Unreliable Guide To Locking¶
- Author:
Rusty Russell
Introduction¶
Welcome, to Rusty's Remarkably Unreliable Guide to Kernel Locking issues. This document describes the locking systems in the Linux Kernel in 2.6.
With the wide availability of HyperThreading, and preemption in the Linux Kernel, everyone hacking on the kernel needs to know the fundamentals of concurrency and locking for SMP.
The Problem With Concurrency¶
(Skip this if you know what a Race Condition is).
In a normal program, you can increment a counter like so:
very_important_count++;
This is what they would expect to happen:
Instance 1 |
Instance 2 |
---|---|
read very_important_count (5) |
|
add 1 (6) |
|
write very_important_count (6) |
|
read very_important_count (6) |
|
add 1 (7) |
|
write very_important_count (7) |
This is what might happen:
Instance 1 |
Instance 2 |
---|---|
read very_important_count (5) |
|
read very_important_count (5) |
|
add 1 (6) |
|
add 1 (6) |
|
write very_important_count (6) |
|
write very_important_count (6) |
Race Conditions and Critical Regions¶
This overlap, where the result depends on the relative timing of multiple tasks, is called a race condition. The piece of code containing the concurrency issue is called a critical region. And especially since Linux starting running on SMP machines, they became one of the major issues in kernel design and implementation.
Preemption can have the same effect, even if there is only one CPU: by preempting one task during the critical region, we have exactly the same race condition. In this case the thread which preempts might run the critical region itself.
The solution is to recognize when these simultaneous accesses occur, and use locks to make sure that only one instance can enter the critical region at any time. There are many friendly primitives in the Linux kernel to help you do this. And then there are the unfriendly primitives, but I'll pretend they don't exist.
Locking in the Linux Kernel¶
If I could give you one piece of advice on locking: keep it simple.
Be reluctant to introduce new locks.
Two Main Types of Kernel Locks: Spinlocks and Mutexes¶
There are two main types of kernel locks. The fundamental type is the
spinlock (include/asm/spinlock.h
), which is a very simple
single-holder lock: if you can't get the spinlock, you keep trying
(spinning) until you can. Spinlocks are very small and fast, and can be
used anywhere.
The second type is a mutex (include/linux/mutex.h
): it is like a
spinlock, but you may block holding a mutex. If you can't lock a mutex,
your task will suspend itself, and be woken up when the mutex is
released. This means the CPU can do something else while you are
waiting. There are many cases when you simply can't sleep (see
What Functions Are Safe To Call From Interrupts?),
and so have to use a spinlock instead.
Neither type of lock is recursive: see Deadlock: Simple and Advanced.
Locks and Uniprocessor Kernels¶
For kernels compiled without CONFIG_SMP
, and without
CONFIG_PREEMPT
spinlocks do not exist at all. This is an excellent
design decision: when no-one else can run at the same time, there is no
reason to have a lock.
If the kernel is compiled without CONFIG_SMP
, but CONFIG_PREEMPT
is set, then spinlocks simply disable preemption, which is sufficient to
prevent any races. For most purposes, we can think of preemption as
equivalent to SMP, and not worry about it separately.
You should always test your locking code with CONFIG_SMP
and
CONFIG_PREEMPT
enabled, even if you don't have an SMP test box,
because it will still catch some kinds of locking bugs.
Mutexes still exist, because they are required for synchronization between user contexts, as we will see below.
Locking Only In User Context¶
If you have a data structure which is only ever accessed from user
context, then you can use a simple mutex (include/linux/mutex.h
) to
protect it. This is the most trivial case: you initialize the mutex.
Then you can call mutex_lock_interruptible()
to grab the
mutex, and mutex_unlock()
to release it. There is also a
mutex_lock()
, which should be avoided, because it will
not return if a signal is received.
Example: net/netfilter/nf_sockopt.c
allows registration of new
setsockopt() and getsockopt() calls, with
nf_register_sockopt(). Registration and de-registration
are only done on module load and unload (and boot time, where there is
no concurrency), and the list of registrations is only consulted for an
unknown setsockopt() or getsockopt() system
call. The nf_sockopt_mutex
is perfect to protect this, especially
since the setsockopt and getsockopt calls may well sleep.
Locking Between User Context and Softirqs¶
If a softirq shares data with user context, you have two problems.
Firstly, the current user context can be interrupted by a softirq, and
secondly, the critical region could be entered from another CPU. This is
where spin_lock_bh() (include/linux/spinlock.h
) is
used. It disables softirqs on that CPU, then grabs the lock.
spin_unlock_bh() does the reverse. (The '_bh' suffix is
a historical reference to "Bottom Halves", the old name for software
interrupts. It should really be called spin_lock_softirq()' in a
perfect world).
Note that you can also use spin_lock_irq() or spin_lock_irqsave() here, which stop hardware interrupts as well: see Hard IRQ Context.
This works perfectly for UP as well: the spin lock vanishes, and this
macro simply becomes local_bh_disable()
(include/linux/interrupt.h
), which protects you from the softirq
being run.
Locking Between User Context and Tasklets¶
This is exactly the same as above, because tasklets are actually run from a softirq.
Locking Between User Context and Timers¶
This, too, is exactly the same as above, because timers are actually run from a softirq. From a locking point of view, tasklets and timers are identical.
Locking Between Tasklets/Timers¶
Sometimes a tasklet or timer might want to share data with another tasklet or timer.
The Same Tasklet/Timer¶
Since a tasklet is never run on two CPUs at once, you don't need to worry about your tasklet being reentrant (running twice at once), even on SMP.
Different Tasklets/Timers¶
If another tasklet/timer wants to share data with your tasklet or timer , you will both need to use spin_lock() and spin_unlock() calls. spin_lock_bh() is unnecessary here, as you are already in a tasklet, and none will be run on the same CPU.
Locking Between Softirqs¶
Often a softirq might want to share data with itself or a tasklet/timer.
The Same Softirq¶
The same softirq can run on the other CPUs: you can use a per-CPU array (see Per-CPU Data) for better performance. If you're going so far as to use a softirq, you probably care about scalable performance enough to justify the extra complexity.
You'll need to use spin_lock() and spin_unlock() for shared data.
Different Softirqs¶
You'll need to use spin_lock() and spin_unlock() for shared data, whether it be a timer, tasklet, different softirq or the same or another softirq: any of them could be running on a different CPU.
Hard IRQ Context¶
Hardware interrupts usually communicate with a tasklet or softirq. Frequently this involves putting work in a queue, which the softirq will take out.
Locking Between Hard IRQ and Softirqs/Tasklets¶
If a hardware irq handler shares data with a softirq, you have two concerns. Firstly, the softirq processing can be interrupted by a hardware interrupt, and secondly, the critical region could be entered by a hardware interrupt on another CPU. This is where spin_lock_irq() is used. It is defined to disable interrupts on that cpu, then grab the lock. spin_unlock_irq() does the reverse.
The irq handler does not need to use spin_lock_irq(), because the softirq cannot run while the irq handler is running: it can use spin_lock(), which is slightly faster. The only exception would be if a different hardware irq handler uses the same lock: spin_lock_irq() will stop that from interrupting us.
This works perfectly for UP as well: the spin lock vanishes, and this
macro simply becomes local_irq_disable()
(include/asm/smp.h
), which protects you from the softirq/tasklet/BH
being run.
spin_lock_irqsave() (include/linux/spinlock.h
) is a
variant which saves whether interrupts were on or off in a flags word,
which is passed to spin_unlock_irqrestore(). This means
that the same code can be used inside an hard irq handler (where
interrupts are already off) and in softirqs (where the irq disabling is
required).
Note that softirqs (and hence tasklets and timers) are run on return from hardware interrupts, so spin_lock_irq() also stops these. In that sense, spin_lock_irqsave() is the most general and powerful locking function.
Locking Between Two Hard IRQ Handlers¶
It is rare to have to share data between two IRQ handlers, but if you do, spin_lock_irqsave() should be used: it is architecture-specific whether all interrupts are disabled inside irq handlers themselves.
Cheat Sheet For Locking¶
Pete Zaitcev gives the following summary:
If you are in a process context (any syscall) and want to lock other process out, use a mutex. You can take a mutex and sleep (
copy_from_user()
orkmalloc(x,GFP_KERNEL)
).Otherwise (== data can be touched in an interrupt), use spin_lock_irqsave() and spin_unlock_irqrestore().
Avoid holding spinlock for more than 5 lines of code and across any function call (except accessors like readb()).
Table of Minimum Requirements¶
The following table lists the minimum locking requirements between various contexts. In some cases, the same context can only be running on one CPU at a time, so no locking is required for that context (eg. a particular thread can only run on one CPU at a time, but if it needs shares data with another thread, locking is required).
Remember the advice above: you can always use spin_lock_irqsave(), which is a superset of all other spinlock primitives.
. |
IRQ Handler A |
IRQ Handler B |
Softirq A |
Softirq B |
Tasklet A |
Tasklet B |
Timer A |
Timer B |
User Context A |
User Context B |
---|---|---|---|---|---|---|---|---|---|---|
IRQ Handler A |
None |
|||||||||
IRQ Handler B |
SLIS |
None |
||||||||
Softirq A |
SLI |
SLI |
SL |
|||||||
Softirq B |
SLI |
SLI |
SL |
SL |
||||||
Tasklet A |
SLI |
SLI |
SL |
SL |
None |
|||||
Tasklet B |
SLI |
SLI |
SL |
SL |
SL |
None |
||||
Timer A |
SLI |
SLI |
SL |
SL |
SL |
SL |
None |
|||
Timer B |
SLI |
SLI |
SL |
SL |
SL |
SL |
SL |
None |
||
User Context A |
SLI |
SLI |
SLBH |
SLBH |
SLBH |
SLBH |
SLBH |
SLBH |
None |
|
User Context B |
SLI |
SLI |
SLBH |
SLBH |
SLBH |
SLBH |
SLBH |
SLBH |
MLI |
None |
Table: Table of Locking Requirements
SLIS |
spin_lock_irqsave |
SLI |
spin_lock_irq |
SL |
spin_lock |
SLBH |
spin_lock_bh |
MLI |
mutex_lock_interruptible |
Table: Legend for Locking Requirements Table
The trylock Functions¶
There are functions that try to acquire a lock only once and immediately return a value telling about success or failure to acquire the lock. They can be used if you need no access to the data protected with the lock when some other thread is holding the lock. You should acquire the lock later if you then need access to the data protected with the lock.
spin_trylock() does not spin but returns non-zero if it acquires the spinlock on the first try or 0 if not. This function can be used in all contexts like spin_lock(): you must have disabled the contexts that might interrupt you and acquire the spin lock.
mutex_trylock()
does not suspend your task but returns
non-zero if it could lock the mutex on the first try or 0 if not. This
function cannot be safely used in hardware or software interrupt
contexts despite not sleeping.
Common Examples¶
Let's step through a simple example: a cache of number to name mappings. The cache keeps a count of how often each of the objects is used, and when it gets full, throws out the least used one.
All In User Context¶
For our first example, we assume that all operations are in user context (ie. from system calls), so we can sleep. This means we can use a mutex to protect the cache and all the objects within it. Here's the code:
#include <linux/list.h>
#include <linux/slab.h>
#include <linux/string.h>
#include <linux/mutex.h>
#include <asm/errno.h>
struct object
{
struct list_head list;
int id;
char name[32];
int popularity;
};
/* Protects the cache, cache_num, and the objects within it */
static DEFINE_MUTEX(cache_lock);
static LIST_HEAD(cache);
static unsigned int cache_num = 0;
#define MAX_CACHE_SIZE 10
/* Must be holding cache_lock */
static struct object *__cache_find(int id)
{
struct object *i;
list_for_each_entry(i, &cache, list)
if (i->id == id) {
i->popularity++;
return i;
}
return NULL;
}
/* Must be holding cache_lock */
static void __cache_delete(struct object *obj)
{
BUG_ON(!obj);
list_del(&obj->list);
kfree(obj);
cache_num--;
}
/* Must be holding cache_lock */
static void __cache_add(struct object *obj)
{
list_add(&obj->list, &cache);
if (++cache_num > MAX_CACHE_SIZE) {
struct object *i, *outcast = NULL;
list_for_each_entry(i, &cache, list) {
if (!outcast || i->popularity < outcast->popularity)
outcast = i;
}
__cache_delete(outcast);
}
}
int cache_add(int id, const char *name)
{
struct object *obj;
if ((obj = kmalloc(sizeof(*obj), GFP_KERNEL)) == NULL)
return -ENOMEM;
strscpy(obj->name, name, sizeof(obj->name));
obj->id = id;
obj->popularity = 0;
mutex_lock(&cache_lock);
__cache_add(obj);
mutex_unlock(&cache_lock);
return 0;
}
void cache_delete(int id)
{
mutex_lock(&cache_lock);
__cache_delete(__cache_find(id));
mutex_unlock(&cache_lock);
}
int cache_find(int id, char *name)
{
struct object *obj;
int ret = -ENOENT;
mutex_lock(&cache_lock);
obj = __cache_find(id);
if (obj) {
ret = 0;
strcpy(name, obj->name);
}
mutex_unlock(&cache_lock);
return ret;
}
Note that we always make sure we have the cache_lock when we add, delete, or look up the cache: both the cache infrastructure itself and the contents of the objects are protected by the lock. In this case it's easy, since we copy the data for the user, and never let them access the objects directly.
There is a slight (and common) optimization here: in cache_add() we set up the fields of the object before grabbing the lock. This is safe, as no-one else can access it until we put it in cache.
Accessing From Interrupt Context¶
Now consider the case where cache_find() can be called from interrupt context: either a hardware interrupt or a softirq. An example would be a timer which deletes object from the cache.
The change is shown below, in standard patch format: the -
are lines
which are taken away, and the +
are lines which are added.
--- cache.c.usercontext 2003-12-09 13:58:54.000000000 +1100
+++ cache.c.interrupt 2003-12-09 14:07:49.000000000 +1100
@@ -12,7 +12,7 @@
int popularity;
};
-static DEFINE_MUTEX(cache_lock);
+static DEFINE_SPINLOCK(cache_lock);
static LIST_HEAD(cache);
static unsigned int cache_num = 0;
#define MAX_CACHE_SIZE 10
@@ -55,6 +55,7 @@
int cache_add(int id, const char *name)
{
struct object *obj;
+ unsigned long flags;
if ((obj = kmalloc(sizeof(*obj), GFP_KERNEL)) == NULL)
return -ENOMEM;
@@ -63,30 +64,33 @@
obj->id = id;
obj->popularity = 0;
- mutex_lock(&cache_lock);
+ spin_lock_irqsave(&cache_lock, flags);
__cache_add(obj);
- mutex_unlock(&cache_lock);
+ spin_unlock_irqrestore(&cache_lock, flags);
return 0;
}
void cache_delete(int id)
{
- mutex_lock(&cache_lock);
+ unsigned long flags;
+
+ spin_lock_irqsave(&cache_lock, flags);
__cache_delete(__cache_find(id));
- mutex_unlock(&cache_lock);
+ spin_unlock_irqrestore(&cache_lock, flags);
}
int cache_find(int id, char *name)
{
struct object *obj;
int ret = -ENOENT;
+ unsigned long flags;
- mutex_lock(&cache_lock);
+ spin_lock_irqsave(&cache_lock, flags);
obj = __cache_find(id);
if (obj) {
ret = 0;
strcpy(name, obj->name);
}
- mutex_unlock(&cache_lock);
+ spin_unlock_irqrestore(&cache_lock, flags);
return ret;
}
Note that the spin_lock_irqsave() will turn off interrupts if they are on, otherwise does nothing (if we are already in an interrupt handler), hence these functions are safe to call from any context.
Unfortunately, cache_add() calls kmalloc()
with the GFP_KERNEL
flag, which is only legal in user context. I
have assumed that cache_add() is still only called in
user context, otherwise this should become a parameter to
cache_add().
Exposing Objects Outside This File¶
If our objects contained more information, it might not be sufficient to copy the information in and out: other parts of the code might want to keep pointers to these objects, for example, rather than looking up the id every time. This produces two problems.
The first problem is that we use the cache_lock
to protect objects:
we'd need to make this non-static so the rest of the code can use it.
This makes locking trickier, as it is no longer all in one place.
The second problem is the lifetime problem: if another structure keeps a pointer to an object, it presumably expects that pointer to remain valid. Unfortunately, this is only guaranteed while you hold the lock, otherwise someone might call cache_delete() and even worse, add another object, re-using the same address.
As there is only one lock, you can't hold it forever: no-one else would get any work done.
The solution to this problem is to use a reference count: everyone who has a pointer to the object increases it when they first get the object, and drops the reference count when they're finished with it. Whoever drops it to zero knows it is unused, and can actually delete it.
Here is the code:
--- cache.c.interrupt 2003-12-09 14:25:43.000000000 +1100
+++ cache.c.refcnt 2003-12-09 14:33:05.000000000 +1100
@@ -7,6 +7,7 @@
struct object
{
struct list_head list;
+ unsigned int refcnt;
int id;
char name[32];
int popularity;
@@ -17,6 +18,35 @@
static unsigned int cache_num = 0;
#define MAX_CACHE_SIZE 10
+static void __object_put(struct object *obj)
+{
+ if (--obj->refcnt == 0)
+ kfree(obj);
+}
+
+static void __object_get(struct object *obj)
+{
+ obj->refcnt++;
+}
+
+void object_put(struct object *obj)
+{
+ unsigned long flags;
+
+ spin_lock_irqsave(&cache_lock, flags);
+ __object_put(obj);
+ spin_unlock_irqrestore(&cache_lock, flags);
+}
+
+void object_get(struct object *obj)
+{
+ unsigned long flags;
+
+ spin_lock_irqsave(&cache_lock, flags);
+ __object_get(obj);
+ spin_unlock_irqrestore(&cache_lock, flags);
+}
+
/* Must be holding cache_lock */
static struct object *__cache_find(int id)
{
@@ -35,6 +65,7 @@
{
BUG_ON(!obj);
list_del(&obj->list);
+ __object_put(obj);
cache_num--;
}
@@ -63,6 +94,7 @@
strscpy(obj->name, name, sizeof(obj->name));
obj->id = id;
obj->popularity = 0;
+ obj->refcnt = 1; /* The cache holds a reference */
spin_lock_irqsave(&cache_lock, flags);
__cache_add(obj);
@@ -79,18 +111,15 @@
spin_unlock_irqrestore(&cache_lock, flags);
}
-int cache_find(int id, char *name)
+struct object *cache_find(int id)
{
struct object *obj;
- int ret = -ENOENT;
unsigned long flags;
spin_lock_irqsave(&cache_lock, flags);
obj = __cache_find(id);
- if (obj) {
- ret = 0;
- strcpy(name, obj->name);
- }
+ if (obj)
+ __object_get(obj);
spin_unlock_irqrestore(&cache_lock, flags);
- return ret;
+ return obj;
}
We encapsulate the reference counting in the standard 'get' and 'put' functions. Now we can return the object itself from cache_find() which has the advantage that the user can now sleep holding the object (eg. to copy_to_user() to name to userspace).
The other point to note is that I said a reference should be held for every pointer to the object: thus the reference count is 1 when first inserted into the cache. In some versions the framework does not hold a reference count, but they are more complicated.
Using Atomic Operations For The Reference Count¶
In practice, atomic_t
would usually be used for refcnt. There are a
number of atomic operations defined in include/asm/atomic.h
: these
are guaranteed to be seen atomically from all CPUs in the system, so no
lock is required. In this case, it is simpler than using spinlocks,
although for anything non-trivial using spinlocks is clearer. The
atomic_inc()
and atomic_dec_and_test()
are used instead of the standard increment and decrement operators, and
the lock is no longer used to protect the reference count itself.
--- cache.c.refcnt 2003-12-09 15:00:35.000000000 +1100
+++ cache.c.refcnt-atomic 2003-12-11 15:49:42.000000000 +1100
@@ -7,7 +7,7 @@
struct object
{
struct list_head list;
- unsigned int refcnt;
+ atomic_t refcnt;
int id;
char name[32];
int popularity;
@@ -18,33 +18,15 @@
static unsigned int cache_num = 0;
#define MAX_CACHE_SIZE 10
-static void __object_put(struct object *obj)
-{
- if (--obj->refcnt == 0)
- kfree(obj);
-}
-
-static void __object_get(struct object *obj)
-{
- obj->refcnt++;
-}
-
void object_put(struct object *obj)
{
- unsigned long flags;
-
- spin_lock_irqsave(&cache_lock, flags);
- __object_put(obj);
- spin_unlock_irqrestore(&cache_lock, flags);
+ if (atomic_dec_and_test(&obj->refcnt))
+ kfree(obj);
}
void object_get(struct object *obj)
{
- unsigned long flags;
-
- spin_lock_irqsave(&cache_lock, flags);
- __object_get(obj);
- spin_unlock_irqrestore(&cache_lock, flags);
+ atomic_inc(&obj->refcnt);
}
/* Must be holding cache_lock */
@@ -65,7 +47,7 @@
{
BUG_ON(!obj);
list_del(&obj->list);
- __object_put(obj);
+ object_put(obj);
cache_num--;
}
@@ -94,7 +76,7 @@
strscpy(obj->name, name, sizeof(obj->name));
obj->id = id;
obj->popularity = 0;
- obj->refcnt = 1; /* The cache holds a reference */
+ atomic_set(&obj->refcnt, 1); /* The cache holds a reference */
spin_lock_irqsave(&cache_lock, flags);
__cache_add(obj);
@@ -119,7 +101,7 @@
spin_lock_irqsave(&cache_lock, flags);
obj = __cache_find(id);
if (obj)
- __object_get(obj);
+ object_get(obj);
spin_unlock_irqrestore(&cache_lock, flags);
return obj;
}
Protecting The Objects Themselves¶
In these examples, we assumed that the objects (except the reference counts) never changed once they are created. If we wanted to allow the name to change, there are three possibilities:
You can make
cache_lock
non-static, and tell people to grab that lock before changing the name in any object.You can provide a cache_obj_rename() which grabs this lock and changes the name for the caller, and tell everyone to use that function.
You can make the
cache_lock
protect only the cache itself, and use another lock to protect the name.
Theoretically, you can make the locks as fine-grained as one lock for every field, for every object. In practice, the most common variants are:
One lock which protects the infrastructure (the
cache
list in this example) and all the objects. This is what we have done so far.One lock which protects the infrastructure (including the list pointers inside the objects), and one lock inside the object which protects the rest of that object.
Multiple locks to protect the infrastructure (eg. one lock per hash chain), possibly with a separate per-object lock.
Here is the "lock-per-object" implementation:
--- cache.c.refcnt-atomic 2003-12-11 15:50:54.000000000 +1100
+++ cache.c.perobjectlock 2003-12-11 17:15:03.000000000 +1100
@@ -6,11 +6,17 @@
struct object
{
+ /* These two protected by cache_lock. */
struct list_head list;
+ int popularity;
+
atomic_t refcnt;
+
+ /* Doesn't change once created. */
int id;
+
+ spinlock_t lock; /* Protects the name */
char name[32];
- int popularity;
};
static DEFINE_SPINLOCK(cache_lock);
@@ -77,6 +84,7 @@
obj->id = id;
obj->popularity = 0;
atomic_set(&obj->refcnt, 1); /* The cache holds a reference */
+ spin_lock_init(&obj->lock);
spin_lock_irqsave(&cache_lock, flags);
__cache_add(obj);
Note that I decide that the popularity count should be protected by the
cache_lock
rather than the per-object lock: this is because it (like
the struct list_head
inside the object)
is logically part of the infrastructure. This way, I don't need to grab
the lock of every object in __cache_add() when seeking
the least popular.
I also decided that the id member is unchangeable, so I don't need to grab each object lock in __cache_find() to examine the id: the object lock is only used by a caller who wants to read or write the name field.
Note also that I added a comment describing what data was protected by which locks. This is extremely important, as it describes the runtime behavior of the code, and can be hard to gain from just reading. And as Alan Cox says, “Lock data, not code”.
Common Problems¶
Deadlock: Simple and Advanced¶
There is a coding bug where a piece of code tries to grab a spinlock twice: it will spin forever, waiting for the lock to be released (spinlocks, rwlocks and mutexes are not recursive in Linux). This is trivial to diagnose: not a stay-up-five-nights-talk-to-fluffy-code-bunnies kind of problem.
For a slightly more complex case, imagine you have a region shared by a softirq and user context. If you use a spin_lock() call to protect it, it is possible that the user context will be interrupted by the softirq while it holds the lock, and the softirq will then spin forever trying to get the same lock.
Both of these are called deadlock, and as shown above, it can occur even
with a single CPU (although not on UP compiles, since spinlocks vanish
on kernel compiles with CONFIG_SMP
=n. You'll still get data
corruption in the second example).
This complete lockup is easy to diagnose: on SMP boxes the watchdog
timer or compiling with DEBUG_SPINLOCK
set
(include/linux/spinlock.h
) will show this up immediately when it
happens.
A more complex problem is the so-called 'deadly embrace', involving two or more locks. Say you have a hash table: each entry in the table is a spinlock, and a chain of hashed objects. Inside a softirq handler, you sometimes want to alter an object from one place in the hash to another: you grab the spinlock of the old hash chain and the spinlock of the new hash chain, and delete the object from the old one, and insert it in the new one.
There are two problems here. First, if your code ever tries to move the object to the same chain, it will deadlock with itself as it tries to lock it twice. Secondly, if the same softirq on another CPU is trying to move another object in the reverse direction, the following could happen:
CPU 1 |
CPU 2 |
---|---|
Grab lock A -> OK |
Grab lock B -> OK |
Grab lock B -> spin |
Grab lock A -> spin |
Table: Consequences
The two CPUs will spin forever, waiting for the other to give up their lock. It will look, smell, and feel like a crash.
Preventing Deadlock¶
Textbooks will tell you that if you always lock in the same order, you will never get this kind of deadlock. Practice will tell you that this approach doesn't scale: when I create a new lock, I don't understand enough of the kernel to figure out where in the 5000 lock hierarchy it will fit.
The best locks are encapsulated: they never get exposed in headers, and are never held around calls to non-trivial functions outside the same file. You can read through this code and see that it will never deadlock, because it never tries to grab another lock while it has that one. People using your code don't even need to know you are using a lock.
A classic problem here is when you provide callbacks or hooks: if you call these with the lock held, you risk simple deadlock, or a deadly embrace (who knows what the callback will do?).
Overzealous Prevention Of Deadlocks¶
Deadlocks are problematic, but not as bad as data corruption. Code which grabs a read lock, searches a list, fails to find what it wants, drops the read lock, grabs a write lock and inserts the object has a race condition.
Racing Timers: A Kernel Pastime¶
Timers can produce their own special problems with races. Consider a collection of objects (list, hash, etc) where each object has a timer which is due to destroy it.
If you want to destroy the entire collection (say on module removal), you might do the following:
/* THIS CODE BAD BAD BAD BAD: IF IT WAS ANY WORSE IT WOULD USE
HUNGARIAN NOTATION */
spin_lock_bh(&list_lock);
while (list) {
struct foo *next = list->next;
timer_delete(&list->timer);
kfree(list);
list = next;
}
spin_unlock_bh(&list_lock);
Sooner or later, this will crash on SMP, because a timer can have just gone off before the spin_lock_bh(), and it will only get the lock after we spin_unlock_bh(), and then try to free the element (which has already been freed!).
This can be avoided by checking the result of
timer_delete()
: if it returns 1, the timer has been deleted.
If 0, it means (in this case) that it is currently running, so we can
do:
retry:
spin_lock_bh(&list_lock);
while (list) {
struct foo *next = list->next;
if (!timer_delete(&list->timer)) {
/* Give timer a chance to delete this */
spin_unlock_bh(&list_lock);
goto retry;
}
kfree(list);
list = next;
}
spin_unlock_bh(&list_lock);
Another common problem is deleting timers which restart themselves (by
calling add_timer()
at the end of their timer function).
Because this is a fairly common case which is prone to races, you should
use timer_delete_sync()
(include/linux/timer.h
) to handle this case.
Before freeing a timer, timer_shutdown()
or timer_shutdown_sync()
should be
called which will keep it from being rearmed. Any subsequent attempt to
rearm the timer will be silently ignored by the core code.
Locking Speed¶
There are three main things to worry about when considering speed of some code which does locking. First is concurrency: how many things are going to be waiting while someone else is holding a lock. Second is the time taken to actually acquire and release an uncontended lock. Third is using fewer, or smarter locks. I'm assuming that the lock is used fairly often: otherwise, you wouldn't be concerned about efficiency.
Concurrency depends on how long the lock is usually held: you should hold the lock for as long as needed, but no longer. In the cache example, we always create the object without the lock held, and then grab the lock only when we are ready to insert it in the list.
Acquisition times depend on how much damage the lock operations do to the pipeline (pipeline stalls) and how likely it is that this CPU was the last one to grab the lock (ie. is the lock cache-hot for this CPU): on a machine with more CPUs, this likelihood drops fast. Consider a 700MHz Intel Pentium III: an instruction takes about 0.7ns, an atomic increment takes about 58ns, a lock which is cache-hot on this CPU takes 160ns, and a cacheline transfer from another CPU takes an additional 170 to 360ns. (These figures from Paul McKenney's Linux Journal RCU article).
These two aims conflict: holding a lock for a short time might be done by splitting locks into parts (such as in our final per-object-lock example), but this increases the number of lock acquisitions, and the results are often slower than having a single lock. This is another reason to advocate locking simplicity.
The third concern is addressed below: there are some methods to reduce the amount of locking which needs to be done.
Read/Write Lock Variants¶
Both spinlocks and mutexes have read/write variants: rwlock_t
and
struct rw_semaphore
. These divide
users into two classes: the readers and the writers. If you are only
reading the data, you can get a read lock, but to write to the data you
need the write lock. Many people can hold a read lock, but a writer must
be sole holder.
If your code divides neatly along reader/writer lines (as our cache code
does), and the lock is held by readers for significant lengths of time,
using these locks can help. They are slightly slower than the normal
locks though, so in practice rwlock_t
is not usually worthwhile.
Avoiding Locks: Read Copy Update¶
There is a special method of read/write locking called Read Copy Update. Using RCU, the readers can avoid taking a lock altogether: as we expect our cache to be read more often than updated (otherwise the cache is a waste of time), it is a candidate for this optimization.
How do we get rid of read locks? Getting rid of read locks means that
writers may be changing the list underneath the readers. That is
actually quite simple: we can read a linked list while an element is
being added if the writer adds the element very carefully. For example,
adding new
to a single linked list called list
:
new->next = list->next;
wmb();
list->next = new;
The wmb() is a write memory barrier. It ensures that the
first operation (setting the new element's next
pointer) is complete
and will be seen by all CPUs, before the second operation is (putting
the new element into the list). This is important, since modern
compilers and modern CPUs can both reorder instructions unless told
otherwise: we want a reader to either not see the new element at all, or
see the new element with the next
pointer correctly pointing at the
rest of the list.
Fortunately, there is a function to do this for standard
struct list_head
lists:
list_add_rcu()
(include/linux/list.h
).
Removing an element from the list is even simpler: we replace the pointer to the old element with a pointer to its successor, and readers will either see it, or skip over it.
list->next = old->next;
There is list_del_rcu()
(include/linux/list.h
) which
does this (the normal version poisons the old object, which we don't
want).
The reader must also be careful: some CPUs can look through the next
pointer to start reading the contents of the next element early, but
don't realize that the pre-fetched contents is wrong when the next
pointer changes underneath them. Once again, there is a
list_for_each_entry_rcu()
(include/linux/list.h
)
to help you. Of course, writers can just use
list_for_each_entry()
, since there cannot be two
simultaneous writers.
Our final dilemma is this: when can we actually destroy the removed
element? Remember, a reader might be stepping through this element in
the list right now: if we free this element and the next
pointer
changes, the reader will jump off into garbage and crash. We need to
wait until we know that all the readers who were traversing the list
when we deleted the element are finished. We use
call_rcu()
to register a callback which will actually
destroy the object once all pre-existing readers are finished.
Alternatively, synchronize_rcu()
may be used to block
until all pre-existing are finished.
But how does Read Copy Update know when the readers are finished? The
method is this: firstly, the readers always traverse the list inside
rcu_read_lock()
/rcu_read_unlock()
pairs:
these simply disable preemption so the reader won't go to sleep while
reading the list.
RCU then waits until every other CPU has slept at least once: since readers cannot sleep, we know that any readers which were traversing the list during the deletion are finished, and the callback is triggered. The real Read Copy Update code is a little more optimized than this, but this is the fundamental idea.
--- cache.c.perobjectlock 2003-12-11 17:15:03.000000000 +1100
+++ cache.c.rcupdate 2003-12-11 17:55:14.000000000 +1100
@@ -1,15 +1,18 @@
#include <linux/list.h>
#include <linux/slab.h>
#include <linux/string.h>
+#include <linux/rcupdate.h>
#include <linux/mutex.h>
#include <asm/errno.h>
struct object
{
- /* These two protected by cache_lock. */
+ /* This is protected by RCU */
struct list_head list;
int popularity;
+ struct rcu_head rcu;
+
atomic_t refcnt;
/* Doesn't change once created. */
@@ -40,7 +43,7 @@
{
struct object *i;
- list_for_each_entry(i, &cache, list) {
+ list_for_each_entry_rcu(i, &cache, list) {
if (i->id == id) {
i->popularity++;
return i;
@@ -49,19 +52,25 @@
return NULL;
}
+/* Final discard done once we know no readers are looking. */
+static void cache_delete_rcu(void *arg)
+{
+ object_put(arg);
+}
+
/* Must be holding cache_lock */
static void __cache_delete(struct object *obj)
{
BUG_ON(!obj);
- list_del(&obj->list);
- object_put(obj);
+ list_del_rcu(&obj->list);
cache_num--;
+ call_rcu(&obj->rcu, cache_delete_rcu);
}
/* Must be holding cache_lock */
static void __cache_add(struct object *obj)
{
- list_add(&obj->list, &cache);
+ list_add_rcu(&obj->list, &cache);
if (++cache_num > MAX_CACHE_SIZE) {
struct object *i, *outcast = NULL;
list_for_each_entry(i, &cache, list) {
@@ -104,12 +114,11 @@
struct object *cache_find(int id)
{
struct object *obj;
- unsigned long flags;
- spin_lock_irqsave(&cache_lock, flags);
+ rcu_read_lock();
obj = __cache_find(id);
if (obj)
object_get(obj);
- spin_unlock_irqrestore(&cache_lock, flags);
+ rcu_read_unlock();
return obj;
}
Note that the reader will alter the popularity member in
__cache_find(), and now it doesn't hold a lock. One
solution would be to make it an atomic_t
, but for this usage, we
don't really care about races: an approximate result is good enough, so
I didn't change it.
The result is that cache_find() requires no synchronization with any other functions, so is almost as fast on SMP as it would be on UP.
There is a further optimization possible here: remember our original cache code, where there were no reference counts and the caller simply held the lock whenever using the object? This is still possible: if you hold the lock, no one can delete the object, so you don't need to get and put the reference count.
Now, because the 'read lock' in RCU is simply disabling preemption, a caller which always has preemption disabled between calling cache_find() and object_put() does not need to actually get and put the reference count: we could expose __cache_find() by making it non-static, and such callers could simply call that.
The benefit here is that the reference count is not written to: the object is not altered in any way, which is much faster on SMP machines due to caching.
Per-CPU Data¶
Another technique for avoiding locking which is used fairly widely is to duplicate information for each CPU. For example, if you wanted to keep a count of a common condition, you could use a spin lock and a single counter. Nice and simple.
If that was too slow (it's usually not, but if you've got a really big
machine to test on and can show that it is), you could instead use a
counter for each CPU, then none of them need an exclusive lock. See
DEFINE_PER_CPU(), get_cpu_var() and
put_cpu_var() (include/linux/percpu.h
).
Of particular use for simple per-cpu counters is the local_t
type,
and the cpu_local_inc() and related functions, which are
more efficient than simple code on some architectures
(include/asm/local.h
).
Note that there is no simple, reliable way of getting an exact value of such a counter, without introducing more locks. This is not a problem for some uses.
Data Which Mostly Used By An IRQ Handler¶
If data is always accessed from within the same IRQ handler, you don't need a lock at all: the kernel already guarantees that the irq handler will not run simultaneously on multiple CPUs.
Manfred Spraul points out that you can still do this, even if the data is very occasionally accessed in user context or softirqs/tasklets. The irq handler doesn't use a lock, and all other accesses are done as so:
mutex_lock(&lock);
disable_irq(irq);
...
enable_irq(irq);
mutex_unlock(&lock);
The disable_irq()
prevents the irq handler from running
(and waits for it to finish if it's currently running on other CPUs).
The spinlock prevents any other accesses happening at the same time.
Naturally, this is slower than just a spin_lock_irq()
call, so it only makes sense if this type of access happens extremely
rarely.
What Functions Are Safe To Call From Interrupts?¶
Many functions in the kernel sleep (ie. call schedule()) directly or indirectly: you can never call them while holding a spinlock, or with preemption disabled. This also means you need to be in user context: calling them from an interrupt is illegal.
Some Functions Which Sleep¶
The most common ones are listed below, but you usually have to read the code to find out if other calls are safe. If everyone else who calls it can sleep, you probably need to be able to sleep, too. In particular, registration and deregistration functions usually expect to be called from user context, and can sleep.
Accesses to userspace:
copy_from_user()
copy_to_user()
kmalloc(GP_KERNEL) <kmalloc>`
mutex_lock_interruptible()
andmutex_lock()
There is a
mutex_trylock()
which does not sleep. Still, it must not be used inside interrupt context since its implementation is not safe for that.mutex_unlock()
will also never sleep. It cannot be used in interrupt context either since a mutex must be released by the same task that acquired it.
Some Functions Which Don't Sleep¶
Some functions are safe to call from any context, or holding almost any lock.
Mutex API reference¶
-
mutex_init¶
mutex_init (mutex)
initialize the mutex
Parameters
mutex
the mutex to be initialized
Description
Initialize the mutex to unlocked state.
It is not allowed to initialize an already locked mutex.
-
bool mutex_is_locked(struct mutex *lock)¶
is the mutex locked
Parameters
struct mutex *lock
the mutex to be queried
Description
Returns true if the mutex is locked, false if unlocked.
-
void mutex_lock(struct mutex *lock)¶
acquire the mutex
Parameters
struct mutex *lock
the mutex to be acquired
Description
Lock the mutex exclusively for this task. If the mutex is not available right now, it will sleep until it can get it.
The mutex must later on be released by the same task that
acquired it. Recursive locking is not allowed. The task
may not exit without first unlocking the mutex. Also, kernel
memory where the mutex resides must not be freed with
the mutex still locked. The mutex must first be initialized
(or statically defined) before it can be locked. memset()
-ing
the mutex to 0 is not allowed.
(The CONFIG_DEBUG_MUTEXES .config option turns on debugging checks that will enforce the restrictions and will also do deadlock debugging)
This function is similar to (but not equivalent to) down().
-
void mutex_unlock(struct mutex *lock)¶
release the mutex
Parameters
struct mutex *lock
the mutex to be released
Description
Unlock a mutex that has been locked by this task previously.
This function must not be used in interrupt context. Unlocking of a not locked mutex is not allowed.
This function is similar to (but not equivalent to) up().
-
void ww_mutex_unlock(struct ww_mutex *lock)¶
release the w/w mutex
Parameters
struct ww_mutex *lock
the mutex to be released
Description
Unlock a mutex that has been locked by this task previously with any of the ww_mutex_lock* functions (with or without an acquire context). It is forbidden to release the locks after releasing the acquire context.
This function must not be used in interrupt context. Unlocking of a unlocked mutex is not allowed.
-
int ww_mutex_trylock(struct ww_mutex *ww, struct ww_acquire_ctx *ww_ctx)¶
tries to acquire the w/w mutex with optional acquire context
Parameters
struct ww_mutex *ww
mutex to lock
struct ww_acquire_ctx *ww_ctx
optional w/w acquire context
Description
Trylocks a mutex with the optional acquire context; no deadlock detection is possible. Returns 1 if the mutex has been acquired successfully, 0 otherwise.
Unlike ww_mutex_lock, no deadlock handling is performed. However, if a ctx is specified, -EALREADY handling may happen in calls to ww_mutex_trylock.
A mutex acquired with this function must be released with ww_mutex_unlock.
-
int mutex_lock_interruptible(struct mutex *lock)¶
Acquire the mutex, interruptible by signals.
Parameters
struct mutex *lock
The mutex to be acquired.
Description
Lock the mutex like mutex_lock()
. If a signal is delivered while the
process is sleeping, this function will return without acquiring the
mutex.
Context
Process context.
Return
0 if the lock was successfully acquired or -EINTR
if a
signal arrived.
-
int mutex_lock_killable(struct mutex *lock)¶
Acquire the mutex, interruptible by fatal signals.
Parameters
struct mutex *lock
The mutex to be acquired.
Description
Lock the mutex like mutex_lock()
. If a signal which will be fatal to
the current process is delivered while the process is sleeping, this
function will return without acquiring the mutex.
Context
Process context.
Return
0 if the lock was successfully acquired or -EINTR
if a
fatal signal arrived.
-
void mutex_lock_io(struct mutex *lock)¶
Acquire the mutex and mark the process as waiting for I/O
Parameters
struct mutex *lock
The mutex to be acquired.
Description
Lock the mutex like mutex_lock()
. While the task is waiting for this
mutex, it will be accounted as being in the IO wait state by the
scheduler.
Context
Process context.
-
int mutex_trylock(struct mutex *lock)¶
try to acquire the mutex, without waiting
Parameters
struct mutex *lock
the mutex to be acquired
Description
Try to acquire the mutex atomically. Returns 1 if the mutex has been acquired successfully, and 0 on contention.
This function must not be used in interrupt context. The mutex must be released by the same task that acquired it.
NOTE
this function follows the spin_trylock() convention, so it is negated from the down_trylock() return values! Be careful about this when converting semaphore users to mutexes.
-
int atomic_dec_and_mutex_lock(atomic_t *cnt, struct mutex *lock)¶
return holding mutex if we dec to 0
Parameters
atomic_t *cnt
the atomic which we are to dec
struct mutex *lock
the mutex to return holding if we dec to 0
Description
return true and hold lock if we dec to 0, return false otherwise
Futex API reference¶
-
struct futex_hash_bucket *futex_hash(union futex_key *key)¶
Return the hash bucket in the global hash
Parameters
union futex_key *key
Pointer to the futex key for which the hash is calculated
Description
We hash on the keys returned from get_futex_key (see below) and return the corresponding hash bucket in the global hash.
-
struct hrtimer_sleeper *futex_setup_timer(ktime_t *time, struct hrtimer_sleeper *timeout, int flags, u64 range_ns)¶
set up the sleeping hrtimer.
Parameters
ktime_t *time
ptr to the given timeout value
struct hrtimer_sleeper *timeout
the hrtimer_sleeper structure to be set up
int flags
futex flags
u64 range_ns
optional range in ns
Return
- Initialized hrtimer_sleeper structure or NULL if no timeout
value given
-
int get_futex_key(u32 __user *uaddr, bool fshared, union futex_key *key, enum futex_access rw)¶
Get parameters which are the keys for a futex
Parameters
u32 __user *uaddr
virtual address of the futex
bool fshared
false for a PROCESS_PRIVATE futex, true for PROCESS_SHARED
union futex_key *key
address where result is stored.
enum futex_access rw
mapping needs to be read/write (values: FUTEX_READ, FUTEX_WRITE)
Return
a negative error code or 0
Description
The key words are stored in key on success.
For shared mappings (when fshared), the key is:
( inode->i_sequence, page->index, offset_within_page )
[ also see get_inode_sequence_number() ]
For private mappings (or when !fshared), the key is:
( current->mm, address, 0 )
This allows (cross process, where applicable) identification of the futex without keeping the page pinned for the duration of the FUTEX_WAIT.
lock_page()
might sleep, the caller should not hold a spinlock.
-
int fault_in_user_writeable(u32 __user *uaddr)¶
Fault in user address and verify RW access
Parameters
u32 __user *uaddr
pointer to faulting user space address
Description
Slow path to fixup the fault we just took in the atomic write access to uaddr.
We have no generic implementation of a non-destructive write to the user address. We know that we faulted in the atomic pagefault disabled section so we can as well avoid the #PF overhead by calling get_user_pages() right away.
-
struct futex_q *futex_top_waiter(struct futex_hash_bucket *hb, union futex_key *key)¶
Return the highest priority waiter on a futex
Parameters
struct futex_hash_bucket *hb
the hash bucket the futex_q's reside in
union futex_key *key
the futex key (to distinguish it from other futex futex_q's)
Description
Must be called with the hb lock held.
-
void wait_for_owner_exiting(int ret, struct task_struct *exiting)¶
Block until the owner has exited
Parameters
int ret
owner's current futex lock status
struct task_struct *exiting
Pointer to the exiting task
Description
Caller must hold a refcount on exiting.
Parameters
struct futex_q *q
The futex_q to unqueue
Description
The q->lock_ptr must not be NULL and must be held by the caller.
Parameters
struct futex_q *q
The futex_q to unqueue
Description
The q->lock_ptr must not be held by the caller. A call to futex_unqueue()
must
be paired with exactly one earlier call to futex_queue()
.
Return
1 - if the futex_q was still queued (and we removed unqueued it);
0 - if the futex_q was already removed by the waking thread
-
void futex_exit_recursive(struct task_struct *tsk)¶
Set the tasks futex state to FUTEX_STATE_DEAD
Parameters
struct task_struct *tsk
task to set the state on
Description
Set the futex exit state of the task lockless. The futex waiter code observes that state when a task is exiting and loops until the task has actually finished the futex cleanup. The worst case for this is that the waiter runs through the wait loop until the state becomes visible.
This is called from the recursive fault handling path in make_task_dead().
This is best effort. Either the futex exit code has run already or not. If the OWNER_DIED bit has been set on the futex then the waiter can take it over. If not, the problem is pushed back to user space. If the futex exit code did not run yet, then an already queued waiter might block forever, but there is nothing which can be done about that.
-
struct futex_q¶
The hashed futex queue entry, one per waiting task
Definition:
struct futex_q {
struct plist_node list;
struct task_struct *task;
spinlock_t *lock_ptr;
union futex_key key;
struct futex_pi_state *pi_state;
struct rt_mutex_waiter *rt_waiter;
union futex_key *requeue_pi_key;
u32 bitset;
atomic_t requeue_state;
#ifdef CONFIG_PREEMPT_RT;
struct rcuwait requeue_wait;
#endif;
};
Members
list
priority-sorted list of tasks waiting on this futex
task
the task waiting on the futex
lock_ptr
the hash bucket lock
key
the key the futex is hashed on
pi_state
optional priority inheritance state
rt_waiter
rt_waiter storage for use with requeue_pi
requeue_pi_key
the requeue_pi target futex key
bitset
bitset for the optional bitmasked wakeup
requeue_state
State field for futex_requeue_pi()
requeue_wait
RCU wait for futex_requeue_pi() (RT only)
Description
We use this hashed waitqueue, instead of a normal wait_queue_entry_t, so we can wake only the relevant ones (hashed queues may be shared).
A futex_q has a woken state, just like tasks have TASK_RUNNING.
It is considered woken when plist_node_empty(q->list
) || q->lock_ptr == 0.
The order of wakeup is always to make the first condition true, then
the second.
PI futexes are typically woken before they are removed from the hash list via the rt_mutex code. See futex_unqueue_pi().
-
int futex_match(union futex_key *key1, union futex_key *key2)¶
Check whether two futex keys are equal
Parameters
union futex_key *key1
Pointer to key1
union futex_key *key2
Pointer to key2
Description
Return 1 if two futex_keys are equal, 0 otherwise.
-
void futex_queue(struct futex_q *q, struct futex_hash_bucket *hb)¶
Enqueue the futex_q on the futex_hash_bucket
Parameters
struct futex_q *q
The futex_q to enqueue
struct futex_hash_bucket *hb
The destination hash bucket
Description
The hb->lock must be held by the caller, and is released here. A call to
futex_queue()
is typically paired with exactly one call to futex_unqueue()
. The
exceptions involve the PI related operations, which may use futex_unqueue_pi()
or nothing if the unqueue is done as part of the wake process and the unqueue
state is implicit in the state of woken task (see futex_wait_requeue_pi()
for
an example).
-
struct futex_vector¶
Auxiliary struct for futex_waitv()
Definition:
struct futex_vector {
struct futex_waitv w;
struct futex_q q;
};
Members
w
Userspace provided data
q
Kernel side data
Description
Struct used to build an array with all data need for futex_waitv()
-
int futex_lock_pi_atomic(u32 __user *uaddr, struct futex_hash_bucket *hb, union futex_key *key, struct futex_pi_state **ps, struct task_struct *task, struct task_struct **exiting, int set_waiters)¶
Atomic work required to acquire a pi aware futex
Parameters
u32 __user *uaddr
the pi futex user address
struct futex_hash_bucket *hb
the pi futex hash bucket
union futex_key *key
the futex key associated with uaddr and hb
struct futex_pi_state **ps
the pi_state pointer where we store the result of the lookup
struct task_struct *task
the task to perform the atomic lock work for. This will be "current" except in the case of requeue pi.
struct task_struct **exiting
Pointer to store the task pointer of the owner task which is in the middle of exiting
int set_waiters
force setting the FUTEX_WAITERS bit (1) or not (0)
Return
0 - ready to wait;
1 - acquired the lock;
<0 - error
Description
The hb->lock must be held by the caller.
exiting is only set when the return value is -EBUSY. If so, this holds a refcount on the exiting task on return and the caller needs to drop it after waiting for the exit to complete.
-
int fixup_pi_owner(u32 __user *uaddr, struct futex_q *q, int locked)¶
Post lock pi_state and corner case management
Parameters
u32 __user *uaddr
user address of the futex
struct futex_q *q
futex_q (contains pi_state and access to the rt_mutex)
int locked
if the attempt to take the rt_mutex succeeded (1) or not (0)
Description
After attempting to lock an rt_mutex, this function is called to cleanup the pi_state owner as well as handle race conditions that may allow us to acquire the lock. Must be called with the hb lock held.
Return
1 - success, lock taken;
0 - success, lock not taken;
<0 - on error (-EFAULT)
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void requeue_futex(struct futex_q *q, struct futex_hash_bucket *hb1, struct futex_hash_bucket *hb2, union futex_key *key2)¶
Requeue a futex_q from one hb to another
Parameters
struct futex_q *q
the futex_q to requeue
struct futex_hash_bucket *hb1
the source hash_bucket
struct futex_hash_bucket *hb2
the target hash_bucket
union futex_key *key2
the new key for the requeued futex_q
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void requeue_pi_wake_futex(struct futex_q *q, union futex_key *key, struct futex_hash_bucket *hb)¶
Wake a task that acquired the lock during requeue
Parameters
struct futex_q *q
the futex_q
union futex_key *key
the key of the requeue target futex
struct futex_hash_bucket *hb
the hash_bucket of the requeue target futex
Description
During futex_requeue, with requeue_pi=1, it is possible to acquire the target futex if it is uncontended or via a lock steal.
Set q::key to the requeue target futex key so the waiter can detect the wakeup on the right futex.
Dequeue q from the hash bucket.
Set q::rt_waiter to NULL so the woken up task can detect atomic lock acquisition.
Set the q->lock_ptr to the requeue target hb->lock for the case that the waiter has to fixup the pi state.
Complete the requeue state so the waiter can make progress. After this point the waiter task can return from the syscall immediately in case that the pi state does not have to be fixed up.
Wake the waiter task.
Must be called with both q->lock_ptr and hb->lock held.
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int futex_proxy_trylock_atomic(u32 __user *pifutex, struct futex_hash_bucket *hb1, struct futex_hash_bucket *hb2, union futex_key *key1, union futex_key *key2, struct futex_pi_state **ps, struct task_struct **exiting, int set_waiters)¶
Attempt an atomic lock for the top waiter
Parameters
u32 __user *pifutex
the user address of the to futex
struct futex_hash_bucket *hb1
the from futex hash bucket, must be locked by the caller
struct futex_hash_bucket *hb2
the to futex hash bucket, must be locked by the caller
union futex_key *key1
the from futex key
union futex_key *key2
the to futex key
struct futex_pi_state **ps
address to store the pi_state pointer
struct task_struct **exiting
Pointer to store the task pointer of the owner task which is in the middle of exiting
int set_waiters
force setting the FUTEX_WAITERS bit (1) or not (0)
Description
Try and get the lock on behalf of the top waiter if we can do it atomically.
Wake the top waiter if we succeed. If the caller specified set_waiters,
then direct futex_lock_pi_atomic()
to force setting the FUTEX_WAITERS bit.
hb1 and hb2 must be held by the caller.
exiting is only set when the return value is -EBUSY. If so, this holds a refcount on the exiting task on return and the caller needs to drop it after waiting for the exit to complete.
Return
0 - failed to acquire the lock atomically;
>0 - acquired the lock, return value is vpid of the top_waiter
<0 - error
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int futex_requeue(u32 __user *uaddr1, unsigned int flags, u32 __user *uaddr2, int nr_wake, int nr_requeue, u32 *cmpval, int requeue_pi)¶
Requeue waiters from uaddr1 to uaddr2
Parameters
u32 __user *uaddr1
source futex user address
unsigned int flags
futex flags (FLAGS_SHARED, etc.)
u32 __user *uaddr2
target futex user address
int nr_wake
number of waiters to wake (must be 1 for requeue_pi)
int nr_requeue
number of waiters to requeue (0-INT_MAX)
u32 *cmpval
uaddr1 expected value (or
NULL
)int requeue_pi
if we are attempting to requeue from a non-pi futex to a pi futex (pi to pi requeue is not supported)
Description
Requeue waiters on uaddr1 to uaddr2. In the requeue_pi case, try to acquire uaddr2 atomically on behalf of the top waiter.
Return
>=0 - on success, the number of tasks requeued or woken;
<0 - on error
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int handle_early_requeue_pi_wakeup(struct futex_hash_bucket *hb, struct futex_q *q, struct hrtimer_sleeper *timeout)¶
Handle early wakeup on the initial futex
Parameters
struct futex_hash_bucket *hb
the hash_bucket futex_q was original enqueued on
struct futex_q *q
the futex_q woken while waiting to be requeued
struct hrtimer_sleeper *timeout
the timeout associated with the wait (NULL if none)
Description
Determine the cause for the early wakeup.
Return
-EWOULDBLOCK or -ETIMEDOUT or -ERESTARTNOINTR
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int futex_wait_requeue_pi(u32 __user *uaddr, unsigned int flags, u32 val, ktime_t *abs_time, u32 bitset, u32 __user *uaddr2)¶
Wait on uaddr and take uaddr2
Parameters
u32 __user *uaddr
the futex we initially wait on (non-pi)
unsigned int flags
futex flags (FLAGS_SHARED, FLAGS_CLOCKRT, etc.), they must be the same type, no requeueing from private to shared, etc.
u32 val
the expected value of uaddr
ktime_t *abs_time
absolute timeout
u32 bitset
32 bit wakeup bitset set by userspace, defaults to all
u32 __user *uaddr2
the pi futex we will take prior to returning to user-space
Description
The caller will wait on uaddr and will be requeued by futex_requeue()
to
uaddr2 which must be PI aware and unique from uaddr. Normal wakeup will wake
on uaddr2 and complete the acquisition of the rt_mutex prior to returning to
userspace. This ensures the rt_mutex maintains an owner when it has waiters;
without one, the pi logic would not know which task to boost/deboost, if
there was a need to.
We call schedule in futex_wait_queue()
when we enqueue and return there
via the following--
1) wakeup on uaddr2 after an atomic lock acquisition by futex_requeue()
2) wakeup on uaddr2 after a requeue
3) signal
4) timeout
If 3, cleanup and return -ERESTARTNOINTR.
If 2, we may then block on trying to take the rt_mutex and return via: 5) successful lock 6) signal 7) timeout 8) other lock acquisition failure
If 6, return -EWOULDBLOCK (restarting the syscall would do the same).
If 4 or 7, we cleanup and return with -ETIMEDOUT.
Return
0 - On success;
<0 - On error
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void futex_wait_queue(struct futex_hash_bucket *hb, struct futex_q *q, struct hrtimer_sleeper *timeout)¶
futex_queue()
and wait for wakeup, timeout, or signal
Parameters
struct futex_hash_bucket *hb
the futex hash bucket, must be locked by the caller
struct futex_q *q
the futex_q to queue up on
struct hrtimer_sleeper *timeout
the prepared hrtimer_sleeper, or null for no timeout
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int unqueue_multiple(struct futex_vector *v, int count)¶
Remove various futexes from their hash bucket
Parameters
struct futex_vector *v
The list of futexes to unqueue
int count
Number of futexes in the list
Description
Helper to unqueue a list of futexes. This can't fail.
Return
>=0 - Index of the last futex that was awoken;
- -1
No futex was awoken
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int futex_wait_multiple_setup(struct futex_vector *vs, int count, int *woken)¶
Prepare to wait and enqueue multiple futexes
Parameters
struct futex_vector *vs
The futex list to wait on
int count
The size of the list
int *woken
Index of the last woken futex, if any. Used to notify the caller that it can return this index to userspace (return parameter)
Description
Prepare multiple futexes in a single step and enqueue them. This may fail if the futex list is invalid or if any futex was already awoken. On success the task is ready to interruptible sleep.
Return
1 - One of the futexes was woken by another thread
0 - Success
<0 - -EFAULT, -EWOULDBLOCK or -EINVAL
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void futex_sleep_multiple(struct futex_vector *vs, unsigned int count, struct hrtimer_sleeper *to)¶
Check sleeping conditions and sleep
Parameters
struct futex_vector *vs
List of futexes to wait for
unsigned int count
Length of vs
struct hrtimer_sleeper *to
Timeout
Description
Sleep if and only if the timeout hasn't expired and no futex on the list has been woken up.
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int futex_wait_multiple(struct futex_vector *vs, unsigned int count, struct hrtimer_sleeper *to)¶
Prepare to wait on and enqueue several futexes
Parameters
struct futex_vector *vs
The list of futexes to wait on
unsigned int count
The number of objects
struct hrtimer_sleeper *to
Timeout before giving up and returning to userspace
Description
Entry point for the FUTEX_WAIT_MULTIPLE futex operation, this function sleeps on a group of futexes and returns on the first futex that is wake, or after the timeout has elapsed.
Return
>=0 - Hint to the futex that was awoken
<0 - On error
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int futex_wait_setup(u32 __user *uaddr, u32 val, unsigned int flags, struct futex_q *q, struct futex_hash_bucket **hb)¶
Prepare to wait on a futex
Parameters
u32 __user *uaddr
the futex userspace address
u32 val
the expected value
unsigned int flags
futex flags (FLAGS_SHARED, etc.)
struct futex_q *q
the associated futex_q
struct futex_hash_bucket **hb
storage for hash_bucket pointer to be returned to caller
Description
Setup the futex_q and locate the hash_bucket. Get the futex value and compare it with the expected value. Handle atomic faults internally. Return with the hb lock held on success, and unlocked on failure.
Return
0 - uaddr contains val and hb has been locked;
<1 - -EFAULT or -EWOULDBLOCK (uaddr does not contain val) and hb is unlocked
Further reading¶
Documentation/locking/spinlocks.rst
: Linus Torvalds' spinlocking tutorial in the kernel sources.Unix Systems for Modern Architectures: Symmetric Multiprocessing and Caching for Kernel Programmers:
Curt Schimmel's very good introduction to kernel level locking (not written for Linux, but nearly everything applies). The book is expensive, but really worth every penny to understand SMP locking. [ISBN: 0201633388]
Thanks¶
Thanks to Telsa Gwynne for DocBooking, neatening and adding style.
Thanks to Martin Pool, Philipp Rumpf, Stephen Rothwell, Paul Mackerras, Ruedi Aschwanden, Alan Cox, Manfred Spraul, Tim Waugh, Pete Zaitcev, James Morris, Robert Love, Paul McKenney, John Ashby for proofreading, correcting, flaming, commenting.
Thanks to the cabal for having no influence on this document.
Glossary¶
- preemption
Prior to 2.5, or when
CONFIG_PREEMPT
is unset, processes in user context inside the kernel would not preempt each other (ie. you had that CPU until you gave it up, except for interrupts). With the addition ofCONFIG_PREEMPT
in 2.5.4, this changed: when in user context, higher priority tasks can "cut in": spinlocks were changed to disable preemption, even on UP.- bh
Bottom Half: for historical reasons, functions with '_bh' in them often now refer to any software interrupt, e.g. spin_lock_bh() blocks any software interrupt on the current CPU. Bottom halves are deprecated, and will eventually be replaced by tasklets. Only one bottom half will be running at any time.
- Hardware Interrupt / Hardware IRQ
Hardware interrupt request. in_hardirq() returns true in a hardware interrupt handler.
- Interrupt Context
Not user context: processing a hardware irq or software irq. Indicated by the in_interrupt() macro returning true.
- SMP
Symmetric Multi-Processor: kernels compiled for multiple-CPU machines. (
CONFIG_SMP=y
).- Software Interrupt / softirq
Software interrupt handler. in_hardirq() returns false; in_softirq() returns true. Tasklets and softirqs both fall into the category of 'software interrupts'.
Strictly speaking a softirq is one of up to 32 enumerated software interrupts which can run on multiple CPUs at once. Sometimes used to refer to tasklets as well (ie. all software interrupts).
- tasklet
A dynamically-registrable software interrupt, which is guaranteed to only run on one CPU at a time.
- timer
A dynamically-registrable software interrupt, which is run at (or close to) a given time. When running, it is just like a tasklet (in fact, they are called from the
TIMER_SOFTIRQ
).- UP
Uni-Processor: Non-SMP. (
CONFIG_SMP=n
).- User Context
The kernel executing on behalf of a particular process (ie. a system call or trap) or kernel thread. You can tell which process with the
current
macro.) Not to be confused with userspace. Can be interrupted by software or hardware interrupts.- Userspace
A process executing its own code outside the kernel.