x75
+37 &rq->lock 297 [<ffffffff8104ba65>] try_to_wake_up+0x127/0x25a
+38 &rq->lock 360 [<ffffffff8103c4c5>] select_task_rq_fair+0x1f0/0x74a
+39 &rq->lock 428 [<ffffffff81045f98>] scheduler_tick+0x46/0x1fb
+40 ---------
+41 &rq->lock 77 [<ffffffff8103bfc4>] task_rq_lock+0x43/0x75
+42 &rq->lock 174 [<ffffffff8104ba65>] try_to_wake_up+0x127/0x25a
+43 &rq->lock 4715 [<ffffffff8103ed4b>] double_rq_lock+0x42/0x54
+44 &rq->lock 893 [<ffffffff81340524>] schedule+0x157/0x7b8
+45
+46...........................................................................................................................................................................................................................
+47
+48 &rq->lock/1: 1526 11488 0.33 388.73 136294.31 11.86 21461 38404 0.00 37.93 109388.53 2.84
+49 -----------
+50 &rq->lock/1 11526 [<ffffffff8103ed58>] double_rq_lock+0x4f/0x54
+51 -----------
+52 &rq->lock/1 5645 [<ffffffff8103ed4b>] double_rq_lock+0x42/0x54
+53 &rq->lock/1 1224 [<ffffffff81340524>] schedule+0x157/0x7b8
+54 &rq->lock/1 4336 [<ffffffff8103ed58>] double_rq_lock+0x4f/0x54
+55 &rq->lock/1 181 [<ffffffff8104ba65>] try_to_wake_up+0x127/0x25a
+
+Line 48 shows statistics for the second subclass (/1) of &rq->lock class
+(subclass starts from 0), since in this case, as line 50 suggests,
+double_rq_lock actually acquires a nested lock of two spinlocks.
+
+View the top contending locks:
+
+# grep : /proc/lock_stat | head
+ clockevents_lock: 2926159 2947636 0.15 46882.81 1784540466.34 605.41 3381345 3879161 0.00 2260.97 53178395.68 13.71
+ tick_broadcast_lock: 346460 346717 0.18 2257.43 39364622.71 113.54 3642919 4242696 0.00 2263.79 49173646.60 11.59
+ &mapping->i_mmap_mutex: 203896 203899 3.36 645530.05 31767507988.39 155800.21 3361776 8893984 0.17 2254.15 14110121.02 1.59
+ &rq->lock: 135014 136909 0.18 606.09 842160.68 6.15 1540728 10436146 0.00 728.72 17606683.41 1.69
+ &(&zone->lru_lock)->rlock: 93000 94934 0.16 59.18 188253.78 1.98 1199912 3809894 0.15 391.40 3559518.81 0.93
+ tasklist_lock-W: 40667 41130 0.23 1189.42 428980.51 10.43 270278 510106 0.16 653.51 3939674.91 7.72
+ tasklist_lock-R: 21298 21305 0.20 1310.05 215511.12 10.12 186204 241258 0.14 1162.33 1179779.23 4.89
+ rcu_node_1: 47656 49022 0.16 635.41 193616.41 3.95 844888 1865423 0.00 764.26 1656226.96 0.89
+ &(&dentry->d_lockref.lock)->rlock: 39791 40179 0.15 1302.08 88851.96 2.21 2790851 12527025 0.10 1910.75 3379714.27 0.27
+ rcu_node_0: 29203 30064 0.16 786.55 1555573.00 51.74 88963 244254 0.00 398.87 428872.51 1.76
+
+Clear the statistics:
+
+# echo 0 > /proc/lock_stat
diff --git a/Documentation/locking/mutex-design.txt b/Documentation/locking/mutex-design.txt
new file mode 100644
index 0000000..ee231ed
--- /dev/null
+++ b/Documentation/locking/mutex-design.txt
@@ -0,0 +1,157 @@
+Generic Mutex Subsystem
+
+started by Ingo Molnar <[email protected]>
+updated by Davidlohr Bueso <[email protected]>
+
+What are mutexes?
+-----------------
+
+In the Linux kernel, mutexes refer to a particular locking primitive
+that enforces serialization on shared memory systems, and not only to
+the generic term referring to 'mutual exclusion' found in academia
+or similar theoretical text books. Mutexes are sleeping locks which
+behave similarly to binary semaphores, and were introduced in 2006[1]
+as an alternative to these. This new data structure provided a number
+of advantages, including simpler interfaces, and at that time smaller
+code (see Disadvantages).
+
+[1] http://lwn.net/Articles/164802/
+
+Implementation
+--------------
+
+Mutexes are represented by 'struct mutex', defined in include/linux/mutex.h
+and implemented in kernel/locking/mutex.c. These locks use a three
+state atomic counter (->count) to represent the different possible
+transitions that can occur during the lifetime of a lock:
+
+ 1: unlocked
+ 0: locked, no waiters
+ negative: locked, with potential waiters
+
+In its most basic form it also includes a wait-queue and a spinlock
+that serializes access to it. CONFIG_SMP systems can also include
+a pointer to the lock task owner (->owner) as well as a spinner MCS
+lock (->osq), both described below in (ii).
+
+When acquiring a mutex, there are three possible paths that can be
+taken, depending on the state of the lock:
+
+(i) fastpath: tries to atomically acquire the lock by decrementing the
+ counter. If it was already taken by another task it goes to the next
+ possible path. This logic is architecture specific. On x86-64, the
+ locking fastpath is 2 instructions:
+
+ 0000000000000e10 <mutex_lock>:
+ e21: f0 ff 0b lock decl (%rbx)
+ e24: 79 08 jns e2e <mutex_lock+0x1e>
+
+ the unlocking fastpath is equally tight:
+
+ 0000000000000bc0 <mutex_unlock>:
+ bc8: f0 ff 07 lock incl (%rdi)
+ bcb: 7f 0a jg bd7 <mutex_unlock+0x17>
+
+
+(ii) midpath: aka optimistic spinning, tries to spin for acquisition
+ while the lock owner is running and there are no other tasks ready
+ to run that have higher priority (need_resched). The rationale is
+ that if the lock owner is running, it is likely to release the lock
+ soon. The mutex spinners are queued up using MCS lock so that only
+ one spinner can compete for the mutex.
+
+ The MCS lock (proposed by Mellor-Crummey and Scott) is a simple spinlock
+ with the desirable properties of being fair and with each cpu trying
+ to acquire the lock spinning on a local variable. It avoids expensive
+ cacheline bouncing that common test-and-set spinlock implementations
+ incur. An MCS-like lock is specially tailored for optimistic spinning
+ for sleeping lock implementation. An important feature of the customized
+ MCS lock is that it has the extra property that spinners are able to exit
+ the MCS spinlock queue when they need to reschedule. This further helps
+ avoid situations where MCS spinners that need to reschedule would continue
+ waiting to spin on mutex owner, only to go directly to slowpath upon
+ obtaining the MCS lock.
+
+
+(iii) slowpath: last resort, if the lock is still unable to be acquired,
+ the task is added to the wait-queue and sleeps until woken up by the
+ unlock path. Under normal circumstances it blocks as TASK_UNINTERRUPTIBLE.
+
+While formally kernel mutexes are sleepable locks, it is path (ii) that
+makes them more practically a hybrid type. By simply not interrupting a
+task and busy-waiting for a few cycles instead of immediately sleeping,
+the performance of this lock has been seen to significantly improve a
+number of workloads. Note that this technique is also used for rw-semaphores.
+
+Semantics
+---------
+
+The mutex subsystem checks and enforces the following rules:
+
+ - Only one task can hold the mutex at a time.
+ - Only the owner can unlock the mutex.
+ - Multiple unlocks are not permitted.
+ - Recursive locking/unlocking is not permitted.
+ - A mutex must only be initialized via the API (see below).
+ - A task may not exit with a mutex held.
+ - Memory areas where held locks reside must not be freed.
+ - Held mutexes must not be reinitialized.
+ - Mutexes may not be used in hardware or software interrupt
+ contexts such as tasklets and timers.
+
+These semantics are fully enforced when CONFIG DEBUG_MUTEXES is enabled.
+In addition, the mutex debugging code also implements a number of other
+features that make lock debugging easier and faster:
+
+ - Uses symbolic names of mutexes, whenever they are printed
+ in debug output.
+ - Point-of-acquire tracking, symbolic lookup of function names,
+ list of all locks held in the system, printout of them.
+ - Owner tracking.
+ - Detects self-recursing locks and prints out all relevant info.
+ - Detects multi-task circular deadlocks and prints out all affected
+ locks and tasks (and only those tasks).
+
+
+Interfaces
+----------
+Statically define the mutex:
+ DEFINE_MUTEX(name);
+
+Dynamically initialize the mutex:
+ mutex_init(mutex);
+
+Acquire the mutex, uninterruptible:
+ void mutex_lock(struct mutex *lock);
+ void mutex_lock_nested(struct mutex *lock, unsigned int subclass);
+ int mutex_trylock(struct mutex *lock);
+
+Acquire the mutex, interruptible:
+ int mutex_lock_interruptible_nested(struct mutex *lock,
+ unsigned int subclass);
+ int mutex_lock_interruptible(struct mutex *lock);
+
+Acquire the mutex, interruptible, if dec to 0:
+ int atomic_dec_and_mutex_lock(atomic_t *cnt, struct mutex *lock);
+
+Unlock the mutex:
+ void mutex_unlock(struct mutex *lock);
+
+Test if the mutex is taken:
+ int mutex_is_locked(struct mutex *lock);
+
+Disadvantages
+-------------
+
+Unlike its original design and purpose, 'struct mutex' is larger than
+most locks in the kernel. E.g: on x86-64 it is 40 bytes, almost twice
+as large as 'struct semaphore' (24 bytes) and 8 bytes shy of the
+'struct rw_semaphore' variant. Larger structure sizes mean more CPU
+cache and memory footprint.
+
+When to use mutexes
+-------------------
+
+Unless the strict semantics of mutexes are unsuitable and/or the critical
+region prevents the lock from being shared, always prefer them to any other
+locking primitive.
diff --git a/Documentation/locking/rt-mutex-design.txt b/Documentation/locking/rt-mutex-design.txt
new file mode 100644
index 0000000..8666070
--- /dev/null
+++ b/Documentation/locking/rt-mutex-design.txt
@@ -0,0 +1,781 @@
+#
+# Copyright (c) 2006 Steven Rostedt
+# Licensed under the GNU Free Documentation License, Version 1.2
+#
+
+RT-mutex implementation design
+------------------------------
+
+This document tries to describe the design of the rtmutex.c implementation.
+It doesn't describe the reasons why rtmutex.c exists. For that please see
+Documentation/rt-mutex.txt. Although this document does explain problems
+that happen without this code, but that is in the concept to understand
+what the code actually is doing.
+
+The goal of this document is to help others understand the priority
+inheritance (PI) algorithm that is used, as well as reasons for the
+decisions that were made to implement PI in the manner that was done.
+
+
+Unbounded Priority Inversion
+----------------------------
+
+Priority inversion is when a lower priority process executes while a higher
+priority process wants to run. This happens for several reasons, and
+most of the time it can't be helped. Anytime a high priority process wants
+to use a resource that a lower priority process has (a mutex for example),
+the high priority process must wait until the lower priority process is done
+with the resource. This is a priority inversion. What we want to prevent
+is something called unbounded priority inversion. That is when the high
+priority process is prevented from running by a lower priority process for
+an undetermined amount of time.
+
+The classic example of unbounded priority inversion is where you have three
+processes, let's call them processes A, B, and C, where A is the highest
+priority process, C is the lowest, and B is in between. A tries to grab a lock
+that C owns and must wait and lets C run to release the lock. But in the
+meantime, B executes, and since B is of a higher priority than C, it preempts C,
+but by doing so, it is in fact preempting A which is a higher priority process.
+Now there's no way of knowing how long A will be sleeping waiting for C
+to release the lock, because for all we know, B is a CPU hog and will
+never give C a chance to release the lock. This is called unbounded priority
+inversion.
+
+Here's a little ASCII art to show the problem.
+
+ grab lock L1 (owned by C)
+ |
+A ---+
+ C preempted by B
+ |
+C +----+
+
+B +-------->
+ B now keeps A from running.
+
+
+Priority Inheritance (PI)
+-------------------------
+
+There are several ways to solve this issue, but other ways are out of scope
+for this document. Here we only discuss PI.
+
+PI is where a process inherits the priority of another process if the other
+process blocks on a lock owned by the current process. To make this easier
+to understand, let's use the previous example, with processes A, B, and C again.
+
+This time, when A blocks on the lock owned by C, C would inherit the priority
+of A. So now if B becomes runnable, it would not preempt C, since C now has
+the high priority of A. As soon as C releases the lock, it loses its
+inherited priority, and A then can continue with the resource that C had.
+
+Terminology
+-----------
+
+Here I explain some terminology that is used in this document to help describe
+the design that is used to implement PI.
+
+PI chain - The PI chain is an ordered series of locks and processes that cause
+ processes to inherit priorities from a previous process that is
+ blocked on one of its locks. This is described in more detail
+ later in this document.
+
+mutex - In this document, to differentiate from locks that implement
+ PI and spin locks that are used in the PI code, from now on
+ the PI locks will be called a mutex.
+
+lock - In this document from now on, I will use the term lock when
+ referring to spin locks that are used to protect parts of the PI
+ algorithm. These locks disable preemption for UP (when
+ CONFIG_PREEMPT is enabled) and on SMP prevents multiple CPUs from
+ entering critical sections simultaneously.
+
+spin lock - Same as lock above.
+
+waiter - A waiter is a struct that is stored on the stack of a blocked
+ process. Since the scope of the waiter is within the code for
+ a process being blocked on the mutex, it is fine to allocate
+ the waiter on the process's stack (local variable). This
+ structure holds a pointer to the task, as well as the mutex that
+ the task is blocked on. It also has the plist node structures to
+ place the task in the waiter_list of a mutex as well as the
+ pi_list of a mutex owner task (described below).
+
+ waiter is sometimes used in reference to the task that is waiting
+ on a mutex. This is the same as waiter->task.
+
+waiters - A list of processes that are blocked on a mutex.
+
+top waiter - The highest priority process waiting on a specific mutex.
+
+top pi waiter - The highest priority process waiting on one of the mutexes
+ that a specific process owns.
+
+Note: task and process are used interchangeably in this document, mostly to
+ differentiate between two processes that are being described together.
+
+
+PI chain
+--------
+
+The PI chain is a list of processes and mutexes that may cause priority
+inheritance to take place. Multiple chains may converge, but a chain
+would never diverge, since a process can't be blocked on more than one
+mutex at a time.
+
+Example:
+
+ Process: A, B, C, D, E
+ Mutexes: L1, L2, L3, L4
+
+ A owns: L1
+ B blocked on L1
+ B owns L2
+ C blocked on L2
+ C owns L3
+ D blocked on L3
+ D owns L4
+ E blocked on L4
+
+The chain would be:
+
+ E->L4->D->L3->C->L2->B->L1->A
+
+To show where two chains merge, we could add another process F and
+another mutex L5 where B owns L5 and F is blocked on mutex L5.
+
+The chain for F would be:
+
+ F->L5->B->L1->A
+
+Since a process may own more than one mutex, but never be blocked on more than
+one, the chains merge.
+
+Here we show both chains:
+
+ E->L4->D->L3->C->L2-+
+ |
+ +->B->L1->A
+ |
+ F->L5-+
+
+For PI to work, the processes at the right end of these chains (or we may
+also call it the Top of the chain) must be equal to or higher in priority
+than the processes to the left or below in the chain.
+
+Also since a mutex may have more than one process blocked on it, we can
+have multiple chains merge at mutexes. If we add another process G that is
+blocked on mutex L2:
+
+ G->L2->B->L1->A
+
+And once again, to show how this can grow I will show the merging chains
+again.
+
+ E->L4->D->L3->C-+
+ +->L2-+
+ | |
+ G-+ +->B->L1->A
+ |
+ F->L5-+
+
+
+Plist
+-----
+
+Before I go further and talk about how the PI chain is stored through lists
+on both mutexes and processes, I'll explain the plist. This is similar to
+the struct list_head functionality that is already in the kernel.
+The implementation of plist is out of scope for this document, but it is
+very important to understand what it does.
+
+There are a few differences between plist and list, the most important one
+being that plist is a priority sorted linked list. This means that the
+priorities of the plist are sorted, such that it takes O(1) to retrieve the
+highest priority item in the list. Obviously this is useful to store processes
+based on their priorities.
+
+Another difference, which is important for implementation, is that, unlike
+list, the head of the list is a different element than the nodes of a list.
+So the head of the list is declared as struct plist_head and nodes that will
+be added to the list are declared as struct plist_node.
+
+
+Mutex Waiter List
+-----------------
+
+Every mutex keeps track of all the waiters that are blocked on itself. The mutex
+has a plist to store these waiters by priority. This list is protected by
+a spin lock that is located in the struct of the mutex. This lock is called
+wait_lock. Since the modification of the waiter list is never done in
+interrupt context, the wait_lock can be taken without disabling interrupts.
+
+
+Task PI List
+------------
+
+To keep track of the PI chains, each process has its own PI list. This is
+a list of all top waiters of the mutexes that are owned by the process.
+Note that this list only holds the top waiters and not all waiters that are
+blocked on mutexes owned by the process.
+
+The top of the task's PI list is always the highest priority task that
+is waiting on a mutex that is owned by the task. So if the task has
+inherited a priority, it will always be the priority of the task that is
+at the top of this list.
+
+This list is stored in the task structure of a process as a plist called
+pi_list. This list is protected by a spin lock also in the task structure,
+called pi_lock. This lock may also be taken in interrupt context, so when
+locking the pi_lock, interrupts must be disabled.
+
+
+Depth of the PI Chain
+---------------------
+
+The maximum depth of the PI chain is not dynamic, and could actually be
+defined. But is very complex to figure it out, since it depends on all
+the nesting of mutexes. Let's look at the example where we have 3 mutexes,
+L1, L2, and L3, and four separate functions func1, func2, func3 and func4.
+The following shows a locking order of L1->L2->L3, but may not actually
+be directly nested that way.
+
+void func1(void)
+{
+ mutex_lock(L1);
+
+ /* do anything */
+
+ mutex_unlock(L1);
+}
+
+void func2(void)
+{
+ mutex_lock(L1);
+ mutex_lock(L2);
+
+ /* do something */
+
+ mutex_unlock(L2);
+ mutex_unlock(L1);
+}
+
+void func3(void)
+{
+ mutex_lock(L2);
+ mutex_lock(L3);
+
+ /* do something else */
+
+ mutex_unlock(L3);
+ mutex_unlock(L2);
+}
+
+void func4(void)
+{
+ mutex_lock(L3);
+
+ /* do something again */
+
+ mutex_unlock(L3);
+}
+
+Now we add 4 processes that run each of these functions separately.
+Processes A, B, C, and D which run functions func1, func2, func3 and func4
+respectively, and such that D runs first and A last. With D being preempted
+in func4 in the "do something again" area, we have a locking that follows:
+
+D owns L3
+ C blocked on L3
+ C owns L2
+ B blocked on L2
+ B owns L1
+ A blocked on L1
+
+And thus we have the chain A->L1->B->L2->C->L3->D.
+
+This gives us a PI depth of 4 (four processes), but looking at any of the
+functions individually, it seems as though they only have at most a locking
+depth of two. So, although the locking depth is defined at compile time,
+it still is very difficult to find the possibilities of that depth.
+
+Now since mutexes can be defined by user-land applications, we don't want a DOS
+type of application that nests large amounts of mutexes to create a large
+PI chain, and have the code holding spin locks while looking at a large
+amount of data. So to prevent this, the implementation not only implements
+a maximum lock depth, but also only holds at most two different locks at a
+time, as it walks the PI chain. More about this below.
+
+
+Mutex owner and flags
+---------------------
+
+The mutex structure contains a pointer to the owner of the mutex. If the
+mutex is not owned, this owner is set to NULL. Since all architectures
+have the task structure on at least a four byte alignment (and if this is
+not true, the rtmutex.c code will be broken!), this allows for the two
+least significant bits to be used as flags. This part is also described
+in Documentation/rt-mutex.txt, but will also be briefly described here.
+
+Bit 0 is used as the "Pending Owner" flag. This is described later.
+Bit 1 is used as the "Has Waiters" flags. This is also described later
+ in more detail, but is set whenever there are waiters on a mutex.
+
+
+cmpxchg Tricks
+--------------
+
+Some architectures implement an atomic cmpxchg (Compare and Exchange). This
+is used (when applicable) to keep the fast path of grabbing and releasing
+mutexes short.
+
+cmpxchg is basically the following function performed atomically:
+
+unsigned long _cmpxchg(unsigned long *A, unsigned long *B, unsigned long *C)
+{
+ unsigned long T = *A;
+ if (*A == *B) {
+ *A = *C;
+ }
+ return T;
+}
+#define cmpxchg(a,b,c) _cmpxchg(&a,&b,&c)
+
+This is really nice to have, since it allows you to only update a variable
+if the variable is what you expect it to be. You know if it succeeded if
+the return value (the old value of A) is equal to B.
+
+The macro rt_mutex_cmpxchg is used to try to lock and unlock mutexes. If
+the architecture does not support CMPXCHG, then this macro is simply set
+to fail every time. But if CMPXCHG is supported, then this will
+help out extremely to keep the fast path short.
+
+The use of rt_mutex_cmpxchg with the flags in the owner field help optimize
+the system for architectures that support it. This will also be explained
+later in this document.
+
+
+Priority adjustments
+--------------------
+
+The implementation of the PI code in rtmutex.c has several places that a
+process must adjust its priority. With the help of the pi_list of a
+process this is rather easy to know what needs to be adjusted.
+
+The functions implementing the task adjustments are rt_mutex_adjust_prio,
+__rt_mutex_adjust_prio (same as the former, but expects the task pi_lock
+to already be taken), rt_mutex_getprio, and rt_mutex_setprio.
+
+rt_mutex_getprio and rt_mutex_setprio are only used in __rt_mutex_adjust_prio.
+
+rt_mutex_getprio returns the priority that the task should have. Either the
+task's own normal priority, or if a process of a higher priority is waiting on
+a mutex owned by the task, then that higher priority should be returned.
+Since the pi_list of a task holds an order by priority list of all the top
+waiters of all the mutexes that the task owns, rt_mutex_getprio simply needs
+to compare the top pi waiter to its own normal priority, and return the higher
+priority back.
+
+(Note: if looking at the code, you will notice that the lower number of
+ prio is returned. This is because the prio field in the task structure
+ is an inverse order of the actual priority. So a "prio" of 5 is
+ of higher priority than a "prio" of 10.)
+
+__rt_mutex_adjust_prio examines the result of rt_mutex_getprio, and if the
+result does not equal the task's current priority, then rt_mutex_setprio
+is called to adjust the priority of the task to the new priority.
+Note that rt_mutex_setprio is defined in kernel/sched/core.c to implement the
+actual change in priority.
+
+It is interesting to note that __rt_mutex_adjust_prio can either increase
+or decrease the priority of the task. In the case that a higher priority
+process has just blocked on a mutex owned by the task, __rt_mutex_adjust_prio
+would increase/boost the task's priority. But if a higher priority task
+were for some reason to leave the mutex (timeout or signal), this same function
+would decrease/unboost the priority of the task. That is because the pi_list
+always contains the highest priority task that is waiting on a mutex owned
+by the task, so we only need to compare the priority of that top pi waiter
+to the normal priority of the given task.
+
+
+High level overview of the PI chain walk
+----------------------------------------
+
+The PI chain walk is implemented by the function rt_mutex_adjust_prio_chain.
+
+The implementation has gone through several iterations, and has ended up
+with what we believe is the best. It walks the PI chain by only grabbing
+at most two locks at a time, and is very efficient.
+
+The rt_mutex_adjust_prio_chain can be used either to boost or lower process
+priorities.
+
+rt_mutex_adjust_prio_chain is called with a task to be checked for PI
+(de)boosting (the owner of a mutex that a process is blocking on), a flag to
+check for deadlocking, the mutex that the task owns, and a pointer to a waiter
+that is the process's waiter struct that is blocked on the mutex (although this
+parameter may be NULL for deboosting).
+
+For this explanation, I will not mention deadlock detection. This explanation
+will try to stay at a high level.
+
+When this function is called, there are no locks held. That also means
+that the state of the owner and lock can change when entered into this function.
+
+Before this function is called, the task has already had rt_mutex_adjust_prio
+performed on it. This means that the task is set to the priority that it
+should be at, but the plist nodes of the task's waiter have not been updated
+with the new priorities, and that this task may not be in the proper locations
+in the pi_lists and wait_lists that the task is blocked on. This function
+solves all that.
+
+A loop is entered, where task is the owner to be checked for PI changes that
+was passed by parameter (for the first iteration). The pi_lock of this task is
+taken to prevent any more changes to the pi_list of the task. This also
+prevents new tasks from completing the blocking on a mutex that is owned by this
+task.
+
+If the task is not blocked on a mutex then the loop is exited. We are at
+the top of the PI chain.
+
+A check is now done to see if the original waiter (the process that is blocked
+on the current mutex) is the top pi waiter of the task. That is, is this
+waiter on the top of the task's pi_list. If it is not, it either means that
+there is another process higher in priority that is blocked on one of the
+mutexes that the task owns, or that the waiter has just woken up via a signal
+or timeout and has left the PI chain. In either case, the loop is exited, since
+we don't need to do any more changes to the priority of the current task, or any
+task that owns a mutex that this current task is waiting on. A priority chain
+walk is only needed when a new top pi waiter is made to a task.
+
+The next check sees if the task's waiter plist node has the priority equal to
+the priority the task is set at. If they are equal, then we are done with
+the loop. Remember that the function started with the priority of the
+task adjusted, but the plist nodes that hold the task in other processes
+pi_lists have not been adjusted.
+
+Next, we look at the mutex that the task is blocked on. The mutex's wait_lock
+is taken. This is done by a spin_trylock, because the locking order of the
+pi_lock and wait_lock goes in the opposite direction. If we fail to grab the
+lock, the pi_lock is released, and we restart the loop.
+
+Now that we have both the pi_lock of the task as well as the wait_lock of
+the mutex the task is blocked on, we update the task's waiter's plist node
+that is located on the mutex's wait_list.
+
+Now we release the pi_lock of the task.
+
+Next the owner of the mutex has its pi_lock taken, so we can update the
+task's entry in the owner's pi_list. If the task is the highest priority
+process on the mutex's wait_list, then we remove the previous top waiter
+from the owner's pi_list, and replace it with the task.
+
+Note: It is possible that the task was the current top waiter on the mutex,
+ in which case the task is not yet on the pi_list of the waiter. This
+ is OK, since plist_del does nothing if the plist node is not on any
+ list.
+
+If the task was not the top waiter of the mutex, but it was before we
+did the priority updates, that means we are deboosting/lowering the
+task. In this case, the task is removed from the pi_list of the owner,
+and the new top waiter is added.
+
+Lastly, we unlock both the pi_lock of the task, as well as the mutex's
+wait_lock, and continue the loop again. On the next iteration of the
+loop, the previous owner of the mutex will be the task that will be
+processed.
+
+Note: One might think that the owner of this mutex might have changed
+ since we just grab the mutex's wait_lock. And one could be right.
+ The important thing to remember is that the owner could not have
+ become the task that is being processed in the PI chain, since
+ we have taken that task's pi_lock at the beginning of the loop.
+ So as long as there is an owner of this mutex that is not the same
+ process as the tasked being worked on, we are OK.
+
+ Looking closely at the code, one might be confused. The check for the
+ end of the PI chain is when the task isn't blocked on anything or the
+ task's waiter structure "task" element is NULL. This check is
+ protected only by the task's pi_lock. But the code to unlock the mutex
+ sets the task's waiter structure "task" element to NULL with only
+ the protection of the mutex's wait_lock, which was not taken yet.
+ Isn't this a race condition if the task becomes the new owner?
+
+ The answer is No! The trick is the spin_trylock of the mutex's
+ wait_lock. If we fail that lock, we release the pi_lock of the
+ task and continue the loop, doing the end of PI chain check again.
+
+ In the code to release the lock, the wait_lock of the mutex is held
+ the entire time, and it is not let go when we grab the pi_lock of the
+ new owner of the mutex. So if the switch of a new owner were to happen
+ after the check for end of the PI chain and the grabbing of the
+ wait_lock, the unlocking code would spin on the new owner's pi_lock
+ but never give up the wait_lock. So the PI chain loop is guaranteed to
+ fail the spin_trylock on the wait_lock, release the pi_lock, and
+ try again.
+
+ If you don't quite understand the above, that's OK. You don't have to,
+ unless you really want to make a proof out of it ;)
+
+
+Pending Owners and Lock stealing
+--------------------------------
+
+One of the flags in the owner field of the mutex structure is "Pending Owner".
+What this means is that an owner was chosen by the process releasing the
+mutex, but that owner has yet to wake up and actually take the mutex.
+
+Why is this important? Why can't we just give the mutex to another process
+and be done with it?
+
+The PI code is to help with real-time processes, and to let the highest
+priority process run as long as possible with little latencies and delays.
+If a high priority process owns a mutex that a lower priority process is
+blocked on, when the mutex is released it would be given to the lower priority
+process. What if the higher priority process wants to take that mutex again.
+The high priority process would fail to take that mutex that it just gave up
+and it would need to boost the lower priority process to run with full
+latency of that critical section (since the low priority process just entered
+it).
+
+There's no reason a high priority process that gives up a mutex should be
+penalized if it tries to take that mutex again. If the new owner of the
+mutex has not woken up yet, there's no reason that the higher priority process
+could not take that mutex away.
+
+To solve this, we introduced Pending Ownership and Lock Stealing. When a
+new process is given a mutex that it was blocked on, it is only given
+pending ownership. This means that it's the new owner, unless a higher
+priority process comes in and tries to grab that mutex. If a higher priority
+process does come along and wants that mutex, we let the higher priority
+process "steal" the mutex from the pending owner (only if it is still pending)
+and continue with the mutex.
+
+
+Taking of a mutex (The walk through)
+------------------------------------
+
+OK, now let's take a look at the detailed walk through of what happens when
+taking a mutex.
+
+The first thing that is tried is the fast taking of the mutex. This is
+done when we have CMPXCHG enabled (otherwise the fast taking automatically
+fails). Only when the owner field of the mutex is NULL can the lock be
+taken with the CMPXCHG and nothing else needs to be done.
+
+If there is contention on the lock, whether it is owned or pending owner
+we go about the slow path (rt_mutex_slowlock).
+
+The slow path function is where the task's waiter structure is created on
+the stack. This is because the waiter structure is only needed for the
+scope of this function. The waiter structure holds the nodes to store
+the task on the wait_list of the mutex, and if need be, the pi_list of
+the owner.
+
+The wait_lock of the mutex is taken since the slow path of unlocking the
+mutex also takes this lock.
+
+We then call try_to_take_rt_mutex. This is where the architecture that
+does not implement CMPXCHG would always grab the lock (if there's no
+contention).
+
+try_to_take_rt_mutex is used every time the task tries to grab a mutex in the
+slow path. The first thing that is done here is an atomic setting of
+the "Has Waiters" flag of the mutex's owner field. Yes, this could really
+be false, because if the mutex has no owner, there are no waiters and
+the current task also won't have any waiters. But we don't have the lock
+yet, so we assume we are going to be a waiter. The reason for this is to
+play nice for those architectures that do have CMPXCHG. By setting this flag
+now, the owner of the mutex can't release the mutex without going into the
+slow unlock path, and it would then need to grab the wait_lock, which this
+code currently holds. So setting the "Has Waiters" flag forces the owner
+to synchronize with this code.
+
+Now that we know that we can't have any races with the owner releasing the
+mutex, we check to see if we can take the ownership. This is done if the
+mutex doesn't have a owner, or if we can steal the mutex from a pending
+owner. Let's look at the situations we have here.
+
+ 1) Has owner that is pending
+ ----------------------------
+
+ The mutex has a owner, but it hasn't woken up and the mutex flag
+ "Pending Owner" is set. The first check is to see if the owner isn't the
+ current task. This is because this function is also used for the pending
+ owner to grab the mutex. When a pending owner wakes up, it checks to see
+ if it can take the mutex, and this is done if the owner is already set to
+ itself. If so, we succeed and leave the function, clearing the "Pending
+ Owner" bit.
+
+ If the pending owner is not current, we check to see if the current priority is
+ higher than the pending owner. If not, we fail the function and return.
+
+ There's also something special about a pending owner. That is a pending owner
+ is never blocked on a mutex. So there is no PI chain to worry about. It also
+ means that if the mutex doesn't have any waiters, there's no accounting needed
+ to update the pending owner's pi_list, since we only worry about processes
+ blocked on the current mutex.
+
+ If there are waiters on this mutex, and we just stole the ownership, we need
+ to take the top waiter, remove it from the pi_list of the pending owner, and
+ add it to the current pi_list. Note that at this moment, the pending owner
+ is no longer on the list of waiters. This is fine, since the pending owner
+ would add itself back when it realizes that it had the ownership stolen
+ from itself. When the pending owner tries to grab the mutex, it will fail
+ in try_to_take_rt_mutex if the owner field points to another process.
+
+ 2) No owner
+ -----------
+
+ If there is no owner (or we successfully stole the lock), we set the owner
+ of the mutex to current, and set the flag of "Has Waiters" if the current
+ mutex actually has waiters, or we clear the flag if it doesn't. See, it was
+ OK that we set that flag early, since now it is cleared.
+
+ 3) Failed to grab ownership
+ ---------------------------
+
+ The most interesting case is when we fail to take ownership. This means that
+ there exists an owner, or there's a pending owner with equal or higher
+ priority than the current task.
+
+We'll continue on the failed case.
+
+If the mutex has a timeout, we set up a timer to go off to break us out
+of this mutex if we failed to get it after a specified amount of time.
+
+Now we enter a loop that will continue to try to take ownership of the mutex, or
+fail from a timeout or signal.
+
+Once again we try to take the mutex. This will usually fail the first time
+in the loop, since it had just failed to get the mutex. But the second time
+in the loop, this would likely succeed, since the task would likely be
+the pending owner.
+
+If the mutex is TASK_INTERRUPTIBLE a check for signals and timeout is done
+here.
+
+The waiter structure has a "task" field that points to the task that is blocked
+on the mutex. This field can be NULL the first time it goes through the loop
+or if the task is a pending owner and had its mutex stolen. If the "task"
+field is NULL then we need to set up the accounting for it.
+
+Task blocks on mutex
+--------------------
+
+The accounting of a mutex and process is done with the waiter structure of
+the process. The "task" field is set to the process, and the "lock" field
+to the mutex. The plist nodes are initialized to the processes current
+priority.
+
+Since the wait_lock was taken at the entry of the slow lock, we can safely
+add the waiter to the wait_list. If the current process is the highest
+priority process currently waiting on this mutex, then we remove the
+previous top waiter process (if it exists) from the pi_list of the owner,
+and add the current process to that list. Since the pi_list of the owner
+has changed, we call rt_mutex_adjust_prio on the owner to see if the owner
+should adjust its priority accordingly.
+
+If the owner is also blocked on a lock, and had its pi_list changed
+(or deadlock checking is on), we unlock the wait_lock of the mutex and go ahead
+and run rt_mutex_adjust_prio_chain on the owner, as described earlier.
+
+Now all locks are released, and if the current process is still blocked on a
+mutex (waiter "task" field is not NULL), then we go to sleep (call schedule).
+
+Waking up in the loop
+---------------------
+
+The schedule can then wake up for a few reasons.
+ 1) we were given pending ownership of the mutex.
+ 2) we received a signal and was TASK_INTERRUPTIBLE
+ 3) we had a timeout and was TASK_INTERRUPTIBLE
+
+In any of these cases, we continue the loop and once again try to grab the
+ownership of the mutex. If we succeed, we exit the loop, otherwise we continue
+and on signal and timeout, will exit the loop, or if we had the mutex stolen
+we just simply add ourselves back on the lists and go back to sleep.
+
+Note: For various reasons, because of timeout and signals, the steal mutex
+ algorithm needs to be careful. This is because the current process is
+ still on the wait_list. And because of dynamic changing of priorities,
+ especially on SCHED_OTHER tasks, the current process can be the
+ highest priority task on the wait_list.
+
+Failed to get mutex on Timeout or Signal
+----------------------------------------
+
+If a timeout or signal occurred, the waiter's "task" field would not be
+NULL and the task needs to be taken off the wait_list of the mutex and perhaps
+pi_list of the owner. If this process was a high priority process, then
+the rt_mutex_adjust_prio_chain needs to be executed again on the owner,
+but this time it will be lowering the priorities.
+
+
+Unlocking the Mutex
+-------------------
+
+The unlocking of a mutex also has a fast path for those architectures with
+CMPXCHG. Since the taking of a mutex on contention always sets the
+"Has Waiters" flag of the mutex's owner, we use this to know if we need to
+take the slow path when unlocking the mutex. If the mutex doesn't have any
+waiters, the owner field of the mutex would equal the current process and
+the mutex can be unlocked by just replacing the owner field with NULL.
+
+If the owner field has the "Has Waiters" bit set (or CMPXCHG is not available),
+the slow unlock path is taken.
+
+The first thing done in the slow unlock path is to take the wait_lock of the
+mutex. This synchronizes the locking and unlocking of the mutex.
+
+A check is made to see if the mutex has waiters or not. On architectures that
+do not have CMPXCHG, this is the location that the owner of the mutex will
+determine if a waiter needs to be awoken or not. On architectures that
+do have CMPXCHG, that check is done in the fast path, but it is still needed
+in the slow path too. If a waiter of a mutex woke up because of a signal
+or timeout between the time the owner failed the fast path CMPXCHG check and
+the grabbing of the wait_lock, the mutex may not have any waiters, thus the
+owner still needs to make this check. If there are no waiters then the mutex
+owner field is set to NULL, the wait_lock is released and nothing more is
+needed.
+
+If there are waiters, then we need to wake one up and give that waiter
+pending ownership.
+
+On the wake up code, the pi_lock of the current owner is taken. The top
+waiter of the lock is found and removed from the wait_list of the mutex
+as well as the pi_list of the current owner. The task field of the new
+pending owner's waiter structure is set to NULL, and the owner field of the
+mutex is set to the new owner with the "Pending Owner" bit set, as well
+as the "Has Waiters" bit if there still are other processes blocked on the
+mutex.
+
+The pi_lock of the previous owner is released, and the new pending owner's
+pi_lock is taken. Remember that this is the trick to prevent the race
+condition in rt_mutex_adjust_prio_chain from adding itself as a waiter
+on the mutex.
+
+We now clear the "pi_blocked_on" field of the new pending owner, and if
+the mutex still has waiters pending, we add the new top waiter to the pi_list
+of the pending owner.
+
+Finally we unlock the pi_lock of the pending owner and wake it up.
+
+
+Contact
+-------
+
+For updates on this document, please email Steven Rostedt <[email protected]>
+
+
+Credits
+-------
+
+Author: Steven Rostedt <[email protected]>
+
+Reviewers: Ingo Molnar, Thomas Gleixner, Thomas Duetsch, and Randy Dunlap
+
+Updates
+-------
+
+This document was originally written for 2.6.17-rc3-mm1
diff --git a/Documentation/locking/rt-mutex.txt b/Documentation/locking/rt-mutex.txt
new file mode 100644
index 0000000..243393d
--- /dev/null
+++ b/Documentation/locking/rt-mutex.txt
@@ -0,0 +1,79 @@
+RT-mutex subsystem with PI support
+----------------------------------
+
+RT-mutexes with priority inheritance are used to support PI-futexes,
+which enable pthread_mutex_t priority inheritance attributes
+(PTHREAD_PRIO_INHERIT). [See Documentation/pi-futex.txt for more details
+about PI-futexes.]
+
+This technology was developed in the -rt tree and streamlined for
+pthread_mutex support.
+
+Basic principles:
+-----------------
+
+RT-mutexes extend the semantics of simple mutexes by the priority
+inheritance protocol.
+
+A low priority owner of a rt-mutex inherits the priority of a higher
+priority waiter until the rt-mutex is released. If the temporarily
+boosted owner blocks on a rt-mutex itself it propagates the priority
+boosting to the owner of the other rt_mutex it gets blocked on. The
+priority boosting is immediately removed once the rt_mutex has been
+unlocked.
+
+This approach allows us to shorten the block of high-prio tasks on
+mutexes which protect shared resources. Priority inheritance is not a
+magic bullet for poorly designed applications, but it allows
+well-designed applications to use userspace locks in critical parts of
+an high priority thread, without losing determinism.
+
+The enqueueing of the waiters into the rtmutex waiter list is done in
+priority order. For same priorities FIFO order is chosen. For each
+rtmutex, only the top priority waiter is enqueued into the owner's
+priority waiters list. This list too queues in priority order. Whenever
+the top priority waiter of a task changes (for example it timed out or
+got a signal), the priority of the owner task is readjusted. [The
+priority enqueueing is handled by "plists", see include/linux/plist.h
+for more details.]
+
+RT-mutexes are optimized for fastpath operations and have no internal
+locking overhead when locking an uncontended mutex or unlocking a mutex
+without waiters. The optimized fastpath operations require cmpxchg
+support. [If that is not available then the rt-mutex internal spinlock
+is used]
+
+The state of the rt-mutex is tracked via the owner field of the rt-mutex
+structure:
+
+rt_mutex->owner holds the task_struct pointer of the owner. Bit 0 and 1
+are used to keep track of the "owner is pending" and "rtmutex has
+waiters" state.
+
+ owner bit1 bit0
+ NULL 0 0 mutex is free (fast acquire possible)
+ NULL 0 1 invalid state
+ NULL 1 0 Transitional state*
+ NULL 1 1 invalid state
+ taskpointer 0 0 mutex is held (fast release possible)
+ taskpointer 0 1 task is pending owner
+ taskpointer 1 0 mutex is held and has waiters
+ taskpointer 1 1 task is pending owner and mutex has waiters
+
+Pending-ownership handling is a performance optimization:
+pending-ownership is assigned to the first (highest priority) waiter of
+the mutex, when the mutex is released. The thread is woken up and once
+it starts executing it can acquire the mutex. Until the mutex is taken
+by it (bit 0 is cleared) a competing higher priority thread can "steal"
+the mutex which puts the woken up thread back on the waiters list.
+
+The pending-ownership optimization is especially important for the
+uninterrupted workflow of high-prio tasks which repeatedly
+takes/releases locks that have lower-prio waiters. Without this
+optimization the higher-prio thread would ping-pong to the lower-prio
+task [because at unlock time we always assign a new owner].
+
+(*) The "mutex has waiters" bit gets set to take the lock. If the lock
+doesn't already have an owner, this bit is quickly cleared if there are
+no waiters. So this is a transitional state to synchronize with looking
+at the owner field of the mutex and the mutex owner releasing the lock.
diff --git a/Documentation/locking/spinlocks.txt b/Documentation/locking/spinlocks.txt
new file mode 100644
index 0000000..97eaf57
--- /dev/null
+++ b/Documentation/locking/spinlocks.txt
@@ -0,0 +1,167 @@
+Lesson 1: Spin locks
+
+The most basic primitive for locking is spinlock.
+
+static DEFINE_SPINLOCK(xxx_lock);
+
+ unsigned long flags;
+
+ spin_lock_irqsave(&xxx_lock, flags);
+ ... critical section here ..
+ spin_unlock_irqrestore(&xxx_lock, flags);
+
+The above is always safe. It will disable interrupts _locally_, but the
+spinlock itself will guarantee the global lock, so it will guarantee that
+there is only one thread-of-control within the region(s) protected by that
+lock. This works well even under UP also, so the code does _not_ need to
+worry about UP vs SMP issues: the spinlocks work correctly under both.
+
+ NOTE! Implications of spin_locks for memory are further described in:
+
+ Documentation/memory-barriers.txt
+ (5) LOCK operations.
+ (6) UNLOCK operations.
+
+The above is usually pretty simple (you usually need and want only one
+spinlock for most things - using more than one spinlock can make things a
+lot more complex and even slower and is usually worth it only for
+sequences that you _know_ need to be split up: avoid it at all cost if you
+aren't sure).
+
+This is really the only really hard part about spinlocks: once you start
+using spinlocks they tend to expand to areas you might not have noticed
+before, because you have to make sure the spinlocks correctly protect the
+shared data structures _everywhere_ they are used. The spinlocks are most
+easily added to places that are completely independent of other code (for
+example, internal driver data structures that nobody else ever touches).
+
+ NOTE! The spin-lock is safe only when you _also_ use the lock itself
+ to do locking across CPU's, which implies that EVERYTHING that
+ touches a shared variable has to agree about the spinlock they want
+ to use.
+
+----
+
+Lesson 2: reader-writer spinlocks.
+
+If your data accesses have a very natural pattern where you usually tend
+to mostly read from the shared variables, the reader-writer locks
+(rw_lock) versions of the spinlocks are sometimes useful. They allow multiple
+readers to be in the same critical region at once, but if somebody wants
+to change the variables it has to get an exclusive write lock.
+
+ NOTE! reader-writer locks require more atomic memory operations than
+ simple spinlocks. Unless the reader critical section is long, you
+ are better off just using spinlocks.
+
+The routines look the same as above:
+
+ rwlock_t xxx_lock = __RW_LOCK_UNLOCKED(xxx_lock);
+
+ unsigned long flags;
+
+ read_lock_irqsave(&xxx_lock, flags);
+ .. critical section that only reads the info ...
+ read_unlock_irqrestore(&xxx_lock, flags);
+
+ write_lock_irqsave(&xxx_lock, flags);
+ .. read and write exclusive access to the info ...
+ write_unlock_irqrestore(&xxx_lock, flags);
+
+The above kind of lock may be useful for complex data structures like
+linked lists, especially searching for entries without changing the list
+itself. The read lock allows many concurrent readers. Anything that
+_changes_ the list will have to get the write lock.
+
+ NOTE! RCU is better for list traversal, but requires careful
+ attention to design detail (see Documentation/RCU/listRCU.txt).
+
+Also, you cannot "upgrade" a read-lock to a write-lock, so if you at _any_
+time need to do any changes (even if you don't do it every time), you have
+to get the write-lock at the very beginning.
+
+ NOTE! We are working hard to remove reader-writer spinlocks in most
+ cases, so please don't add a new one without consensus. (Instead, see
+ Documentation/RCU/rcu.txt for complete information.)
+
+----
+
+Lesson 3: spinlocks revisited.
+
+The single spin-lock primitives above are by no means the only ones. They
+are the most safe ones, and the ones that work under all circumstances,
+but partly _because_ they are safe they are also fairly slow. They are slower
+than they'd need to be, because they do have to disable interrupts
+(which is just a single instruction on a x86, but it's an expensive one -
+and on other architectures it can be worse).
+
+If you have a case where you have to protect a data structure across
+several CPU's and you want to use spinlocks you can potentially use
+cheaper versions of the spinlocks. IFF you know that the spinlocks are
+never used in interrupt handlers, you can use the non-irq versions:
+
+ spin_lock(&lock);
+ ...
+ spin_unlock(&lock);
+
+(and the equivalent read-write versions too, of course). The spinlock will
+guarantee the same kind of exclusive access, and it will be much faster.
+This is useful if you know that the data in question is only ever
+manipulated from a "process context", ie no interrupts involved.
+
+The reasons you mustn't use these versions if you have interrupts that
+play with the spinlock is that you can get deadlocks:
+
+ spin_lock(&lock);
+ ...
+ <- interrupt comes in:
+ spin_lock(&lock);
+
+where an interrupt tries to lock an already locked variable. This is ok if
+the other interrupt happens on another CPU, but it is _not_ ok if the
+interrupt happens on the same CPU that already holds the lock, because the
+lock will obviously never be released (because the interrupt is waiting
+for the lock, and the lock-holder is interrupted by the interrupt and will
+not continue until the interrupt has been processed).
+
+(This is also the reason why the irq-versions of the spinlocks only need
+to disable the _local_ interrupts - it's ok to use spinlocks in interrupts
+on other CPU's, because an interrupt on another CPU doesn't interrupt the
+CPU that holds the lock, so the lock-holder can continue and eventually
+releases the lock).
+
+Note that you can be clever with read-write locks and interrupts. For
+example, if you know that the interrupt only ever gets a read-lock, then
+you can use a non-irq version of read locks everywhere - because they
+don't block on each other (and thus there is no dead-lock wrt interrupts.
+But when you do the write-lock, you have to use the irq-safe version.
+
+For an example of being clever with rw-locks, see the "waitqueue_lock"
+handling in kernel/sched/core.c - nothing ever _changes_ a wait-queue from
+within an interrupt, they only read the queue in order to know whom to
+wake up. So read-locks are safe (which is good: they are very common
+indeed), while write-locks need to protect themselves against interrupts.
+
+ Linus
+
+----
+
+Reference information:
+
+For dynamic initialization, use spin_lock_init() or rwlock_init() as
+appropriate:
+
+ spinlock_t xxx_lock;
+ rwlock_t xxx_rw_lock;
+
+ static int __init xxx_init(void)
+ {
+ spin_lock_init(&xxx_lock);
+ rwlock_init(&xxx_rw_lock);
+ ...
+ }
+
+ module_init(xxx_init);
+
+For static initialization, use DEFINE_SPINLOCK() / DEFINE_RWLOCK() or
+__SPIN_LOCK_UNLOCKED() / __RW_LOCK_UNLOCKED() as appropriate.
diff --git a/Documentation/locking/ww-mutex-design.txt b/Documentation/locking/ww-mutex-design.txt
new file mode 100644
index 0000000..8a112dc
--- /dev/null
+++ b/Documentation/locking/ww-mutex-design.txt
@@ -0,0 +1,344 @@
+Wait/Wound Deadlock-Proof Mutex Design
+======================================
+
+Please read mutex-design.txt first, as it applies to wait/wound mutexes too.
+
+Motivation for WW-Mutexes
+-------------------------
+
+GPU's do operations that commonly involve many buffers. Those buffers
+can be shared across contexts/processes, exist in different memory
+domains (for example VRAM vs system memory), and so on. And with
+PRIME / dmabuf, they can even be shared across devices. So there are
+a handful of situations where the driver needs to wait for buffers to
+become ready. If you think about this in terms of waiting on a buffer
+mutex for it to become available, this presents a problem because
+there is no way to guarantee that buffers appear in a execbuf/batch in
+the same order in all contexts. That is directly under control of
+userspace, and a result of the sequence of GL calls that an application
+makes. Which results in the potential for deadlock. The problem gets
+more complex when you consider that the kernel may need to migrate the
+buffer(s) into VRAM before the GPU operates on the buffer(s), which
+may in turn require evicting some other buffers (and you don't want to
+evict other buffers which are already queued up to the GPU), but for a
+simplified understanding of the problem you can ignore this.
+
+The algorithm that the TTM graphics subsystem came up with for dealing with
+this problem is quite simple. For each group of buffers (execbuf) that need
+to be locked, the caller would be assigned a unique reservation id/ticket,
+from a global counter. In case of deadlock while locking all the buffers
+associated with a execbuf, the one with the lowest reservation ticket (i.e.
+the oldest task) wins, and the one with the higher reservation id (i.e. the
+younger task) unlocks all of the buffers that it has already locked, and then
+tries again.
+
+In the RDBMS literature this deadlock handling approach is called wait/wound:
+The older tasks waits until it can acquire the contended lock. The younger tasks
+needs to back off and drop all the locks it is currently holding, i.e. the
+younger task is wounded.
+
+Concepts
+--------
+
+Compared to normal mutexes two additional concepts/objects show up in the lock
+interface for w/w mutexes:
+
+Acquire context: To ensure eventual forward progress it is important the a task
+trying to acquire locks doesn't grab a new reservation id, but keeps the one it
+acquired when starting the lock acquisition. This ticket is stored in the
+acquire context. Furthermore the acquire context keeps track of debugging state
+to catch w/w mutex interface abuse.
+
+W/w class: In contrast to normal mutexes the lock class needs to be explicit for
+w/w mutexes, since it is required to initialize the acquire context.
+
+Furthermore there are three different class of w/w lock acquire functions:
+
+* Normal lock acquisition with a context, using ww_mutex_lock.
+
+* Slowpath lock acquisition on the contending lock, used by the wounded task
+ after having dropped all already acquired locks. These functions have the
+ _slow postfix.
+
+ From a simple semantics point-of-view the _slow functions are not strictly
+ required, since simply calling the normal ww_mutex_lock functions on the
+ contending lock (after having dropped all other already acquired locks) will
+ work correctly. After all if no other ww mutex has been acquired yet there's
+ no deadlock potential and hence the ww_mutex_lock call will block and not
+ prematurely return -EDEADLK. The advantage of the _slow functions is in
+ interface safety:
+ - ww_mutex_lock has a __must_check int return type, whereas ww_mutex_lock_slow
+ has a void return type. Note that since ww mutex code needs loops/retries
+ anyway the __must_check doesn't result in spurious warnings, even though the
+ very first lock operation can never fail.
+ - When full debugging is enabled ww_mutex_lock_slow checks that all acquired
+ ww mutex have been released (preventing deadlocks) and makes sure that we
+ block on the contending lock (preventing spinning through the -EDEADLK
+ slowpath until the contended lock can be acquired).
+
+* Functions to only acquire a single w/w mutex, which results in the exact same
+ semantics as a normal mutex. This is done by calling ww_mutex_lock with a NULL
+ context.
+
+ Again this is not strictly required. But often you only want to acquire a
+ single lock in which case it's pointless to set up an acquire context (and so
+ better to avoid grabbing a deadlock avoidance ticket).
+
+Of course, all the usual variants for handling wake-ups due to signals are also
+provided.
+
+Usage
+-----
+
+Three different ways to acquire locks within the same w/w class. Common
+definitions for methods #1 and #2:
+
+static DEFINE_WW_CLASS(ww_class);
+
+struct obj {
+ struct ww_mutex lock;
+ /* obj data */
+};
+
+struct obj_entry {
+ struct list_head head;
+ struct obj *obj;
+};
+
+Method 1, using a list in execbuf->buffers that's not allowed to be reordered.
+This is useful if a list of required objects is already tracked somewhere.
+Furthermore the lock helper can use propagate the -EALREADY return code back to
+the caller as a signal that an object is twice on the list. This is useful if
+the list is constructed from userspace input and the ABI requires userspace to
+not have duplicate entries (e.g. for a gpu commandbuffer submission ioctl).
+
+int lock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
+{
+ struct obj *res_obj = NULL;
+ struct obj_entry *contended_entry = NULL;
+ struct obj_entry *entry;
+
+ ww_acquire_init(ctx, &ww_class);
+
+retry:
+ list_for_each_entry (entry, list, head) {
+ if (entry->obj == res_obj) {
+ res_obj = NULL;
+ continue;
+ }
+ ret = ww_mutex_lock(&entry->obj->lock, ctx);
+ if (ret < 0) {
+ contended_entry = entry;
+ goto err;
+ }
+ }
+
+ ww_acquire_done(ctx);
+ return 0;
+
+err:
+ list_for_each_entry_continue_reverse (entry, list, head)
+ ww_mutex_unlock(&entry->obj->lock);
+
+ if (res_obj)
+ ww_mutex_unlock(&res_obj->lock);
+
+ if (ret == -EDEADLK) {
+ /* we lost out in a seqno race, lock and retry.. */
+ ww_mutex_lock_slow(&contended_entry->obj->lock, ctx);
+ res_obj = contended_entry->obj;
+ goto retry;
+ }
+ ww_acquire_fini(ctx);
+
+ return ret;
+}
+
+Method 2, using a list in execbuf->buffers that can be reordered. Same semantics
+of duplicate entry detection using -EALREADY as method 1 above. But the
+list-reordering allows for a bit more idiomatic code.
+
+int lock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
+{
+ struct obj_entry *entry, *entry2;
+
+ ww_acquire_init(ctx, &ww_class);
+
+ list_for_each_entry (entry, list, head) {
+ ret = ww_mutex_lock(&entry->obj->lock, ctx);
+ if (ret < 0) {
+ entry2 = entry;
+
+ list_for_each_entry_continue_reverse (entry2, list, head)
+ ww_mutex_unlock(&entry2->obj->lock);
+
+ if (ret != -EDEADLK) {
+ ww_acquire_fini(ctx);
+ return ret;
+ }
+
+ /* we lost out in a seqno race, lock and retry.. */
+ ww_mutex_lock_slow(&entry->obj->lock, ctx);
+
+ /*
+ * Move buf to head of the list, this will point
+ * buf->next to the first unlocked entry,
+ * restarting the for loop.
+ */
+ list_del(&entry->head);
+ list_add(&entry->head, list);
+ }
+ }
+
+ ww_acquire_done(ctx);
+ return 0;
+}
+
+Unlocking works the same way for both methods #1 and #2:
+
+void unlock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
+{
+ struct obj_entry *entry;
+
+ list_for_each_entry (entry, list, head)
+ ww_mutex_unlock(&entry->obj->lock);
+
+ ww_acquire_fini(ctx);
+}
+
+Method 3 is useful if the list of objects is constructed ad-hoc and not upfront,
+e.g. when adjusting edges in a graph where each node has its own ww_mutex lock,
+and edges can only be changed when holding the locks of all involved nodes. w/w
+mutexes are a natural fit for such a case for two reasons:
+- They can handle lock-acquisition in any order which allows us to start walking
+ a graph from a starting point and then iteratively discovering new edges and
+ locking down the nodes those edges connect to.
+- Due to the -EALREADY return code signalling that a given objects is already
+ held there's no need for additional book-keeping to break cycles in the graph
+ or keep track off which looks are already held (when using more than one node
+ as a starting point).
+
+Note that this approach differs in two important ways from the above methods:
+- Since the list of objects is dynamically constructed (and might very well be
+ different when retrying due to hitting the -EDEADLK wound condition) there's
+ no need to keep any object on a persistent list when it's not locked. We can
+ therefore move the list_head into the object itself.
+- On the other hand the dynamic object list construction also means that the -EALREADY return
+ code can't be propagated.
+
+Note also that methods #1 and #2 and method #3 can be combined, e.g. to first lock a
+list of starting nodes (passed in from userspace) using one of the above
+methods. And then lock any additional objects affected by the operations using
+method #3 below. The backoff/retry procedure will be a bit more involved, since
+when the dynamic locking step hits -EDEADLK we also need to unlock all the
+objects acquired with the fixed list. But the w/w mutex debug checks will catch
+any interface misuse for these cases.
+
+Also, method 3 can't fail the lock acquisition step since it doesn't return
+-EALREADY. Of course this would be different when using the _interruptible
+variants, but that's outside of the scope of these examples here.
+
+struct obj {
+ struct ww_mutex ww_mutex;
+ struct list_head locked_list;
+};
+
+static DEFINE_WW_CLASS(ww_class);
+
+void __unlock_objs(struct list_head *list)
+{
+ struct obj *entry, *temp;
+
+ list_for_each_entry_safe (entry, temp, list, locked_list) {
+ /* need to do that before unlocking, since only the current lock holder is
+ allowed to use object */
+ list_del(&entry->locked_list);
+ ww_mutex_unlock(entry->ww_mutex)
+ }
+}
+
+void lock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
+{
+ struct obj *obj;
+
+ ww_acquire_init(ctx, &ww_class);
+
+retry:
+ /* re-init loop start state */
+ loop {
+ /* magic code which walks over a graph and decides which objects
+ * to lock */
+
+ ret = ww_mutex_lock(obj->ww_mutex, ctx);
+ if (ret == -EALREADY) {
+ /* we have that one already, get to the next object */
+ continue;
+ }
+ if (ret == -EDEADLK) {
+ __unlock_objs(list);
+
+ ww_mutex_lock_slow(obj, ctx);
+ list_add(&entry->locked_list, list);
+ goto retry;
+ }
+
+ /* locked a new object, add it to the list */
+ list_add_tail(&entry->locked_list, list);
+ }
+
+ ww_acquire_done(ctx);
+ return 0;
+}
+
+void unlock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
+{
+ __unlock_objs(list);
+ ww_acquire_fini(ctx);
+}
+
+Method 4: Only lock one single objects. In that case deadlock detection and
+prevention is obviously overkill, since with grabbing just one lock you can't
+produce a deadlock within just one class. To simplify this case the w/w mutex
+api can be used with a NULL context.
+
+Implementation Details
+----------------------
+
+Design:
+ ww_mutex currently encapsulates a struct mutex, this means no extra overhead for
+ normal mutex locks, which are far more common. As such there is only a small
+ increase in code size if wait/wound mutexes are not used.
+
+ In general, not much contention is expected. The locks are typically used to
+ serialize access to resources for devices. The only way to make wakeups
+ smarter would be at the cost of adding a field to struct mutex_waiter. This
+ would add overhead to all cases where normal mutexes are used, and
+ ww_mutexes are generally less performance sensitive.
+
+Lockdep:
+ Special care has been taken to warn for as many cases of api abuse
+ as possible. Some common api abuses will be caught with
+ CONFIG_DEBUG_MUTEXES, but CONFIG_PROVE_LOCKING is recommended.
+
+ Some of the errors which will be warned about:
+ - Forgetting to call ww_acquire_fini or ww_acquire_init.
+ - Attempting to lock more mutexes after ww_acquire_done.
+ - Attempting to lock the wrong mutex after -EDEADLK and
+ unlocking all mutexes.
+ - Attempting to lock the right mutex after -EDEADLK,
+ before unlocking all mutexes.
+
+ - Calling ww_mutex_lock_slow before -EDEADLK was returned.
+
+ - Unlocking mutexes with the wrong unlock function.
+ - Calling one of the ww_acquire_* twice on the same context.
+ - Using a different ww_class for the mutex than for the ww_acquire_ctx.
+ - Normal lockdep errors that can result in deadlocks.
+
+ Some of the lockdep errors that can result in deadlocks:
+ - Calling ww_acquire_init to initialize a second ww_acquire_ctx before
+ having called ww_acquire_fini on the first.
+ - 'normal' deadlocks that can occur.
+
+FIXME: Update this section once we have the TASK_DEADLOCK task state flag magic
+implemented.
diff --git a/Documentation/lockstat.txt b/Documentation/lockstat.txt
deleted file mode 100644
index 72d0106..0000000
--- a/Documentation/lockstat.txt
+++ /dev/null
@@ -1,178 +0,0 @@
-
-LOCK STATISTICS
-
-- WHAT
-
-As the name suggests, it provides statistics on locks.
-
-- WHY
-
-Because things like lock contention can severely impact performance.
-
-- HOW
-
-Lockdep already has hooks in the lock functions and maps lock instances to
-lock classes. We build on that (see Documentation/lockdep-design.txt).
-The graph below shows the relation between the lock functions and the various
-hooks therein.
-
- __acquire
- |
- lock _____
- | \
- | __contended
- | |
- | <wait>
- | _______/
- |/
- |
- __acquired
- |
- .
- <hold>
- .
- |
- __release
- |
- unlock
-
-lock, unlock - the regular lock functions
-__* - the hooks
-<> - states
-
-With these hooks we provide the following statistics:
-
- con-bounces - number of lock contention that involved x-cpu data
- contentions - number of lock acquisitions that had to wait
- wait time min - shortest (non-0) time we ever had to wait for a lock
- max - longest time we ever had to wait for a lock
- total - total time we spend waiting on this lock
- avg - average time spent waiting on this lock
- acq-bounces - number of lock acquisitions that involved x-cpu data
- acquisitions - number of times we took the lock
- hold time min - shortest (non-0) time we ever held the lock
- max - longest time we ever held the lock
- total - total time this lock was held
- avg - average time this lock was held
-
-These numbers are gathered per lock class, per read/write state (when
-applicable).
-
-It also tracks 4 contention points per class. A contention point is a call site
-that had to wait on lock acquisition.
-
- - CONFIGURATION
-
-Lock statistics are enabled via CONFIG_LOCK_STAT.
-
- - USAGE
-
-Enable collection of statistics:
-
-# echo 1 >/proc/sys/kernel/lock_stat
-
-Disable collection of statistics:
-
-# echo 0 >/proc/sys/kernel/lock_stat
-
-Look at the current lock statistics:
-
-( line numbers not part of actual output, done for clarity in the explanation
- below )
-
-# less /proc/lock_stat
-
-01 lock_stat version 0.4
-02-----------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------
-03 class name con-bounces contentions waittime-min waittime-max waittime-total waittime-avg acq-bounces acquisitions holdtime-min holdtime-max holdtime-total holdtime-avg
-04-----------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------
-05
-06 &mm->mmap_sem-W: 46 84 0.26 939.10 16371.53 194.90 47291 2922365 0.16 2220301.69 17464026916.32 5975.99
-07 &mm->mmap_sem-R: 37 100 1.31 299502.61 325629.52 3256.30 212344 34316685 0.10 7744.91 95016910.20 2.77
-08 ---------------
-09 &mm->mmap_sem 1 [<ffffffff811502a7>] khugepaged_scan_mm_slot+0x57/0x280
-19 &mm->mmap_sem 96 [<ffffffff815351c4>] __do_page_fault+0x1d4/0x510
-11 &mm->mmap_sem 34 [<ffffffff81113d77>] vm_mmap_pgoff+0x87/0xd0
-12 &mm->mmap_sem 17 [<ffffffff81127e71>] vm_munmap+0x41/0x80
-13 ---------------
-14 &mm->mmap_sem 1 [<ffffffff81046fda>] dup_mmap+0x2a/0x3f0
-15 &mm->mmap_sem 60 [<ffffffff81129e29>] SyS_mprotect+0xe9/0x250
-16 &mm->mmap_sem 41 [<ffffffff815351c4>] __do_page_fault+0x1d4/0x510
-17 &mm->mmap_sem 68 [<ffffffff81113d77>] vm_mmap_pgoff+0x87/0xd0
-18
-19.............................................................................................................................................................................................................................
-20
-21 unix_table_lock: 110 112 0.21 49.24 163.91 1.46 21094 66312 0.12 624.42 31589.81 0.48
-22 ---------------
-23 unix_table_lock 45 [<ffffffff8150ad8e>] unix_create1+0x16e/0x1b0
-24 unix_table_lock 47 [<ffffffff8150b111>] unix_release_sock+0x31/0x250
-25 unix_table_lock 15 [<ffffffff8150ca37>] unix_find_other+0x117/0x230
-26 unix_table_lock 5 [<ffffffff8150a09f>] unix_autobind+0x11f/0x1b0
-27 ---------------
-28 unix_table_lock 39 [<ffffffff8150b111>] unix_release_sock+0x31/0x250
-29 unix_table_lock 49 [<ffffffff8150ad8e>] unix_create1+0x16e/0x1b0
-30 unix_table_lock 20 [<ffffffff8150ca37>] unix_find_other+0x117/0x230
-31 unix_table_lock 4 [<ffffffff8150a09f>] unix_autobind+0x11f/0x1b0
-
-
-This excerpt shows the first two lock class statistics. Line 01 shows the
-output version - each time the format changes this will be updated. Line 02-04
-show the header with column descriptions. Lines 05-18 and 20-31 show the actual
-statistics. These statistics come in two parts; the actual stats separated by a
-short separator (line 08, 13) from the contention points.
-
-The first lock (05-18) is a read/write lock, and shows two lines above the
-short separator. The contention points don't match the column descriptors,
-they have two: contentions and [<IP>] symbol. The second set of contention
-points are the points we're contending with.
-
-The integer part of the time values is in us.
-
-Dealing with nested locks, subclasses may appear:
-
-32...........................................................................................................................................................................................................................
-33
-34 &rq->lock: 13128 13128 0.43 190.53 103881.26 7.91 97454 3453404 0.00 401.11 13224683.11 3.82
-35 ---------
-36 &rq->lock 645 [<ffffffff8103bfc4>] task_rq_lock+0x43/0x75
-37 &rq->lock 297 [<ffffffff8104ba65>] try_to_wake_up+0x127/0x25a
-38 &rq->lock 360 [<ffffffff8103c4c5>] select_task_rq_fair+0x1f0/0x74a
-39 &rq->lock 428 [<ffffffff81045f98>] scheduler_tick+0x46/0x1fb
-40 ---------
-41 &rq->lock 77 [<ffffffff8103bfc4>] task_rq_lock+0x43/0x75
-42 &rq->lock 174 [<ffffffff8104ba65>] try_to_wake_up+0x127/0x25a
-43 &rq->lock 4715 [<ffffffff8103ed4b>] double_rq_lock+0x42/0x54
-44 &rq->lock 893 [<ffffffff81340524>] schedule+0x157/0x7b8
-45
-46...........................................................................................................................................................................................................................
-47
-48 &rq->lock/1: 1526 11488 0.33 388.73 136294.31 11.86 21461 38404 0.00 37.93 109388.53 2.84
-49 -----------
-50 &rq->lock/1 11526 [<ffffffff8103ed58>] double_rq_lock+0x4f/0x54
-51 -----------
-52 &rq->lock/1 5645 [<ffffffff8103ed4b>] double_rq_lock+0x42/0x54
-53 &rq->lock/1 1224 [<ffffffff81340524>] schedule+0x157/0x7b8
-54 &rq->lock/1 4336 [<ffffffff8103ed58>] double_rq_lock+0x4f/0x54
-55 &rq->lock/1 181 [<ffffffff8104ba65>] try_to_wake_up+0x127/0x25a
-
-Line 48 shows statistics for the second subclass (/1) of &rq->lock class
-(subclass starts from 0), since in this case, as line 50 suggests,
-double_rq_lock actually acquires a nested lock of two spinlocks.
-
-View the top contending locks:
-
-# grep : /proc/lock_stat | head
- clockevents_lock: 2926159 2947636 0.15 46882.81 1784540466.34 605.41 3381345 3879161 0.00 2260.97 53178395.68 13.71
- tick_broadcast_lock: 346460 346717 0.18 2257.43 39364622.71 113.54 3642919 4242696 0.00 2263.79 49173646.60 11.59
- &mapping->i_mmap_mutex: 203896 203899 3.36 645530.05 31767507988.39 155800.21 3361776 8893984 0.17 2254.15 14110121.02 1.59
- &rq->lock: 135014 136909 0.18 606.09 842160.68 6.15 1540728 10436146 0.00 728.72 17606683.41 1.69
- &(&zone->lru_lock)->rlock: 93000 94934 0.16 59.18 188253.78 1.98 1199912 3809894 0.15 391.40 3559518.81 0.93
- tasklist_lock-W: 40667 41130 0.23 1189.42 428980.51 10.43 270278 510106 0.16 653.51 3939674.91 7.72
- tasklist_lock-R: 21298 21305 0.20 1310.05 215511.12 10.12 186204 241258 0.14 1162.33 1179779.23 4.89
- rcu_node_1: 47656 49022 0.16 635.41 193616.41 3.95 844888 1865423 0.00 764.26 1656226.96 0.89
- &(&dentry->d_lockref.lock)->rlock: 39791 40179 0.15 1302.08 88851.96 2.21 2790851 12527025 0.10 1910.75 3379714.27 0.27
- rcu_node_0: 29203 30064 0.16 786.55 1555573.00 51.74 88963 244254 0.00 398.87 428872.51 1.76
-
-Clear the statistics:
-
-# echo 0 > /proc/lock_stat
diff --git a/Documentation/mutex-design.txt b/Documentation/mutex-design.txt
deleted file mode 100644
index ee231ed..0000000
--- a/Documentation/mutex-design.txt
+++ /dev/null
@@ -1,157 +0,0 @@
-Generic Mutex Subsystem
-
-started by Ingo Molnar <[email protected]>
-updated by Davidlohr Bueso <[email protected]>
-
-What are mutexes?
------------------
-
-In the Linux kernel, mutexes refer to a particular locking primitive
-that enforces serialization on shared memory systems, and not only to
-the generic term referring to 'mutual exclusion' found in academia
-or similar theoretical text books. Mutexes are sleeping locks which
-behave similarly to binary semaphores, and were introduced in 2006[1]
-as an alternative to these. This new data structure provided a number
-of advantages, including simpler interfaces, and at that time smaller
-code (see Disadvantages).
-
-[1] http://lwn.net/Articles/164802/
-
-Implementation
---------------
-
-Mutexes are represented by 'struct mutex', defined in include/linux/mutex.h
-and implemented in kernel/locking/mutex.c. These locks use a three
-state atomic counter (->count) to represent the different possible
-transitions that can occur during the lifetime of a lock:
-
- 1: unlocked
- 0: locked, no waiters
- negative: locked, with potential waiters
-
-In its most basic form it also includes a wait-queue and a spinlock
-that serializes access to it. CONFIG_SMP systems can also include
-a pointer to the lock task owner (->owner) as well as a spinner MCS
-lock (->osq), both described below in (ii).
-
-When acquiring a mutex, there are three possible paths that can be
-taken, depending on the state of the lock:
-
-(i) fastpath: tries to atomically acquire the lock by decrementing the
- counter. If it was already taken by another task it goes to the next
- possible path. This logic is architecture specific. On x86-64, the
- locking fastpath is 2 instructions:
-
- 0000000000000e10 <mutex_lock>:
- e21: f0 ff 0b lock decl (%rbx)
- e24: 79 08 jns e2e <mutex_lock+0x1e>
-
- the unlocking fastpath is equally tight:
-
- 0000000000000bc0 <mutex_unlock>:
- bc8: f0 ff 07 lock incl (%rdi)
- bcb: 7f 0a jg bd7 <mutex_unlock+0x17>
-
-
-(ii) midpath: aka optimistic spinning, tries to spin for acquisition
- while the lock owner is running and there are no other tasks ready
- to run that have higher priority (need_resched). The rationale is
- that if the lock owner is running, it is likely to release the lock
- soon. The mutex spinners are queued up using MCS lock so that only
- one spinner can compete for the mutex.
-
- The MCS lock (proposed by Mellor-Crummey and Scott) is a simple spinlock
- with the desirable properties of being fair and with each cpu trying
- to acquire the lock spinning on a local variable. It avoids expensive
- cacheline bouncing that common test-and-set spinlock implementations
- incur. An MCS-like lock is specially tailored for optimistic spinning
- for sleeping lock implementation. An important feature of the customized
- MCS lock is that it has the extra property that spinners are able to exit
- the MCS spinlock queue when they need to reschedule. This further helps
- avoid situations where MCS spinners that need to reschedule would continue
- waiting to spin on mutex owner, only to go directly to slowpath upon
- obtaining the MCS lock.
-
-
-(iii) slowpath: last resort, if the lock is still unable to be acquired,
- the task is added to the wait-queue and sleeps until woken up by the
- unlock path. Under normal circumstances it blocks as TASK_UNINTERRUPTIBLE.
-
-While formally kernel mutexes are sleepable locks, it is path (ii) that
-makes them more practically a hybrid type. By simply not interrupting a
-task and busy-waiting for a few cycles instead of immediately sleeping,
-the performance of this lock has been seen to significantly improve a
-number of workloads. Note that this technique is also used for rw-semaphores.
-
-Semantics
----------
-
-The mutex subsystem checks and enforces the following rules:
-
- - Only one task can hold the mutex at a time.
- - Only the owner can unlock the mutex.
- - Multiple unlocks are not permitted.
- - Recursive locking/unlocking is not permitted.
- - A mutex must only be initialized via the API (see below).
- - A task may not exit with a mutex held.
- - Memory areas where held locks reside must not be freed.
- - Held mutexes must not be reinitialized.
- - Mutexes may not be used in hardware or software interrupt
- contexts such as tasklets and timers.
-
-These semantics are fully enforced when CONFIG DEBUG_MUTEXES is enabled.
-In addition, the mutex debugging code also implements a number of other
-features that make lock debugging easier and faster:
-
- - Uses symbolic names of mutexes, whenever they are printed
- in debug output.
- - Point-of-acquire tracking, symbolic lookup of function names,
- list of all locks held in the system, printout of them.
- - Owner tracking.
- - Detects self-recursing locks and prints out all relevant info.
- - Detects multi-task circular deadlocks and prints out all affected
- locks and tasks (and only those tasks).
-
-
-Interfaces
-----------
-Statically define the mutex:
- DEFINE_MUTEX(name);
-
-Dynamically initialize the mutex:
- mutex_init(mutex);
-
-Acquire the mutex, uninterruptible:
- void mutex_lock(struct mutex *lock);
- void mutex_lock_nested(struct mutex *lock, unsigned int subclass);
- int mutex_trylock(struct mutex *lock);
-
-Acquire the mutex, interruptible:
- int mutex_lock_interruptible_nested(struct mutex *lock,
- unsigned int subclass);
- int mutex_lock_interruptible(struct mutex *lock);
-
-Acquire the mutex, interruptible, if dec to 0:
- int atomic_dec_and_mutex_lock(atomic_t *cnt, struct mutex *lock);
-
-Unlock the mutex:
- void mutex_unlock(struct mutex *lock);
-
-Test if the mutex is taken:
- int mutex_is_locked(struct mutex *lock);
-
-Disadvantages
--------------
-
-Unlike its original design and purpose, 'struct mutex' is larger than
-most locks in the kernel. E.g: on x86-64 it is 40 bytes, almost twice
-as large as 'struct semaphore' (24 bytes) and 8 bytes shy of the
-'struct rw_semaphore' variant. Larger structure sizes mean more CPU
-cache and memory footprint.
-
-When to use mutexes
--------------------
-
-Unless the strict semantics of mutexes are unsuitable and/or the critical
-region prevents the lock from being shared, always prefer them to any other
-locking primitive.
diff --git a/Documentation/rt-mutex-design.txt b/Documentation/rt-mutex-design.txt
deleted file mode 100644
index 8666070..0000000
--- a/Documentation/rt-mutex-design.txt
+++ /dev/null
@@ -1,781 +0,0 @@
-#
-# Copyright (c) 2006 Steven Rostedt
-# Licensed under the GNU Free Documentation License, Version 1.2
-#
-
-RT-mutex implementation design
-------------------------------
-
-This document tries to describe the design of the rtmutex.c implementation.
-It doesn't describe the reasons why rtmutex.c exists. For that please see
-Documentation/rt-mutex.txt. Although this document does explain problems
-that happen without this code, but that is in the concept to understand
-what the code actually is doing.
-
-The goal of this document is to help others understand the priority
-inheritance (PI) algorithm that is used, as well as reasons for the
-decisions that were made to implement PI in the manner that was done.
-
-
-Unbounded Priority Inversion
-----------------------------
-
-Priority inversion is when a lower priority process executes while a higher
-priority process wants to run. This happens for several reasons, and
-most of the time it can't be helped. Anytime a high priority process wants
-to use a resource that a lower priority process has (a mutex for example),
-the high priority process must wait until the lower priority process is done
-with the resource. This is a priority inversion. What we want to prevent
-is something called unbounded priority inversion. That is when the high
-priority process is prevented from running by a lower priority process for
-an undetermined amount of time.
-
-The classic example of unbounded priority inversion is where you have three
-processes, let's call them processes A, B, and C, where A is the highest
-priority process, C is the lowest, and B is in between. A tries to grab a lock
-that C owns and must wait and lets C run to release the lock. But in the
-meantime, B executes, and since B is of a higher priority than C, it preempts C,
-but by doing so, it is in fact preempting A which is a higher priority process.
-Now there's no way of knowing how long A will be sleeping waiting for C
-to release the lock, because for all we know, B is a CPU hog and will
-never give C a chance to release the lock. This is called unbounded priority
-inversion.
-
-Here's a little ASCII art to show the problem.
-
- grab lock L1 (owned by C)
- |
-A ---+
- C preempted by B
- |
-C +----+
-
-B +-------->
- B now keeps A from running.
-
-
-Priority Inheritance (PI)
--------------------------
-
-There are several ways to solve this issue, but other ways are out of scope
-for this document. Here we only discuss PI.
-
-PI is where a process inherits the priority of another process if the other
-process blocks on a lock owned by the current process. To make this easier
-to understand, let's use the previous example, with processes A, B, and C again.
-
-This time, when A blocks on the lock owned by C, C would inherit the priority
-of A. So now if B becomes runnable, it would not preempt C, since C now has
-the high priority of A. As soon as C releases the lock, it loses its
-inherited priority, and A then can continue with the resource that C had.
-
-Terminology
------------
-
-Here I explain some terminology that is used in this document to help describe
-the design that is used to implement PI.
-
-PI chain - The PI chain is an ordered series of locks and processes that cause
- processes to inherit priorities from a previous process that is
- blocked on one of its locks. This is described in more detail
- later in this document.
-
-mutex - In this document, to differentiate from locks that implement
- PI and spin locks that are used in the PI code, from now on
- the PI locks will be called a mutex.
-
-lock - In this document from now on, I will use the term lock when
- referring to spin locks that are used to protect parts of the PI
- algorithm. These locks disable preemption for UP (when
- CONFIG_PREEMPT is enabled) and on SMP prevents multiple CPUs from
- entering critical sections simultaneously.
-
-spin lock - Same as lock above.
-
-waiter - A waiter is a struct that is stored on the stack of a blocked
- process. Since the scope of the waiter is within the code for
- a process being blocked on the mutex, it is fine to allocate
- the waiter on the process's stack (local variable). This
- structure holds a pointer to the task, as well as the mutex that
- the task is blocked on. It also has the plist node structures to
- place the task in the waiter_list of a mutex as well as the
- pi_list of a mutex owner task (described below).
-
- waiter is sometimes used in reference to the task that is waiting
- on a mutex. This is the same as waiter->task.
-
-waiters - A list of processes that are blocked on a mutex.
-
-top waiter - The highest priority process waiting on a specific mutex.
-
-top pi waiter - The highest priority process waiting on one of the mutexes
- that a specific process owns.
-
-Note: task and process are used interchangeably in this document, mostly to
- differentiate between two processes that are being described together.
-
-
-PI chain
---------
-
-The PI chain is a list of processes and mutexes that may cause priority
-inheritance to take place. Multiple chains may converge, but a chain
-would never diverge, since a process can't be blocked on more than one
-mutex at a time.
-
-Example:
-
- Process: A, B, C, D, E
- Mutexes: L1, L2, L3, L4
-
- A owns: L1
- B blocked on L1
- B owns L2
- C blocked on L2
- C owns L3
- D blocked on L3
- D owns L4
- E blocked on L4
-
-The chain would be:
-
- E->L4->D->L3->C->L2->B->L1->A
-
-To show where two chains merge, we could add another process F and
-another mutex L5 where B owns L5 and F is blocked on mutex L5.
-
-The chain for F would be:
-
- F->L5->B->L1->A
-
-Since a process may own more than one mutex, but never be blocked on more than
-one, the chains merge.
-
-Here we show both chains:
-
- E->L4->D->L3->C->L2-+
- |
- +->B->L1->A
- |
- F->L5-+
-
-For PI to work, the processes at the right end of these chains (or we may
-also call it the Top of the chain) must be equal to or higher in priority
-than the processes to the left or below in the chain.
-
-Also since a mutex may have more than one process blocked on it, we can
-have multiple chains merge at mutexes. If we add another process G that is
-blocked on mutex L2:
-
- G->L2->B->L1->A
-
-And once again, to show how this can grow I will show the merging chains
-again.
-
- E->L4->D->L3->C-+
- +->L2-+
- | |
- G-+ +->B->L1->A
- |
- F->L5-+
-
-
-Plist
------
-
-Before I go further and talk about how the PI chain is stored through lists
-on both mutexes and processes, I'll explain the plist. This is similar to
-the struct list_head functionality that is already in the kernel.
-The implementation of plist is out of scope for this document, but it is
-very important to understand what it does.
-
-There are a few differences between plist and list, the most important one
-being that plist is a priority sorted linked list. This means that the
-priorities of the plist are sorted, such that it takes O(1) to retrieve the
-highest priority item in the list. Obviously this is useful to store processes
-based on their priorities.
-
-Another difference, which is important for implementation, is that, unlike
-list, the head of the list is a different element than the nodes of a list.
-So the head of the list is declared as struct plist_head and nodes that will
-be added to the list are declared as struct plist_node.
-
-
-Mutex Waiter List
------------------
-
-Every mutex keeps track of all the waiters that are blocked on itself. The mutex
-has a plist to store these waiters by priority. This list is protected by
-a spin lock that is located in the struct of the mutex. This lock is called
-wait_lock. Since the modification of the waiter list is never done in
-interrupt context, the wait_lock can be taken without disabling interrupts.
-
-
-Task PI List
-------------
-
-To keep track of the PI chains, each process has its own PI list. This is
-a list of all top waiters of the mutexes that are owned by the process.
-Note that this list only holds the top waiters and not all waiters that are
-blocked on mutexes owned by the process.
-
-The top of the task's PI list is always the highest priority task that
-is waiting on a mutex that is owned by the task. So if the task has
-inherited a priority, it will always be the priority of the task that is
-at the top of this list.
-
-This list is stored in the task structure of a process as a plist called
-pi_list. This list is protected by a spin lock also in the task structure,
-called pi_lock. This lock may also be taken in interrupt context, so when
-locking the pi_lock, interrupts must be disabled.
-
-
-Depth of the PI Chain
----------------------
-
-The maximum depth of the PI chain is not dynamic, and could actually be
-defined. But is very complex to figure it out, since it depends on all
-the nesting of mutexes. Let's look at the example where we have 3 mutexes,
-L1, L2, and L3, and four separate functions func1, func2, func3 and func4.
-The following shows a locking order of L1->L2->L3, but may not actually
-be directly nested that way.
-
-void func1(void)
-{
- mutex_lock(L1);
-
- /* do anything */
-
- mutex_unlock(L1);
-}
-
-void func2(void)
-{
- mutex_lock(L1);
- mutex_lock(L2);
-
- /* do something */
-
- mutex_unlock(L2);
- mutex_unlock(L1);
-}
-
-void func3(void)
-{
- mutex_lock(L2);
- mutex_lock(L3);
-
- /* do something else */
-
- mutex_unlock(L3);
- mutex_unlock(L2);
-}
-
-void func4(void)
-{
- mutex_lock(L3);
-
- /* do something again */
-
- mutex_unlock(L3);
-}
-
-Now we add 4 processes that run each of these functions separately.
-Processes A, B, C, and D which run functions func1, func2, func3 and func4
-respectively, and such that D runs first and A last. With D being preempted
-in func4 in the "do something again" area, we have a locking that follows:
-
-D owns L3
- C blocked on L3
- C owns L2
- B blocked on L2
- B owns L1
- A blocked on L1
-
-And thus we have the chain A->L1->B->L2->C->L3->D.
-
-This gives us a PI depth of 4 (four processes), but looking at any of the
-functions individually, it seems as though they only have at most a locking
-depth of two. So, although the locking depth is defined at compile time,
-it still is very difficult to find the possibilities of that depth.
-
-Now since mutexes can be defined by user-land applications, we don't want a DOS
-type of application that nests large amounts of mutexes to create a large
-PI chain, and have the code holding spin locks while looking at a large
-amount of data. So to prevent this, the implementation not only implements
-a maximum lock depth, but also only holds at most two different locks at a
-time, as it walks the PI chain. More about this below.
-
-
-Mutex owner and flags
----------------------
-
-The mutex structure contains a pointer to the owner of the mutex. If the
-mutex is not owned, this owner is set to NULL. Since all architectures
-have the task structure on at least a four byte alignment (and if this is
-not true, the rtmutex.c code will be broken!), this allows for the two
-least significant bits to be used as flags. This part is also described
-in Documentation/rt-mutex.txt, but will also be briefly described here.
-
-Bit 0 is used as the "Pending Owner" flag. This is described later.
-Bit 1 is used as the "Has Waiters" flags. This is also described later
- in more detail, but is set whenever there are waiters on a mutex.
-
-
-cmpxchg Tricks
---------------
-
-Some architectures implement an atomic cmpxchg (Compare and Exchange). This
-is used (when applicable) to keep the fast path of grabbing and releasing
-mutexes short.
-
-cmpxchg is basically the following function performed atomically:
-
-unsigned long _cmpxchg(unsigned long *A, unsigned long *B, unsigned long *C)
-{
- unsigned long T = *A;
- if (*A == *B) {
- *A = *C;
- }
- return T;
-}
-#define cmpxchg(a,b,c) _cmpxchg(&a,&b,&c)
-
-This is really nice to have, since it allows you to only update a variable
-if the variable is what you expect it to be. You know if it succeeded if
-the return value (the old value of A) is equal to B.
-
-The macro rt_mutex_cmpxchg is used to try to lock and unlock mutexes. If
-the architecture does not support CMPXCHG, then this macro is simply set
-to fail every time. But if CMPXCHG is supported, then this will
-help out extremely to keep the fast path short.
-
-The use of rt_mutex_cmpxchg with the flags in the owner field help optimize
-the system for architectures that support it. This will also be explained
-later in this document.
-
-
-Priority adjustments
---------------------
-
-The implementation of the PI code in rtmutex.c has several places that a
-process must adjust its priority. With the help of the pi_list of a
-process this is rather easy to know what needs to be adjusted.
-
-The functions implementing the task adjustments are rt_mutex_adjust_prio,
-__rt_mutex_adjust_prio (same as the former, but expects the task pi_lock
-to already be taken), rt_mutex_getprio, and rt_mutex_setprio.
-
-rt_mutex_getprio and rt_mutex_setprio are only used in __rt_mutex_adjust_prio.
-
-rt_mutex_getprio returns the priority that the task should have. Either the
-task's own normal priority, or if a process of a higher priority is waiting on
-a mutex owned by the task, then that higher priority should be returned.
-Since the pi_list of a task holds an order by priority list of all the top
-waiters of all the mutexes that the task owns, rt_mutex_getprio simply needs
-to compare the top pi waiter to its own normal priority, and return the higher
-priority back.
-
-(Note: if looking at the code, you will notice that the lower number of
- prio is returned. This is because the prio field in the task structure
- is an inverse order of the actual priority. So a "prio" of 5 is
- of higher priority than a "prio" of 10.)
-
-__rt_mutex_adjust_prio examines the result of rt_mutex_getprio, and if the
-result does not equal the task's current priority, then rt_mutex_setprio
-is called to adjust the priority of the task to the new priority.
-Note that rt_mutex_setprio is defined in kernel/sched/core.c to implement the
-actual change in priority.
-
-It is interesting to note that __rt_mutex_adjust_prio can either increase
-or decrease the priority of the task. In the case that a higher priority
-process has just blocked on a mutex owned by the task, __rt_mutex_adjust_prio
-would increase/boost the task's priority. But if a higher priority task
-were for some reason to leave the mutex (timeout or signal), this same function
-would decrease/unboost the priority of the task. That is because the pi_list
-always contains the highest priority task that is waiting on a mutex owned
-by the task, so we only need to compare the priority of that top pi waiter
-to the normal priority of the given task.
-
-
-High level overview of the PI chain walk
-----------------------------------------
-
-The PI chain walk is implemented by the function rt_mutex_adjust_prio_chain.
-
-The implementation has gone through several iterations, and has ended up
-with what we believe is the best. It walks the PI chain by only grabbing
-at most two locks at a time, and is very efficient.
-
-The rt_mutex_adjust_prio_chain can be used either to boost or lower process
-priorities.
-
-rt_mutex_adjust_prio_chain is called with a task to be checked for PI
-(de)boosting (the owner of a mutex that a process is blocking on), a flag to
-check for deadlocking, the mutex that the task owns, and a pointer to a waiter
-that is the process's waiter struct that is blocked on the mutex (although this
-parameter may be NULL for deboosting).
-
-For this explanation, I will not mention deadlock detection. This explanation
-will try to stay at a high level.
-
-When this function is called, there are no locks held. That also means
-that the state of the owner and lock can change when entered into this function.
-
-Before this function is called, the task has already had rt_mutex_adjust_prio
-performed on it. This means that the task is set to the priority that it
-should be at, but the plist nodes of the task's waiter have not been updated
-with the new priorities, and that this task may not be in the proper locations
-in the pi_lists and wait_lists that the task is blocked on. This function
-solves all that.
-
-A loop is entered, where task is the owner to be checked for PI changes that
-was passed by parameter (for the first iteration). The pi_lock of this task is
-taken to prevent any more changes to the pi_list of the task. This also
-prevents new tasks from completing the blocking on a mutex that is owned by this
-task.
-
-If the task is not blocked on a mutex then the loop is exited. We are at
-the top of the PI chain.
-
-A check is now done to see if the original waiter (the process that is blocked
-on the current mutex) is the top pi waiter of the task. That is, is this
-waiter on the top of the task's pi_list. If it is not, it either means that
-there is another process higher in priority that is blocked on one of the
-mutexes that the task owns, or that the waiter has just woken up via a signal
-or timeout and has left the PI chain. In either case, the loop is exited, since
-we don't need to do any more changes to the priority of the current task, or any
-task that owns a mutex that this current task is waiting on. A priority chain
-walk is only needed when a new top pi waiter is made to a task.
-
-The next check sees if the task's waiter plist node has the priority equal to
-the priority the task is set at. If they are equal, then we are done with
-the loop. Remember that the function started with the priority of the
-task adjusted, but the plist nodes that hold the task in other processes
-pi_lists have not been adjusted.
-
-Next, we look at the mutex that the task is blocked on. The mutex's wait_lock
-is taken. This is done by a spin_trylock, because the locking order of the
-pi_lock and wait_lock goes in the opposite direction. If we fail to grab the
-lock, the pi_lock is released, and we restart the loop.
-
-Now that we have both the pi_lock of the task as well as the wait_lock of
-the mutex the task is blocked on, we update the task's waiter's plist node
-that is located on the mutex's wait_list.
-
-Now we release the pi_lock of the task.
-
-Next the owner of the mutex has its pi_lock taken, so we can update the
-task's entry in the owner's pi_list. If the task is the highest priority
-process on the mutex's wait_list, then we remove the previous top waiter
-from the owner's pi_list, and replace it with the task.
-
-Note: It is possible that the task was the current top waiter on the mutex,
- in which case the task is not yet on the pi_list of the waiter. This
- is OK, since plist_del does nothing if the plist node is not on any
- list.
-
-If the task was not the top waiter of the mutex, but it was before we
-did the priority updates, that means we are deboosting/lowering the
-task. In this case, the task is removed from the pi_list of the owner,
-and the new top waiter is added.
-
-Lastly, we unlock both the pi_lock of the task, as well as the mutex's
-wait_lock, and continue the loop again. On the next iteration of the
-loop, the previous owner of the mutex will be the task that will be
-processed.
-
-Note: One might think that the owner of this mutex might have changed
- since we just grab the mutex's wait_lock. And one could be right.
- The important thing to remember is that the owner could not have
- become the task that is being processed in the PI chain, since
- we have taken that task's pi_lock at the beginning of the loop.
- So as long as there is an owner of this mutex that is not the same
- process as the tasked being worked on, we are OK.
-
- Looking closely at the code, one might be confused. The check for the
- end of the PI chain is when the task isn't blocked on anything or the
- task's waiter structure "task" element is NULL. This check is
- protected only by the task's pi_lock. But the code to unlock the mutex
- sets the task's waiter structure "task" element to NULL with only
- the protection of the mutex's wait_lock, which was not taken yet.
- Isn't this a race condition if the task becomes the new owner?
-
- The answer is No! The trick is the spin_trylock of the mutex's
- wait_lock. If we fail that lock, we release the pi_lock of the
- task and continue the loop, doing the end of PI chain check again.
-
- In the code to release the lock, the wait_lock of the mutex is held
- the entire time, and it is not let go when we grab the pi_lock of the
- new owner of the mutex. So if the switch of a new owner were to happen
- after the check for end of the PI chain and the grabbing of the
- wait_lock, the unlocking code would spin on the new owner's pi_lock
- but never give up the wait_lock. So the PI chain loop is guaranteed to
- fail the spin_trylock on the wait_lock, release the pi_lock, and
- try again.
-
- If you don't quite understand the above, that's OK. You don't have to,
- unless you really want to make a proof out of it ;)
-
-
-Pending Owners and Lock stealing
---------------------------------
-
-One of the flags in the owner field of the mutex structure is "Pending Owner".
-What this means is that an owner was chosen by the process releasing the
-mutex, but that owner has yet to wake up and actually take the mutex.
-
-Why is this important? Why can't we just give the mutex to another process
-and be done with it?
-
-The PI code is to help with real-time processes, and to let the highest
-priority process run as long as possible with little latencies and delays.
-If a high priority process owns a mutex that a lower priority process is
-blocked on, when the mutex is released it would be given to the lower priority
-process. What if the higher priority process wants to take that mutex again.
-The high priority process would fail to take that mutex that it just gave up
-and it would need to boost the lower priority process to run with full
-latency of that critical section (since the low priority process just entered
-it).
-
-There's no reason a high priority process that gives up a mutex should be
-penalized if it tries to take that mutex again. If the new owner of the
-mutex has not woken up yet, there's no reason that the higher priority process
-could not take that mutex away.
-
-To solve this, we introduced Pending Ownership and Lock Stealing. When a
-new process is given a mutex that it was blocked on, it is only given
-pending ownership. This means that it's the new owner, unless a higher
-priority process comes in and tries to grab that mutex. If a higher priority
-process does come along and wants that mutex, we let the higher priority
-process "steal" the mutex from the pending owner (only if it is still pending)
-and continue with the mutex.
-
-
-Taking of a mutex (The walk through)
-------------------------------------
-
-OK, now let's take a look at the detailed walk through of what happens when
-taking a mutex.
-
-The first thing that is tried is the fast taking of the mutex. This is
-done when we have CMPXCHG enabled (otherwise the fast taking automatically
-fails). Only when the owner field of the mutex is NULL can the lock be
-taken with the CMPXCHG and nothing else needs to be done.
-
-If there is contention on the lock, whether it is owned or pending owner
-we go about the slow path (rt_mutex_slowlock).
-
-The slow path function is where the task's waiter structure is created on
-the stack. This is because the waiter structure is only needed for the
-scope of this function. The waiter structure holds the nodes to store
-the task on the wait_list of the mutex, and if need be, the pi_list of
-the owner.
-
-The wait_lock of the mutex is taken since the slow path of unlocking the
-mutex also takes this lock.
-
-We then call try_to_take_rt_mutex. This is where the architecture that
-does not implement CMPXCHG would always grab the lock (if there's no
-contention).
-
-try_to_take_rt_mutex is used every time the task tries to grab a mutex in the
-slow path. The first thing that is done here is an atomic setting of
-the "Has Waiters" flag of the mutex's owner field. Yes, this could really
-be false, because if the mutex has no owner, there are no waiters and
-the current task also won't have any waiters. But we don't have the lock
-yet, so we assume we are going to be a waiter. The reason for this is to
-play nice for those architectures that do have CMPXCHG. By setting this flag
-now, the owner of the mutex can't release the mutex without going into the
-slow unlock path, and it would then need to grab the wait_lock, which this
-code currently holds. So setting the "Has Waiters" flag forces the owner
-to synchronize with this code.
-
-Now that we know that we can't have any races with the owner releasing the
-mutex, we check to see if we can take the ownership. This is done if the
-mutex doesn't have a owner, or if we can steal the mutex from a pending
-owner. Let's look at the situations we have here.
-
- 1) Has owner that is pending
- ----------------------------
-
- The mutex has a owner, but it hasn't woken up and the mutex flag
- "Pending Owner" is set. The first check is to see if the owner isn't the
- current task. This is because this function is also used for the pending
- owner to grab the mutex. When a pending owner wakes up, it checks to see
- if it can take the mutex, and this is done if the owner is already set to
- itself. If so, we succeed and leave the function, clearing the "Pending
- Owner" bit.
-
- If the pending owner is not current, we check to see if the current priority is
- higher than the pending owner. If not, we fail the function and return.
-
- There's also something special about a pending owner. That is a pending owner
- is never blocked on a mutex. So there is no PI chain to worry about. It also
- means that if the mutex doesn't have any waiters, there's no accounting needed
- to update the pending owner's pi_list, since we only worry about processes
- blocked on the current mutex.
-
- If there are waiters on this mutex, and we just stole the ownership, we need
- to take the top waiter, remove it from the pi_list of the pending owner, and
- add it to the current pi_list. Note that at this moment, the pending owner
- is no longer on the list of waiters. This is fine, since the pending owner
- would add itself back when it realizes that it had the ownership stolen
- from itself. When the pending owner tries to grab the mutex, it will fail
- in try_to_take_rt_mutex if the owner field points to another process.
-
- 2) No owner
- -----------
-
- If there is no owner (or we successfully stole the lock), we set the owner
- of the mutex to current, and set the flag of "Has Waiters" if the current
- mutex actually has waiters, or we clear the flag if it doesn't. See, it was
- OK that we set that flag early, since now it is cleared.
-
- 3) Failed to grab ownership
- ---------------------------
-
- The most interesting case is when we fail to take ownership. This means that
- there exists an owner, or there's a pending owner with equal or higher
- priority than the current task.
-
-We'll continue on the failed case.
-
-If the mutex has a timeout, we set up a timer to go off to break us out
-of this mutex if we failed to get it after a specified amount of time.
-
-Now we enter a loop that will continue to try to take ownership of the mutex, or
-fail from a timeout or signal.
-
-Once again we try to take the mutex. This will usually fail the first time
-in the loop, since it had just failed to get the mutex. But the second time
-in the loop, this would likely succeed, since the task would likely be
-the pending owner.
-
-If the mutex is TASK_INTERRUPTIBLE a check for signals and timeout is done
-here.
-
-The waiter structure has a "task" field that points to the task that is blocked
-on the mutex. This field can be NULL the first time it goes through the loop
-or if the task is a pending owner and had its mutex stolen. If the "task"
-field is NULL then we need to set up the accounting for it.
-
-Task blocks on mutex
---------------------
-
-The accounting of a mutex and process is done with the waiter structure of
-the process. The "task" field is set to the process, and the "lock" field
-to the mutex. The plist nodes are initialized to the processes current
-priority.
-
-Since the wait_lock was taken at the entry of the slow lock, we can safely
-add the waiter to the wait_list. If the current process is the highest
-priority process currently waiting on this mutex, then we remove the
-previous top waiter process (if it exists) from the pi_list of the owner,
-and add the current process to that list. Since the pi_list of the owner
-has changed, we call rt_mutex_adjust_prio on the owner to see if the owner
-should adjust its priority accordingly.
-
-If the owner is also blocked on a lock, and had its pi_list changed
-(or deadlock checking is on), we unlock the wait_lock of the mutex and go ahead
-and run rt_mutex_adjust_prio_chain on the owner, as described earlier.
-
-Now all locks are released, and if the current process is still blocked on a
-mutex (waiter "task" field is not NULL), then we go to sleep (call schedule).
-
-Waking up in the loop
----------------------
-
-The schedule can then wake up for a few reasons.
- 1) we were given pending ownership of the mutex.
- 2) we received a signal and was TASK_INTERRUPTIBLE
- 3) we had a timeout and was TASK_INTERRUPTIBLE
-
-In any of these cases, we continue the loop and once again try to grab the
-ownership of the mutex. If we succeed, we exit the loop, otherwise we continue
-and on signal and timeout, will exit the loop, or if we had the mutex stolen
-we just simply add ourselves back on the lists and go back to sleep.
-
-Note: For various reasons, because of timeout and signals, the steal mutex
- algorithm needs to be careful. This is because the current process is
- still on the wait_list. And because of dynamic changing of priorities,
- especially on SCHED_OTHER tasks, the current process can be the
- highest priority task on the wait_list.
-
-Failed to get mutex on Timeout or Signal
-----------------------------------------
-
-If a timeout or signal occurred, the waiter's "task" field would not be
-NULL and the task needs to be taken off the wait_list of the mutex and perhaps
-pi_list of the owner. If this process was a high priority process, then
-the rt_mutex_adjust_prio_chain needs to be executed again on the owner,
-but this time it will be lowering the priorities.
-
-
-Unlocking the Mutex
--------------------
-
-The unlocking of a mutex also has a fast path for those architectures with
-CMPXCHG. Since the taking of a mutex on contention always sets the
-"Has Waiters" flag of the mutex's owner, we use this to know if we need to
-take the slow path when unlocking the mutex. If the mutex doesn't have any
-waiters, the owner field of the mutex would equal the current process and
-the mutex can be unlocked by just replacing the owner field with NULL.
-
-If the owner field has the "Has Waiters" bit set (or CMPXCHG is not available),
-the slow unlock path is taken.
-
-The first thing done in the slow unlock path is to take the wait_lock of the
-mutex. This synchronizes the locking and unlocking of the mutex.
-
-A check is made to see if the mutex has waiters or not. On architectures that
-do not have CMPXCHG, this is the location that the owner of the mutex will
-determine if a waiter needs to be awoken or not. On architectures that
-do have CMPXCHG, that check is done in the fast path, but it is still needed
-in the slow path too. If a waiter of a mutex woke up because of a signal
-or timeout between the time the owner failed the fast path CMPXCHG check and
-the grabbing of the wait_lock, the mutex may not have any waiters, thus the
-owner still needs to make this check. If there are no waiters then the mutex
-owner field is set to NULL, the wait_lock is released and nothing more is
-needed.
-
-If there are waiters, then we need to wake one up and give that waiter
-pending ownership.
-
-On the wake up code, the pi_lock of the current owner is taken. The top
-waiter of the lock is found and removed from the wait_list of the mutex
-as well as the pi_list of the current owner. The task field of the new
-pending owner's waiter structure is set to NULL, and the owner field of the
-mutex is set to the new owner with the "Pending Owner" bit set, as well
-as the "Has Waiters" bit if there still are other processes blocked on the
-mutex.
-
-The pi_lock of the previous owner is released, and the new pending owner's
-pi_lock is taken. Remember that this is the trick to prevent the race
-condition in rt_mutex_adjust_prio_chain from adding itself as a waiter
-on the mutex.
-
-We now clear the "pi_blocked_on" field of the new pending owner, and if
-the mutex still has waiters pending, we add the new top waiter to the pi_list
-of the pending owner.
-
-Finally we unlock the pi_lock of the pending owner and wake it up.
-
-
-Contact
--------
-
-For updates on this document, please email Steven Rostedt <[email protected]>
-
-
-Credits
--------
-
-Author: Steven Rostedt <[email protected]>
-
-Reviewers: Ingo Molnar, Thomas Gleixner, Thomas Duetsch, and Randy Dunlap
-
-Updates
--------
-
-This document was originally written for 2.6.17-rc3-mm1
diff --git a/Documentation/rt-mutex.txt b/Documentation/rt-mutex.txt
deleted file mode 100644
index 243393d..0000000
--- a/Documentation/rt-mutex.txt
+++ /dev/null
@@ -1,79 +0,0 @@
-RT-mutex subsystem with PI support
-----------------------------------
-
-RT-mutexes with priority inheritance are used to support PI-futexes,
-which enable pthread_mutex_t priority inheritance attributes
-(PTHREAD_PRIO_INHERIT). [See Documentation/pi-futex.txt for more details
-about PI-futexes.]
-
-This technology was developed in the -rt tree and streamlined for
-pthread_mutex support.
-
-Basic principles:
------------------
-
-RT-mutexes extend the semantics of simple mutexes by the priority
-inheritance protocol.
-
-A low priority owner of a rt-mutex inherits the priority of a higher
-priority waiter until the rt-mutex is released. If the temporarily
-boosted owner blocks on a rt-mutex itself it propagates the priority
-boosting to the owner of the other rt_mutex it gets blocked on. The
-priority boosting is immediately removed once the rt_mutex has been
-unlocked.
-
-This approach allows us to shorten the block of high-prio tasks on
-mutexes which protect shared resources. Priority inheritance is not a
-magic bullet for poorly designed applications, but it allows
-well-designed applications to use userspace locks in critical parts of
-an high priority thread, without losing determinism.
-
-The enqueueing of the waiters into the rtmutex waiter list is done in
-priority order. For same priorities FIFO order is chosen. For each
-rtmutex, only the top priority waiter is enqueued into the owner's
-priority waiters list. This list too queues in priority order. Whenever
-the top priority waiter of a task changes (for example it timed out or
-got a signal), the priority of the owner task is readjusted. [The
-priority enqueueing is handled by "plists", see include/linux/plist.h
-for more details.]
-
-RT-mutexes are optimized for fastpath operations and have no internal
-locking overhead when locking an uncontended mutex or unlocking a mutex
-without waiters. The optimized fastpath operations require cmpxchg
-support. [If that is not available then the rt-mutex internal spinlock
-is used]
-
-The state of the rt-mutex is tracked via the owner field of the rt-mutex
-structure:
-
-rt_mutex->owner holds the task_struct pointer of the owner. Bit 0 and 1
-are used to keep track of the "owner is pending" and "rtmutex has
-waiters" state.
-
- owner bit1 bit0
- NULL 0 0 mutex is free (fast acquire possible)
- NULL 0 1 invalid state
- NULL 1 0 Transitional state*
- NULL 1 1 invalid state
- taskpointer 0 0 mutex is held (fast release possible)
- taskpointer 0 1 task is pending owner
- taskpointer 1 0 mutex is held and has waiters
- taskpointer 1 1 task is pending owner and mutex has waiters
-
-Pending-ownership handling is a performance optimization:
-pending-ownership is assigned to the first (highest priority) waiter of
-the mutex, when the mutex is released. The thread is woken up and once
-it starts executing it can acquire the mutex. Until the mutex is taken
-by it (bit 0 is cleared) a competing higher priority thread can "steal"
-the mutex which puts the woken up thread back on the waiters list.
-
-The pending-ownership optimization is especially important for the
-uninterrupted workflow of high-prio tasks which repeatedly
-takes/releases locks that have lower-prio waiters. Without this
-optimization the higher-prio thread would ping-pong to the lower-prio
-task [because at unlock time we always assign a new owner].
-
-(*) The "mutex has waiters" bit gets set to take the lock. If the lock
-doesn't already have an owner, this bit is quickly cleared if there are
-no waiters. So this is a transitional state to synchronize with looking
-at the owner field of the mutex and the mutex owner releasing the lock.
diff --git a/Documentation/spinlocks.txt b/Documentation/spinlocks.txt
deleted file mode 100644
index 97eaf57..0000000
--- a/Documentation/spinlocks.txt
+++ /dev/null
@@ -1,167 +0,0 @@
-Lesson 1: Spin locks
-
-The most basic primitive for locking is spinlock.
-
-static DEFINE_SPINLOCK(xxx_lock);
-
- unsigned long flags;
-
- spin_lock_irqsave(&xxx_lock, flags);
- ... critical section here ..
- spin_unlock_irqrestore(&xxx_lock, flags);
-
-The above is always safe. It will disable interrupts _locally_, but the
-spinlock itself will guarantee the global lock, so it will guarantee that
-there is only one thread-of-control within the region(s) protected by that
-lock. This works well even under UP also, so the code does _not_ need to
-worry about UP vs SMP issues: the spinlocks work correctly under both.
-
- NOTE! Implications of spin_locks for memory are further described in:
-
- Documentation/memory-barriers.txt
- (5) LOCK operations.
- (6) UNLOCK operations.
-
-The above is usually pretty simple (you usually need and want only one
-spinlock for most things - using more than one spinlock can make things a
-lot more complex and even slower and is usually worth it only for
-sequences that you _know_ need to be split up: avoid it at all cost if you
-aren't sure).
-
-This is really the only really hard part about spinlocks: once you start
-using spinlocks they tend to expand to areas you might not have noticed
-before, because you have to make sure the spinlocks correctly protect the
-shared data structures _everywhere_ they are used. The spinlocks are most
-easily added to places that are completely independent of other code (for
-example, internal driver data structures that nobody else ever touches).
-
- NOTE! The spin-lock is safe only when you _also_ use the lock itself
- to do locking across CPU's, which implies that EVERYTHING that
- touches a shared variable has to agree about the spinlock they want
- to use.
-
-----
-
-Lesson 2: reader-writer spinlocks.
-
-If your data accesses have a very natural pattern where you usually tend
-to mostly read from the shared variables, the reader-writer locks
-(rw_lock) versions of the spinlocks are sometimes useful. They allow multiple
-readers to be in the same critical region at once, but if somebody wants
-to change the variables it has to get an exclusive write lock.
-
- NOTE! reader-writer locks require more atomic memory operations than
- simple spinlocks. Unless the reader critical section is long, you
- are better off just using spinlocks.
-
-The routines look the same as above:
-
- rwlock_t xxx_lock = __RW_LOCK_UNLOCKED(xxx_lock);
-
- unsigned long flags;
-
- read_lock_irqsave(&xxx_lock, flags);
- .. critical section that only reads the info ...
- read_unlock_irqrestore(&xxx_lock, flags);
-
- write_lock_irqsave(&xxx_lock, flags);
- .. read and write exclusive access to the info ...
- write_unlock_irqrestore(&xxx_lock, flags);
-
-The above kind of lock may be useful for complex data structures like
-linked lists, especially searching for entries without changing the list
-itself. The read lock allows many concurrent readers. Anything that
-_changes_ the list will have to get the write lock.
-
- NOTE! RCU is better for list traversal, but requires careful
- attention to design detail (see Documentation/RCU/listRCU.txt).
-
-Also, you cannot "upgrade" a read-lock to a write-lock, so if you at _any_
-time need to do any changes (even if you don't do it every time), you have
-to get the write-lock at the very beginning.
-
- NOTE! We are working hard to remove reader-writer spinlocks in most
- cases, so please don't add a new one without consensus. (Instead, see
- Documentation/RCU/rcu.txt for complete information.)
-
-----
-
-Lesson 3: spinlocks revisited.
-
-The single spin-lock primitives above are by no means the only ones. They
-are the most safe ones, and the ones that work under all circumstances,
-but partly _because_ they are safe they are also fairly slow. They are slower
-than they'd need to be, because they do have to disable interrupts
-(which is just a single instruction on a x86, but it's an expensive one -
-and on other architectures it can be worse).
-
-If you have a case where you have to protect a data structure across
-several CPU's and you want to use spinlocks you can potentially use
-cheaper versions of the spinlocks. IFF you know that the spinlocks are
-never used in interrupt handlers, you can use the non-irq versions:
-
- spin_lock(&lock);
- ...
- spin_unlock(&lock);
-
-(and the equivalent read-write versions too, of course). The spinlock will
-guarantee the same kind of exclusive access, and it will be much faster.
-This is useful if you know that the data in question is only ever
-manipulated from a "process context", ie no interrupts involved.
-
-The reasons you mustn't use these versions if you have interrupts that
-play with the spinlock is that you can get deadlocks:
-
- spin_lock(&lock);
- ...
- <- interrupt comes in:
- spin_lock(&lock);
-
-where an interrupt tries to lock an already locked variable. This is ok if
-the other interrupt happens on another CPU, but it is _not_ ok if the
-interrupt happens on the same CPU that already holds the lock, because the
-lock will obviously never be released (because the interrupt is waiting
-for the lock, and the lock-holder is interrupted by the interrupt and will
-not continue until the interrupt has been processed).
-
-(This is also the reason why the irq-versions of the spinlocks only need
-to disable the _local_ interrupts - it's ok to use spinlocks in interrupts
-on other CPU's, because an interrupt on another CPU doesn't interrupt the
-CPU that holds the lock, so the lock-holder can continue and eventually
-releases the lock).
-
-Note that you can be clever with read-write locks and interrupts. For
-example, if you know that the interrupt only ever gets a read-lock, then
-you can use a non-irq version of read locks everywhere - because they
-don't block on each other (and thus there is no dead-lock wrt interrupts.
-But when you do the write-lock, you have to use the irq-safe version.
-
-For an example of being clever with rw-locks, see the "waitqueue_lock"
-handling in kernel/sched/core.c - nothing ever _changes_ a wait-queue from
-within an interrupt, they only read the queue in order to know whom to
-wake up. So read-locks are safe (which is good: they are very common
-indeed), while write-locks need to protect themselves against interrupts.
-
- Linus
-
-----
-
-Reference information:
-
-For dynamic initialization, use spin_lock_init() or rwlock_init() as
-appropriate:
-
- spinlock_t xxx_lock;
- rwlock_t xxx_rw_lock;
-
- static int __init xxx_init(void)
- {
- spin_lock_init(&xxx_lock);
- rwlock_init(&xxx_rw_lock);
- ...
- }
-
- module_init(xxx_init);
-
-For static initialization, use DEFINE_SPINLOCK() / DEFINE_RWLOCK() or
-__SPIN_LOCK_UNLOCKED() / __RW_LOCK_UNLOCKED() as appropriate.
diff --git a/Documentation/ww-mutex-design.txt b/Documentation/ww-mutex-design.txt
deleted file mode 100644
index 8a112dc..0000000
--- a/Documentation/ww-mutex-design.txt
+++ /dev/null
@@ -1,344 +0,0 @@
-Wait/Wound Deadlock-Proof Mutex Design
-======================================
-
-Please read mutex-design.txt first, as it applies to wait/wound mutexes too.
-
-Motivation for WW-Mutexes
--------------------------
-
-GPU's do operations that commonly involve many buffers. Those buffers
-can be shared across contexts/processes, exist in different memory
-domains (for example VRAM vs system memory), and so on. And with
-PRIME / dmabuf, they can even be shared across devices. So there are
-a handful of situations where the driver needs to wait for buffers to
-become ready. If you think about this in terms of waiting on a buffer
-mutex for it to become available, this presents a problem because
-there is no way to guarantee that buffers appear in a execbuf/batch in
-the same order in all contexts. That is directly under control of
-userspace, and a result of the sequence of GL calls that an application
-makes. Which results in the potential for deadlock. The problem gets
-more complex when you consider that the kernel may need to migrate the
-buffer(s) into VRAM before the GPU operates on the buffer(s), which
-may in turn require evicting some other buffers (and you don't want to
-evict other buffers which are already queued up to the GPU), but for a
-simplified understanding of the problem you can ignore this.
-
-The algorithm that the TTM graphics subsystem came up with for dealing with
-this problem is quite simple. For each group of buffers (execbuf) that need
-to be locked, the caller would be assigned a unique reservation id/ticket,
-from a global counter. In case of deadlock while locking all the buffers
-associated with a execbuf, the one with the lowest reservation ticket (i.e.
-the oldest task) wins, and the one with the higher reservation id (i.e. the
-younger task) unlocks all of the buffers that it has already locked, and then
-tries again.
-
-In the RDBMS literature this deadlock handling approach is called wait/wound:
-The older tasks waits until it can acquire the contended lock. The younger tasks
-needs to back off and drop all the locks it is currently holding, i.e. the
-younger task is wounded.
-
-Concepts
---------
-
-Compared to normal mutexes two additional concepts/objects show up in the lock
-interface for w/w mutexes:
-
-Acquire context: To ensure eventual forward progress it is important the a task
-trying to acquire locks doesn't grab a new reservation id, but keeps the one it
-acquired when starting the lock acquisition. This ticket is stored in the
-acquire context. Furthermore the acquire context keeps track of debugging state
-to catch w/w mutex interface abuse.
-
-W/w class: In contrast to normal mutexes the lock class needs to be explicit for
-w/w mutexes, since it is required to initialize the acquire context.
-
-Furthermore there are three different class of w/w lock acquire functions:
-
-* Normal lock acquisition with a context, using ww_mutex_lock.
-
-* Slowpath lock acquisition on the contending lock, used by the wounded task
- after having dropped all already acquired locks. These functions have the
- _slow postfix.
-
- From a simple semantics point-of-view the _slow functions are not strictly
- required, since simply calling the normal ww_mutex_lock functions on the
- contending lock (after having dropped all other already acquired locks) will
- work correctly. After all if no other ww mutex has been acquired yet there's
- no deadlock potential and hence the ww_mutex_lock call will block and not
- prematurely return -EDEADLK. The advantage of the _slow functions is in
- interface safety:
- - ww_mutex_lock has a __must_check int return type, whereas ww_mutex_lock_slow
- has a void return type. Note that since ww mutex code needs loops/retries
- anyway the __must_check doesn't result in spurious warnings, even though the
- very first lock operation can never fail.
- - When full debugging is enabled ww_mutex_lock_slow checks that all acquired
- ww mutex have been released (preventing deadlocks) and makes sure that we
- block on the contending lock (preventing spinning through the -EDEADLK
- slowpath until the contended lock can be acquired).
-
-* Functions to only acquire a single w/w mutex, which results in the exact same
- semantics as a normal mutex. This is done by calling ww_mutex_lock with a NULL
- context.
-
- Again this is not strictly required. But often you only want to acquire a
- single lock in which case it's pointless to set up an acquire context (and so
- better to avoid grabbing a deadlock avoidance ticket).
-
-Of course, all the usual variants for handling wake-ups due to signals are also
-provided.
-
-Usage
------
-
-Three different ways to acquire locks within the same w/w class. Common
-definitions for methods #1 and #2:
-
-static DEFINE_WW_CLASS(ww_class);
-
-struct obj {
- struct ww_mutex lock;
- /* obj data */
-};
-
-struct obj_entry {
- struct list_head head;
- struct obj *obj;
-};
-
-Method 1, using a list in execbuf->buffers that's not allowed to be reordered.
-This is useful if a list of required objects is already tracked somewhere.
-Furthermore the lock helper can use propagate the -EALREADY return code back to
-the caller as a signal that an object is twice on the list. This is useful if
-the list is constructed from userspace input and the ABI requires userspace to
-not have duplicate entries (e.g. for a gpu commandbuffer submission ioctl).
-
-int lock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
-{
- struct obj *res_obj = NULL;
- struct obj_entry *contended_entry = NULL;
- struct obj_entry *entry;
-
- ww_acquire_init(ctx, &ww_class);
-
-retry:
- list_for_each_entry (entry, list, head) {
- if (entry->obj == res_obj) {
- res_obj = NULL;
- continue;
- }
- ret = ww_mutex_lock(&entry->obj->lock, ctx);
- if (ret < 0) {
- contended_entry = entry;
- goto err;
- }
- }
-
- ww_acquire_done(ctx);
- return 0;
-
-err:
- list_for_each_entry_continue_reverse (entry, list, head)
- ww_mutex_unlock(&entry->obj->lock);
-
- if (res_obj)
- ww_mutex_unlock(&res_obj->lock);
-
- if (ret == -EDEADLK) {
- /* we lost out in a seqno race, lock and retry.. */
- ww_mutex_lock_slow(&contended_entry->obj->lock, ctx);
- res_obj = contended_entry->obj;
- goto retry;
- }
- ww_acquire_fini(ctx);
-
- return ret;
-}
-
-Method 2, using a list in execbuf->buffers that can be reordered. Same semantics
-of duplicate entry detection using -EALREADY as method 1 above. But the
-list-reordering allows for a bit more idiomatic code.
-
-int lock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
-{
- struct obj_entry *entry, *entry2;
-
- ww_acquire_init(ctx, &ww_class);
-
- list_for_each_entry (entry, list, head) {
- ret = ww_mutex_lock(&entry->obj->lock, ctx);
- if (ret < 0) {
- entry2 = entry;
-
- list_for_each_entry_continue_reverse (entry2, list, head)
- ww_mutex_unlock(&entry2->obj->lock);
-
- if (ret != -EDEADLK) {
- ww_acquire_fini(ctx);
- return ret;
- }
-
- /* we lost out in a seqno race, lock and retry.. */
- ww_mutex_lock_slow(&entry->obj->lock, ctx);
-
- /*
- * Move buf to head of the list, this will point
- * buf->next to the first unlocked entry,
- * restarting the for loop.
- */
- list_del(&entry->head);
- list_add(&entry->head, list);
- }
- }
-
- ww_acquire_done(ctx);
- return 0;
-}
-
-Unlocking works the same way for both methods #1 and #2:
-
-void unlock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
-{
- struct obj_entry *entry;
-
- list_for_each_entry (entry, list, head)
- ww_mutex_unlock(&entry->obj->lock);
-
- ww_acquire_fini(ctx);
-}
-
-Method 3 is useful if the list of objects is constructed ad-hoc and not upfront,
-e.g. when adjusting edges in a graph where each node has its own ww_mutex lock,
-and edges can only be changed when holding the locks of all involved nodes. w/w
-mutexes are a natural fit for such a case for two reasons:
-- They can handle lock-acquisition in any order which allows us to start walking
- a graph from a starting point and then iteratively discovering new edges and
- locking down the nodes those edges connect to.
-- Due to the -EALREADY return code signalling that a given objects is already
- held there's no need for additional book-keeping to break cycles in the graph
- or keep track off which looks are already held (when using more than one node
- as a starting point).
-
-Note that this approach differs in two important ways from the above methods:
-- Since the list of objects is dynamically constructed (and might very well be
- different when retrying due to hitting the -EDEADLK wound condition) there's
- no need to keep any object on a persistent list when it's not locked. We can
- therefore move the list_head into the object itself.
-- On the other hand the dynamic object list construction also means that the -EALREADY return
- code can't be propagated.
-
-Note also that methods #1 and #2 and method #3 can be combined, e.g. to first lock a
-list of starting nodes (passed in from userspace) using one of the above
-methods. And then lock any additional objects affected by the operations using
-method #3 below. The backoff/retry procedure will be a bit more involved, since
-when the dynamic locking step hits -EDEADLK we also need to unlock all the
-objects acquired with the fixed list. But the w/w mutex debug checks will catch
-any interface misuse for these cases.
-
-Also, method 3 can't fail the lock acquisition step since it doesn't return
--EALREADY. Of course this would be different when using the _interruptible
-variants, but that's outside of the scope of these examples here.
-
-struct obj {
- struct ww_mutex ww_mutex;
- struct list_head locked_list;
-};
-
-static DEFINE_WW_CLASS(ww_class);
-
-void __unlock_objs(struct list_head *list)
-{
- struct obj *entry, *temp;
-
- list_for_each_entry_safe (entry, temp, list, locked_list) {
- /* need to do that before unlocking, since only the current lock holder is
- allowed to use object */
- list_del(&entry->locked_list);
- ww_mutex_unlock(entry->ww_mutex)
- }
-}
-
-void lock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
-{
- struct obj *obj;
-
- ww_acquire_init(ctx, &ww_class);
-
-retry:
- /* re-init loop start state */
- loop {
- /* magic code which walks over a graph and decides which objects
- * to lock */
-
- ret = ww_mutex_lock(obj->ww_mutex, ctx);
- if (ret == -EALREADY) {
- /* we have that one already, get to the next object */
- continue;
- }
- if (ret == -EDEADLK) {
- __unlock_objs(list);
-
- ww_mutex_lock_slow(obj, ctx);
- list_add(&entry->locked_list, list);
- goto retry;
- }
-
- /* locked a new object, add it to the list */
- list_add_tail(&entry->locked_list, list);
- }
-
- ww_acquire_done(ctx);
- return 0;
-}
-
-void unlock_objs(struct list_head *list, struct ww_acquire_ctx *ctx)
-{
- __unlock_objs(list);
- ww_acquire_fini(ctx);
-}
-
-Method 4: Only lock one single objects. In that case deadlock detection and
-prevention is obviously overkill, since with grabbing just one lock you can't
-produce a deadlock within just one class. To simplify this case the w/w mutex
-api can be used with a NULL context.
-
-Implementation Details
-----------------------
-
-Design:
- ww_mutex currently encapsulates a struct mutex, this means no extra overhead for
- normal mutex locks, which are far more common. As such there is only a small
- increase in code size if wait/wound mutexes are not used.
-
- In general, not much contention is expected. The locks are typically used to
- serialize access to resources for devices. The only way to make wakeups
- smarter would be at the cost of adding a field to struct mutex_waiter. This
- would add overhead to all cases where normal mutexes are used, and
- ww_mutexes are generally less performance sensitive.
-
-Lockdep:
- Special care has been taken to warn for as many cases of api abuse
- as possible. Some common api abuses will be caught with
- CONFIG_DEBUG_MUTEXES, but CONFIG_PROVE_LOCKING is recommended.
-
- Some of the errors which will be warned about:
- - Forgetting to call ww_acquire_fini or ww_acquire_init.
- - Attempting to lock more mutexes after ww_acquire_done.
- - Attempting to lock the wrong mutex after -EDEADLK and
- unlocking all mutexes.
- - Attempting to lock the right mutex after -EDEADLK,
- before unlocking all mutexes.
-
- - Calling ww_mutex_lock_slow before -EDEADLK was returned.
-
- - Unlocking mutexes with the wrong unlock function.
- - Calling one of the ww_acquire_* twice on the same context.
- - Using a different ww_class for the mutex than for the ww_acquire_ctx.
- - Normal lockdep errors that can result in deadlocks.
-
- Some of the lockdep errors that can result in deadlocks:
- - Calling ww_acquire_init to initialize a second ww_acquire_ctx before
- having called ww_acquire_fini on the first.
- - 'normal' deadlocks that can occur.
-
-FIXME: Update this section once we have the TASK_DEADLOCK task state flag magic
-implemented.
diff --git a/MAINTAINERS b/MAINTAINERS
index b4534ee..053d8c2 100644
--- a/MAINTAINERS
+++ b/MAINTAINERS
@@ -5516,8 +5516,8 @@ M: Ingo Molnar <[email protected]>
L: [email protected]
T: git git://git.kernel.org/pub/scm/linux/kernel/git/tip/tip.git core/locking
S: Maintained
-F: Documentation/lockdep*.txt
-F: Documentation/lockstat.txt
+F: Documentation/locking/lockdep*.txt
+F: Documentation/locking/lockstat.txt
F: include/linux/lockdep.h
F: kernel/locking/
diff --git a/drivers/gpu/drm/drm_modeset_lock.c b/drivers/gpu/drm/drm_modeset_lock.c
index 0dc57d5..3a02e5e 100644
--- a/drivers/gpu/drm/drm_modeset_lock.c
+++ b/drivers/gpu/drm/drm_modeset_lock.c
@@ -35,7 +35,7 @@
* of extra utility/tracking out of our acquire-ctx. This is provided
* by drm_modeset_lock / drm_modeset_acquire_ctx.
*
- * For basic principles of ww_mutex, see: Documentation/ww-mutex-design.txt
+ * For basic principles of ww_mutex, see: Documentation/locking/ww-mutex-design.txt
*
* The basic usage pattern is to:
*
diff --git a/include/linux/lockdep.h b/include/linux/lockdep.h
index 008388f9..f388481 100644
--- a/include/linux/lockdep.h
+++ b/include/linux/lockdep.h
@@ -4,7 +4,7 @@
* Copyright (C) 2006,2007 Red Hat, Inc., Ingo Molnar <[email protected]>
* Copyright (C) 2007 Red Hat, Inc., Peter Zijlstra <[email protected]>
*
- * see Documentation/lockdep-design.txt for more details.
+ * see Documentation/locking/lockdep-design.txt for more details.
*/
#ifndef __LINUX_LOCKDEP_H
#define __LINUX_LOCKDEP_H
diff --git a/include/linux/mutex.h b/include/linux/mutex.h
index e4c2941..cc31498 100644
--- a/include/linux/mutex.h
+++ b/include/linux/mutex.h
@@ -133,7 +133,7 @@ static inline int mutex_is_locked(struct mutex *lock)
/*
* See kernel/locking/mutex.c for detailed documentation of these APIs.
- * Also see Documentation/mutex-design.txt.
+ * Also see Documentation/locking/mutex-design.txt.
*/
#ifdef CONFIG_DEBUG_LOCK_ALLOC
extern void mutex_lock_nested(struct mutex *lock, unsigned int subclass);
diff --git a/include/linux/rwsem.h b/include/linux/rwsem.h
index 035d3c5..8f498cd 100644
--- a/include/linux/rwsem.h
+++ b/include/linux/rwsem.h
@@ -149,7 +149,7 @@ extern void downgrade_write(struct rw_semaphore *sem);
* static then another method for expressing nested locking is
* the explicit definition of lock class keys and the use of
* lockdep_set_class() at lock initialization time.
- * See Documentation/lockdep-design.txt for more details.)
+ * See Documentation/locking/lockdep-design.txt for more details.)
*/
extern void down_read_nested(struct rw_semaphore *sem, int subclass);
extern void down_write_nested(struct rw_semaphore *sem, int subclass);
diff --git a/kernel/locking/mutex.c b/kernel/locking/mutex.c
index 0d8b6ed..dadbf88 100644
--- a/kernel/locking/mutex.c
+++ b/kernel/locking/mutex.c
@@ -15,7 +15,7 @@
* by Steven Rostedt, based on work by Gregory Haskins, Peter Morreale
* and Sven Dietrich.
*
- * Also see Documentation/mutex-design.txt.
+ * Also see Documentation/locking/mutex-design.txt.
*/
#include <linux/mutex.h>
#include <linux/ww_mutex.h>
diff --git a/kernel/locking/rtmutex.c b/kernel/locking/rtmutex.c
index a0ea2a1..7c98873 100644
--- a/kernel/locking/rtmutex.c
+++ b/kernel/locking/rtmutex.c
@@ -8,7 +8,7 @@
* Copyright (C) 2005 Kihon Technologies Inc., Steven Rostedt
* Copyright (C) 2006 Esben Nielsen
*
- * See Documentation/rt-mutex-design.txt for details.
+ * See Documentation/locking/rt-mutex-design.txt for details.
*/
#include <linux/spinlock.h>
#include <linux/export.h>
diff --git a/lib/Kconfig.debug b/lib/Kconfig.debug
index 03f461f..269da61 100644
--- a/lib/Kconfig.debug
+++ b/lib/Kconfig.debug
@@ -924,7 +924,7 @@ config PROVE_LOCKING
the proof of observed correctness is also maintained for an
arbitrary combination of these separate locking variants.
- For more details, see Documentation/lockdep-design.txt.
+ For more details, see Documentation/locking/lockdep-design.txt.
config LOCKDEP
bool
@@ -945,7 +945,7 @@ config LOCK_STAT
help
This feature enables tracking lock contention points
- For more details, see Documentation/lockstat.txt
+ For more details, see Documentation/locking/lockstat.txt
This also enables lock events required by "perf lock",
subcommand of perf.
--
1.8.1.4
Fortunately Jason was able to reduce some of the overhead we
had introduced in the original rwsem optimistic spinning -
an it is now the same size as mutexes. Update the documentation
accordingly.
Acked-by: Jason Low <[email protected]>
Signed-off-by: Davidlohr Bueso <[email protected]>
---
Documentation/locking/mutex-design.txt | 6 +++---
1 file changed, 3 insertions(+), 3 deletions(-)
diff --git a/Documentation/locking/mutex-design.txt b/Documentation/locking/mutex-design.txt
index ee231ed..60c482d 100644
--- a/Documentation/locking/mutex-design.txt
+++ b/Documentation/locking/mutex-design.txt
@@ -145,9 +145,9 @@ Disadvantages
Unlike its original design and purpose, 'struct mutex' is larger than
most locks in the kernel. E.g: on x86-64 it is 40 bytes, almost twice
-as large as 'struct semaphore' (24 bytes) and 8 bytes shy of the
-'struct rw_semaphore' variant. Larger structure sizes mean more CPU
-cache and memory footprint.
+as large as 'struct semaphore' (24 bytes) and tied, along with rwsems,
+for the largest lock in the kernel. Larger structure sizes mean more
+CPU cache and memory footprint.
When to use mutexes
-------------------
--
1.8.1.4
When unlocking, we always want to reach the slowpath with the lock's counter
indicating it is unlocked. -- as returned by the asm fastpath call or by
explicitly setting it. While doing so, at least in theory, we can optimize
and allow faster lock stealing.
When unlocking, we always want to reach the slowpath with the lock's counter
indicating it is unlocked. -- as returned by the asm fastpath call or by
explicitly setting it. While doing so, at least in theory, we can optimize
and allow faster lock stealing.
Signed-off-by: Davidlohr Bueso <[email protected]>
---
Changes from v1:
- Moved comment about value of the counter below to make sense only if
the fastpath leaves the counter unlocked.
kernel/locking/mutex.c | 11 +++++++++--
1 file changed, 9 insertions(+), 2 deletions(-)
diff --git a/kernel/locking/mutex.c b/kernel/locking/mutex.c
index ad0e333..93bec48 100644
--- a/kernel/locking/mutex.c
+++ b/kernel/locking/mutex.c
@@ -684,9 +684,16 @@ __mutex_unlock_common_slowpath(struct mutex *lock, int nested)
unsigned long flags;
/*
- * some architectures leave the lock unlocked in the fastpath failure
+ * As a performance measurement, release the lock before doing other
+ * wakeup related duties to follow. This allows other tasks to acquire
+ * the lock sooner, while still handling cleanups in past unlock calls.
+ * This can be done as we do not enforce strict equivalence between the
+ * mutex counter and wait_list.
+ *
+ *
+ * Some architectures leave the lock unlocked in the fastpath failure
* case, others need to leave it locked. In the later case we have to
- * unlock it here
+ * unlock it here - as the lock counter is currently 0 or negative.
*/
if (__mutex_slowpath_needs_to_unlock())
atomic_set(&lock->count, 1);
--
1.8.1.4
... as we clearly inline mcs_spin_lock() now.
Acked-by: Jason Low <[email protected]>
Signed-off-by: Davidlohr Bueso <[email protected]>
---
kernel/locking/mcs_spinlock.h | 3 ---
1 file changed, 3 deletions(-)
diff --git a/kernel/locking/mcs_spinlock.h b/kernel/locking/mcs_spinlock.h
index 23e89c5..4d60986 100644
--- a/kernel/locking/mcs_spinlock.h
+++ b/kernel/locking/mcs_spinlock.h
@@ -56,9 +56,6 @@ do { \
* If the lock has already been acquired, then this will proceed to spin
* on this node->locked until the previous lock holder sets the node->locked
* in mcs_spin_unlock().
- *
- * We don't inline mcs_spin_lock() so that perf can correctly account for the
- * time spent in this lock function.
*/
static inline
void mcs_spin_lock(struct mcs_spinlock **lock, struct mcs_spinlock *node)
--
1.8.1.4
On 07/30/14 13:41, Davidlohr Bueso wrote:
> Specifically:
> Documentation/locking/lockdep-design.txt
> Documentation/locking/lockstat.txt
> Documentation/locking/mutex-design.txt
> Documentation/locking/rt-mutex-design.txt
> Documentation/locking/rt-mutex.txt
> Documentation/locking/spinlocks.txt
> Documentation/locking/ww-mutex-design.txt
>
> Signed-off-by: Davidlohr Bueso <[email protected]>
Did you read and use Documentation/SubmittingPatches? see:
Please use "git diff -M --stat --summary" to generate the diffstat:
the -M enables rename detection, and the summary enables a summary of
new/deleted or renamed files.
With rename detection, the statistics are rather different [...]
because git will notice that a fair number of the changes are renames.
We really want to see the renames when possible. Otherwise it's good.
Acked-by: Randy Dunlap <[email protected]>
Thanks.
> ---
> Documentation/00-INDEX | 2 +
> Documentation/DocBook/kernel-locking.tmpl | 2 +-
> Documentation/lockdep-design.txt | 286 -----------
> Documentation/locking/lockdep-design.txt | 286 +++++++++++
> Documentation/locking/lockstat.txt | 178 +++++++
> Documentation/locking/mutex-design.txt | 157 ++++++
> Documentation/locking/rt-mutex-design.txt | 781 ++++++++++++++++++++++++++++++
> Documentation/locking/rt-mutex.txt | 79 +++
> Documentation/locking/spinlocks.txt | 167 +++++++
> Documentation/locking/ww-mutex-design.txt | 344 +++++++++++++
> Documentation/lockstat.txt | 178 -------
> Documentation/mutex-design.txt | 157 ------
> Documentation/rt-mutex-design.txt | 781 ------------------------------
> Documentation/rt-mutex.txt | 79 ---
> Documentation/spinlocks.txt | 167 -------
> Documentation/ww-mutex-design.txt | 344 -------------
> MAINTAINERS | 4 +-
> drivers/gpu/drm/drm_modeset_lock.c | 2 +-
> include/linux/lockdep.h | 2 +-
> include/linux/mutex.h | 2 +-
> include/linux/rwsem.h | 2 +-
> kernel/locking/mutex.c | 2 +-
> kernel/locking/rtmutex.c | 2 +-
> lib/Kconfig.debug | 4 +-
> 24 files changed, 2005 insertions(+), 2003 deletions(-)
> delete mode 100644 Documentation/lockdep-design.txt
> create mode 100644 Documentation/locking/lockdep-design.txt
> create mode 100644 Documentation/locking/lockstat.txt
> create mode 100644 Documentation/locking/mutex-design.txt
> create mode 100644 Documentation/locking/rt-mutex-design.txt
> create mode 100644 Documentation/locking/rt-mutex.txt
> create mode 100644 Documentation/locking/spinlocks.txt
> create mode 100644 Documentation/locking/ww-mutex-design.txt
> delete mode 100644 Documentation/lockstat.txt
> delete mode 100644 Documentation/mutex-design.txt
> delete mode 100644 Documentation/rt-mutex-design.txt
> delete mode 100644 Documentation/rt-mutex.txt
> delete mode 100644 Documentation/spinlocks.txt
> delete mode 100644 Documentation/ww-mutex-design.txt
--
~Randy
On Wed, 2014-07-30 at 13:41 -0700, Davidlohr Bueso wrote:
> When we fail to acquire the mutex in the fastpath, we end up calling
> __mutex_lock_common(). A *lot* goes on in this function. Move out the
> optimistic spinning code into mutex_optimistic_spin() and simplify
> the former a bit. Furthermore, this is similar to what we have in
> rwsems. No logical changes.
>
> Signed-off-by: Davidlohr Bueso <[email protected]>
Acked-by: Jason Low <[email protected]>
On Wed, Jul 30, 2014 at 02:02:32PM -0700, Randy Dunlap wrote:
> On 07/30/14 13:41, Davidlohr Bueso wrote:
> > Specifically:
> > Documentation/locking/lockdep-design.txt
> > Documentation/locking/lockstat.txt
> > Documentation/locking/mutex-design.txt
> > Documentation/locking/rt-mutex-design.txt
> > Documentation/locking/rt-mutex.txt
> > Documentation/locking/spinlocks.txt
> > Documentation/locking/ww-mutex-design.txt
> >
> > Signed-off-by: Davidlohr Bueso <[email protected]>
>
> Did you read and use Documentation/SubmittingPatches? see:
Not in the last few years, covert changes to that file are likely to be
missed by everybody already doing kernel work for a living.
> Please use "git diff -M --stat --summary" to generate the diffstat:
> the -M enables rename detection, and the summary enables a summary of
> new/deleted or renamed files.
Pretty pointless as I'll pull the lot through quilt, which will destroy
any such niceties.