Atomic operations in QEMU

CPUs perform independent memory operations effectively in random order. but this can be a problem for CPU-CPU interaction (including interactions between QEMU and the guest). Multi-threaded programs use various tools to instruct the compiler and the CPU to restrict the order to something that is consistent with the expectations of the programmer.

The most basic tool is locking. Mutexes, condition variables and semaphores are used in QEMU, and should be the default approach to synchronization. Anything else is considerably harder, but it’s also justified more often than one would like; the most performance-critical parts of QEMU in particular require a very low level approach to concurrency, involving memory barriers and atomic operations. The semantics of concurrent memory accesses are governed by the C11 memory model.

QEMU provides a header, qemu/atomic.h, which wraps C11 atomics to provide better portability and a less verbose syntax. qemu/atomic.h provides macros that fall in three camps:

  • compiler barriers: barrier();
  • weak atomic access and manual memory barriers: atomic_read(), atomic_set(), smp_rmb(), smp_wmb(), smp_mb(), smp_mb_acquire(), smp_mb_release(), smp_read_barrier_depends();
  • sequentially consistent atomic access: everything else.

In general, use of qemu/atomic.h should be wrapped with more easily used data structures (e.g. the lock-free singly-linked list operations QSLIST_INSERT_HEAD_ATOMIC and QSLIST_MOVE_ATOMIC) or synchronization primitives (such as RCU, QemuEvent or QemuLockCnt). Bare use of atomic operations and memory barriers should be limited to inter-thread checking of flags and documented thoroughly.

Compiler memory barrier

barrier() prevents the compiler from moving the memory accesses on either side of it to the other side. The compiler barrier has no direct effect on the CPU, which may then reorder things however it wishes.

barrier() is mostly used within qemu/atomic.h itself. On some architectures, CPU guarantees are strong enough that blocking compiler optimizations already ensures the correct order of execution. In this case, qemu/atomic.h will reduce stronger memory barriers to simple compiler barriers.

Still, barrier() can be useful when writing code that can be interrupted by signal handlers.

Sequentially consistent atomic access

Most of the operations in the qemu/atomic.h header ensure sequential consistency, where “the result of any execution is the same as if the operations of all the processors were executed in some sequential order, and the operations of each individual processor appear in this sequence in the order specified by its program”.

qemu/atomic.h provides the following set of atomic read-modify-write operations:

void atomic_inc(ptr)
void atomic_dec(ptr)
void atomic_add(ptr, val)
void atomic_sub(ptr, val)
void atomic_and(ptr, val)
void atomic_or(ptr, val)

typeof(*ptr) atomic_fetch_inc(ptr)
typeof(*ptr) atomic_fetch_dec(ptr)
typeof(*ptr) atomic_fetch_add(ptr, val)
typeof(*ptr) atomic_fetch_sub(ptr, val)
typeof(*ptr) atomic_fetch_and(ptr, val)
typeof(*ptr) atomic_fetch_or(ptr, val)
typeof(*ptr) atomic_fetch_xor(ptr, val)
typeof(*ptr) atomic_fetch_inc_nonzero(ptr)
typeof(*ptr) atomic_xchg(ptr, val)
typeof(*ptr) atomic_cmpxchg(ptr, old, new)

all of which return the old value of *ptr. These operations are polymorphic; they operate on any type that is as wide as a pointer or smaller.

Similar operations return the new value of *ptr:

typeof(*ptr) atomic_inc_fetch(ptr)
typeof(*ptr) atomic_dec_fetch(ptr)
typeof(*ptr) atomic_add_fetch(ptr, val)
typeof(*ptr) atomic_sub_fetch(ptr, val)
typeof(*ptr) atomic_and_fetch(ptr, val)
typeof(*ptr) atomic_or_fetch(ptr, val)
typeof(*ptr) atomic_xor_fetch(ptr, val)

qemu/atomic.h also provides loads and stores that cannot be reordered with each other:

typeof(*ptr) atomic_mb_read(ptr)
void         atomic_mb_set(ptr, val)

However these do not provide sequential consistency and, in particular, they do not participate in the total ordering enforced by sequentially-consistent operations. For this reason they are deprecated. They should instead be replaced with any of the following (ordered from easiest to hardest):

  • accesses inside a mutex or spinlock
  • lightweight synchronization primitives such as QemuEvent
  • RCU operations (atomic_rcu_read, atomic_rcu_set) when publishing or accessing a new version of a data structure
  • other atomic accesses: atomic_read and atomic_load_acquire for loads, atomic_set and atomic_store_release for stores, smp_mb to forbid reordering subsequent loads before a store.

Weak atomic access and manual memory barriers

Compared to sequentially consistent atomic access, programming with weaker consistency models can be considerably more complicated. The only guarantees that you can rely upon in this case are:

  • atomic accesses will not cause data races (and hence undefined behavior); ordinary accesses instead cause data races if they are concurrent with other accesses of which at least one is a write. In order to ensure this, the compiler will not optimize accesses out of existence, create unsolicited accesses, or perform other similar optimzations.
  • acquire operations will appear to happen, with respect to the other components of the system, before all the LOAD or STORE operations specified afterwards.
  • release operations will appear to happen, with respect to the other components of the system, after all the LOAD or STORE operations specified before.
  • release operations will synchronize with acquire operations; see Acquire/release pairing and the synchronizes-with relation for a detailed explanation.

When using this model, variables are accessed with:

  • atomic_read() and atomic_set(); these prevent the compiler from optimizing accesses out of existence and creating unsolicited accesses, but do not otherwise impose any ordering on loads and stores: both the compiler and the processor are free to reorder them.
  • atomic_load_acquire(), which guarantees the LOAD to appear to happen, with respect to the other components of the system, before all the LOAD or STORE operations specified afterwards. Operations coming before atomic_load_acquire() can still be reordered after it.
  • atomic_store_release(), which guarantees the STORE to appear to happen, with respect to the other components of the system, after all the LOAD or STORE operations specified before. Operations coming after atomic_store_release() can still be reordered before it.

Restrictions to the ordering of accesses can also be specified using the memory barrier macros: smp_rmb(), smp_wmb(), smp_mb(), smp_mb_acquire(), smp_mb_release(), smp_read_barrier_depends().

Memory barriers control the order of references to shared memory. They come in six kinds:

  • smp_rmb() guarantees that all the LOAD operations specified before the barrier will appear to happen before all the LOAD operations specified after the barrier with respect to the other components of the system.

    In other words, smp_rmb() puts a partial ordering on loads, but is not required to have any effect on stores.

  • smp_wmb() guarantees that all the STORE operations specified before the barrier will appear to happen before all the STORE operations specified after the barrier with respect to the other components of the system.

    In other words, smp_wmb() puts a partial ordering on stores, but is not required to have any effect on loads.

  • smp_mb_acquire() guarantees that all the LOAD operations specified before the barrier will appear to happen before all the LOAD or STORE operations specified after the barrier with respect to the other components of the system.

  • smp_mb_release() guarantees that all the STORE operations specified after the barrier will appear to happen after all the LOAD or STORE operations specified before the barrier with respect to the other components of the system.

  • smp_mb() guarantees that all the LOAD and STORE operations specified before the barrier will appear to happen before all the LOAD and STORE operations specified after the barrier with respect to the other components of the system.

    smp_mb() puts a partial ordering on both loads and stores. It is stronger than both a read and a write memory barrier; it implies both smp_mb_acquire() and smp_mb_release(), but it also prevents STOREs coming before the barrier from overtaking LOADs coming after the barrier and vice versa.

  • smp_read_barrier_depends() is a weaker kind of read barrier. On most processors, whenever two loads are performed such that the second depends on the result of the first (e.g., the first load retrieves the address to which the second load will be directed), the processor will guarantee that the first LOAD will appear to happen before the second with respect to the other components of the system. However, this is not always true—for example, it was not true on Alpha processors. Whenever this kind of access happens to shared memory (that is not protected by a lock), a read barrier is needed, and smp_read_barrier_depends() can be used instead of smp_rmb().

    Note that the first load really has to have a _data_ dependency and not a control dependency. If the address for the second load is dependent on the first load, but the dependency is through a conditional rather than actually loading the address itself, then it’s a _control_ dependency and a full read barrier or better is required.

Memory barriers and atomic_load_acquire/atomic_store_release are mostly used when a data structure has one thread that is always a writer and one thread that is always a reader:

thread 1 thread 2
atomic_store_release(&a, x);
atomic_store_release(&b, y);
y = atomic_load_acquire(&b);
x = atomic_load_acquire(&a);

In this case, correctness is easy to check for using the “pairing” trick that is explained below.

Sometimes, a thread is accessing many variables that are otherwise unrelated to each other (for example because, apart from the current thread, exactly one other thread will read or write each of these variables). In this case, it is possible to “hoist” the barriers outside a loop. For example:

before after
n = 0;
for (i = 0; i < 10; i++)
  n += atomic_load_acquire(&a[i]);
n = 0;
for (i = 0; i < 10; i++)
  n += atomic_read(&a[i]);
smp_mb_acquire();
for (i = 0; i < 10; i++)
  atomic_store_release(&a[i], false);
smp_mb_release();
for (i = 0; i < 10; i++)
  atomic_set(&a[i], false);

Splitting a loop can also be useful to reduce the number of barriers:

before after
n = 0;
for (i = 0; i < 10; i++) {
  atomic_store_release(&a[i], false);
  smp_mb();
  n += atomic_read(&b[i]);
}
smp_mb_release();
for (i = 0; i < 10; i++)
  atomic_set(&a[i], false);
smb_mb();
n = 0;
for (i = 0; i < 10; i++)
  n += atomic_read(&b[i]);

In this case, a smp_mb_release() is also replaced with a (possibly cheaper, and clearer as well) smp_wmb():

before after
for (i = 0; i < 10; i++) {
  atomic_store_release(&a[i], false);
  atomic_store_release(&b[i], false);
}
smp_mb_release();
for (i = 0; i < 10; i++)
  atomic_set(&a[i], false);
smb_wmb();
for (i = 0; i < 10; i++)
  atomic_set(&b[i], false);

Acquire/release pairing and the synchronizes-with relation

Atomic operations other than atomic_set() and atomic_read() have either acquire or release semantics [1]. This has two effects:

[1]Read-modify-write operations can have both—acquire applies to the read part, and release to the write.
  • within a thread, they are ordered either before subsequent operations (for acquire) or after previous operations (for release).
  • if a release operation in one thread synchronizes with an acquire operation in another thread, the ordering constraints propagates from the first to the second thread. That is, everything before the release operation in the first thread is guaranteed to happen before everything after the acquire operation in the second thread.

The concept of acquire and release semantics is not exclusive to atomic operations; almost all higher-level synchronization primitives also have acquire or release semantics. For example:

  • pthread_mutex_lock has acquire semantics, pthread_mutex_unlock has release semantics and synchronizes with a pthread_mutex_lock for the same mutex.
  • pthread_cond_signal and pthread_cond_broadcast have release semantics; pthread_cond_wait has both release semantics (synchronizing with pthread_mutex_lock) and acquire semantics (synchronizing with pthread_mutex_unlock and signaling of the condition variable).
  • pthread_create has release semantics and synchronizes with the start of the new thread; pthread_join has acquire semantics and synchronizes with the exiting of the thread.
  • qemu_event_set has release semantics, qemu_event_wait has acquire semantics.

For example, in the following example there are no atomic accesses, but still thread 2 is relying on the synchronizes-with relation between pthread_exit (release) and pthread_join (acquire):

thread 1 thread 2
*a = 1;
pthread_exit(a);
pthread_join(thread1, &a);
x = *a;

Synchronization between threads basically descends from this pairing of a release operation and an acquire operation. Therefore, atomic operations other than atomic_set() and atomic_read() will almost always be paired with another operation of the opposite kind: an acquire operation will pair with a release operation and vice versa. This rule of thumb is extremely useful; in the case of QEMU, however, note that the other operation may actually be in a driver that runs in the guest!

smp_read_barrier_depends(), smp_rmb(), smp_mb_acquire(), atomic_load_acquire() and atomic_rcu_read() all count as acquire operations. smp_wmb(), smp_mb_release(), atomic_store_release() and atomic_rcu_set() all count as release operations. smp_mb() counts as both acquire and release, therefore it can pair with any other atomic operation. Here is an example:

thread 1 thread 2
atomic_set(&a, 1);
smp_wmb();
atomic_set(&b, 2);
x = atomic_read(&b);
smp_rmb();
y = atomic_read(&a);

Note that a load-store pair only counts if the two operations access the same variable: that is, a store-release on a variable x synchronizes with a load-acquire on a variable x, while a release barrier synchronizes with any acquire operation. The following example shows correct synchronization:

thread 1 thread 2
atomic_set(&a, 1);
atomic_store_release(&b, 2);
x = atomic_load_acquire(&b);
y = atomic_read(&a);

Acquire and release semantics of higher-level primitives can also be relied upon for the purpose of establishing the synchronizes with relation.

Note that the “writing” thread is accessing the variables in the opposite order as the “reading” thread. This is expected: stores before a release operation will normally match the loads after the acquire operation, and vice versa. In fact, this happened already in the pthread_exit/pthread_join example above.

Finally, this more complex example has more than two accesses and data dependency barriers. It also does not use atomic accesses whenever there cannot be a data race:

thread 1 thread 2
b[2] = 1;
smp_wmb();
x->i = 2;
smp_wmb();
atomic_set(&a, x);
x = atomic_read(&a);
smp_read_barrier_depends();
y = x->i;
smp_read_barrier_depends();
z = b[y];

Comparison with Linux kernel primitives

Here is a list of differences between Linux kernel atomic operations and memory barriers, and the equivalents in QEMU:

  • atomic operations in Linux are always on a 32-bit int type and use a boxed atomic_t type; atomic operations in QEMU are polymorphic and use normal C types.

  • Originally, atomic_read and atomic_set in Linux gave no guarantee at all. Linux 4.1 updated them to implement volatile semantics via ACCESS_ONCE (or the more recent READ/WRITE_ONCE).

    QEMU’s atomic_read and atomic_set implement C11 atomic relaxed semantics if the compiler supports it, and volatile semantics otherwise. Both semantics prevent the compiler from doing certain transformations; the difference is that atomic accesses are guaranteed to be atomic, while volatile accesses aren’t. Thus, in the volatile case we just cross our fingers hoping that the compiler will generate atomic accesses, since we assume the variables passed are machine-word sized and properly aligned.

    No barriers are implied by atomic_read and atomic_set in either Linux or QEMU.

  • atomic read-modify-write operations in Linux are of three kinds:

    atomic_OP returns void
    atomic_OP_return returns new value of the variable
    atomic_fetch_OP returns the old value of the variable
    atomic_cmpxchg returns the old value of the variable

    In QEMU, the second kind is named atomic_OP_fetch.

  • different atomic read-modify-write operations in Linux imply a different set of memory barriers; in QEMU, all of them enforce sequential consistency.

  • in QEMU, atomic_read() and atomic_set() do not participate in the total ordering enforced by sequentially-consistent operations. This is because QEMU uses the C11 memory model. The following example is correct in Linux but not in QEMU:

    Linux (correct) QEMU (incorrect)
    a = atomic_fetch_add(&x, 2);
    b = READ_ONCE(&y);
    
    a = atomic_fetch_add(&x, 2);
    b = atomic_read(&y);
    

    because the read of y can be moved (by either the processor or the compiler) before the write of x.

    Fixing this requires an smp_mb() memory barrier between the write of x and the read of y. In the common case where only one thread writes x, it is also possible to write it like this:

    QEMU (correct)
    a = atomic_read(&x);
    atomic_set(&x, a + 2);
    smp_mb();
    b = atomic_read(&y);
    

Sources

  • Documentation/memory-barriers.txt from the Linux kernel