28
Locks

From the introduction to concurrency, we saw one of the fundamental
problems in concurrent programming: we would like to execute a series
of instructions atomically, but due to the presence of interrupts on a single
processor (or multiple threads executing on multiple processors concurrently), we couldn’t. In this chapter, we thus attack this problem directly,
with the introduction of something referred to as a lock. Programmers
annotate source code with locks, putting them around critical sections,
and thus ensure that any such critical section executes as if it were a single atomic instruction.

28.1

Locks: The Basic Idea
As an example, assume our critical section looks like this, the canonical
update of a shared variable:
balance = balance + 1;

Of course, other critical sections are possible, such as adding an element to a linked list or other more complex updates to shared structures,
but we’ll just keep to this simple example for now. To use a lock, we add
some code around the critical section like this:
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lock_t mutex; // some globally-allocated lock ’mutex’
...
lock(&mutex);
balance = balance + 1;
unlock(&mutex);

A lock is just a variable, and thus to use one, you must declare a lock
variable of some kind (such as mutex above). This lock variable (or just
“lock” for short) holds the state of the lock at any instant in time. It is either available (or unlocked or free) and thus no thread holds the lock, or
acquired (or locked or held), and thus exactly one thread holds the lock
and presumably is in a critical section. We could store other information
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in the data type as well, such as which thread holds the lock, or a queue
for ordering lock acquisition, but information like that is hidden from the
user of the lock.
The semantics of the lock() and unlock() routines are simple. Calling the routine lock() tries to acquire the lock; if no other thread holds
the lock (i.e., it is free), the thread will acquire the lock and enter the critical section; this thread is sometimes said to be the owner of the lock. If
another thread then calls lock() on that same lock variable (mutex in
this example), it will not return while the lock is held by another thread;
in this way, other threads are prevented from entering the critical section
while the first thread that holds the lock is in there.
Once the owner of the lock calls unlock(), the lock is now available
(free) again. If no other threads are waiting for the lock (i.e., no other
thread has called lock() and is stuck therein), the state of the lock is
simply changed to free. If there are waiting threads (stuck in lock()),
one of them will (eventually) notice (or be informed of) this change of the
lock’s state, acquire the lock, and enter the critical section.
Locks provide some minimal amount of control over scheduling to
programmers. In general, we view threads as entities created by the programmer but scheduled by the OS, in any fashion that the OS chooses.
Locks yield some of that control back to the programmer; by putting
a lock around a section of code, the programmer can guarantee that no
more than a single thread can ever be active within that code. Thus locks
help transform the chaos that is traditional OS scheduling into a more
controlled activity.

28.2

Pthread Locks
The name that the POSIX library uses for a lock is a mutex, as it is used
to provide mutual exclusion between threads, i.e., if one thread is in the
critical section, it excludes the others from entering until it has completed
the section. Thus, when you see the following POSIX threads code, you
should understand that it is doing the same thing as above (we again use
our wrappers that check for errors upon lock and unlock):

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pthread_mutex_t lock = PTHREAD_MUTEX_INITIALIZER;

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Pthread_mutex_lock(&lock);
// wrapper for pthread_mutex_lock()
balance = balance + 1;
Pthread_mutex_unlock(&lock);

You might also notice here that the POSIX version passes a variable
to lock and unlock, as we may be using different locks to protect different
variables. Doing so can increase concurrency: instead of one big lock that
is used any time any critical section is accessed (a coarse-grained locking
strategy), one will often protect different data and data structures with
different locks, thus allowing more threads to be in locked code at once
(a more fine-grained approach).

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28.3

3

Building A Lock
By now, you should have some understanding of how a lock works,
from the perspective of a programmer. But how should we build a lock?
What hardware support is needed? What OS support? It is this set of
questions we address in the rest of this chapter.
The Crux: H OW T O B UILD A L OCK
How can we build an efficient lock? Efficient locks provided mutual
exclusion at low cost, and also might attain a few other properties we
discuss below. What hardware support is needed? What OS support?
To build a working lock, we will need some help from our old friend,
the hardware, as well as our good pal, the OS. Over the years, a number of different hardware primitives have been added to the instruction
sets of various computer architectures; while we won’t study how these
instructions are implemented (that, after all, is the topic of a computer
architecture class), we will study how to use them in order to build a mutual exclusion primitive like a lock. We will also study how the OS gets
involved to complete the picture and enable us to build a sophisticated
locking library.

28.4

Evaluating Locks
Before building any locks, we should first understand what our goals
are, and thus we ask how to evaluate the efficacy of a particular lock
implementation. To evaluate whether a lock works (and works well), we
should first establish some basic criteria. The first is whether the lock does
its basic task, which is to provide mutual exclusion. Basically, does the
lock work, preventing multiple threads from entering a critical section?
The second is fairness. Does each thread contending for the lock get
a fair shot at acquiring it once it is free? Another way to look at this is
by examining the more extreme case: does any thread contending for the
lock starve while doing so, thus never obtaining it?
The final criterion is performance, specifically the time overheads added
by using the lock. There are a few different cases that are worth considering here. One is the case of no contention; when a single thread
is running and grabs and releases the lock, what is the overhead of doing so? Another is the case where multiple threads are contending for
the lock on a single CPU; in this case, are there performance concerns? Finally, how does the lock perform when there are multiple CPUs involved,
and threads on each contending for the lock? By comparing these different scenarios, we can better understand the performance impact of using
various locking techniques, as described below.

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28.5

Controlling Interrupts
One of the earliest solutions used to provide mutual exclusion was
to disable interrupts for critical sections; this solution was invented for
single-processor systems. The code would look like this:

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void lock() {
DisableInterrupts();
}
void unlock() {
EnableInterrupts();
}

Assume we are running on such a single-processor system. By turning off interrupts (using some kind of special hardware instruction) before entering a critical section, we ensure that the code inside the critical
section will not be interrupted, and thus will execute as if it were atomic.
When we are finished, we re-enable interrupts (again, via a hardware instruction) and thus the program proceeds as usual.
The main positive of this approach is its simplicity. You certainly don’t
have to scratch your head too hard to figure out why this works. Without
interruption, a thread can be sure that the code it executes will execute
and that no other thread will interfere with it.
The negatives, unfortunately, are many. First, this approach requires
us to allow any calling thread to perform a privileged operation (turning
interrupts on and off), and thus trust that this facility is not abused. As
you already know, any time we are required to trust an arbitrary program, we are probably in trouble. Here, the trouble manifests in numerous ways: a greedy program could call lock() at the beginning of its
execution and thus monopolize the processor; worse, an errant or malicious program could call lock() and go into an endless loop. In this
latter case, the OS never regains control of the system, and there is only
one recourse: restart the system. Using interrupt disabling as a generalpurpose synchronization solution requires too much trust in applications.
Second, the approach does not work on multiprocessors. If multiple
threads are running on different CPUs, and each try to enter the same
critical section, it does not matter whether interrupts are disabled; threads
will be able to run on other processors, and thus could enter the critical
section. As multiprocessors are now commonplace, our general solution
will have to do better than this.
Third, and probably least important, this approach can be inefficient.
Compared to normal instruction execution, code that masks or unmasks
interrupts tends to be executed slowly by modern CPUs.
For these reasons, turning off interrupts is only used in limited contexts as a mutual-exclusion primitive. For example, in some cases an
operating system itself will sometimes use interrupt masking to guarantee atomicity when accessing its own data structures, or at least to prevent certain messy interrupt handling situations from arising. This usage
makes sense, as the trust issue disappears inside the OS, which always
trusts itself to perform privileged operations anyhow.

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5

A SIDE : D EKKER ’ S AND P ETERSON ’ S A LGORITHMS
In the 1960’s, Dijkstra posed the concurrency problem to his friends,
and one of them, a mathematician named Theodorus Jozef Dekker, came
up with a solution [D68]. Unlike the solutions we discuss here, which use
special hardware instructions and even OS support, Dekker’s approach
uses just loads and stores (assuming they are atomic with respect to each
other, which was true on early hardware).
Dekker’s approach was later refined by Peterson [P81] (and thus “Peterson’s algorithm”), shown here. Once again, just loads and stores are
used, and the idea is to ensure that two threads never enter a critical section at the same time. Here is Peterson’s algorithm (for two threads); see
if you can understand it.
int flag[2];
int turn;
void init() {
flag[0] = flag[1] = 0;
// 1->thread wants to grab lock
turn = 0;
// whose turn? (thread 0 or 1?)
}
void lock() {
flag[self] = 1;
// self: thread ID of caller
turn = 1 - self;
// make it other thread’s turn
while ((flag[1-self] == 1) && (turn == 1 - self))
; // spin-wait
}
void unlock() {
flag[self] = 0;
// simply undo your intent
}

For some reason, developing locks that work without special hardware support became all the rage for a while, giving theory-types a lot
of problems to work on. Of course, this all became quite useless when
people realized it is much easier to assume a little hardware support (and
indeed that support had been around from the very earliest days of multiprocessing). Further, algorithms like the ones above don’t work on modern hardware (due to relaxed memory consistency models), thus making
them even less useful than they were before. Yet more research relegated
to the dustbin of history...

28.6

Test And Set (Atomic Exchange)
Because disabling interrupts does not work on multiple processors,
system designers started to invent hardware support for locking. The
earliest multiprocessor systems, such as the Burroughs B5000 in the early
1960’s [M82], had such support; today all systems provide this type of
support, even for single CPU systems.

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typedef struct __lock_t { int flag; } lock_t;

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void init(lock_t *mutex) {
// 0 -> lock is available, 1 -> held
mutex->flag = 0;
}

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void lock(lock_t *mutex) {
while (mutex->flag == 1) // TEST the flag
; // spin-wait (do nothing)
mutex->flag = 1;
// now SET it!
}

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void unlock(lock_t *mutex) {
mutex->flag = 0;
}

Figure 28.1: First Attempt: A Simple Flag
The simplest bit of hardware support to understand is what is known
as a test-and-set instruction, also known as atomic exchange. To understand how test-and-set works, let’s first try to build a simple lock without
it. In this failed attempt, we use a simple flag variable to denote whether
the lock is held or not.
In this first attempt (Figure 28.1), the idea is quite simple: use a simple
variable to indicate whether some thread has possession of a lock. The
first thread that enters the critical section will call lock(), which tests
whether the flag is equal to 1 (in this case, it is not), and then sets the flag
to 1 to indicate that the thread now holds the lock. When finished with
the critical section, the thread calls unlock() and clears the flag, thus
indicating that the lock is no longer held.
If another thread happens to call lock() while that first thread is in
the critical section, it will simply spin-wait in the while loop for that
thread to call unlock() and clear the flag. Once that first thread does
so, the waiting thread will fall out of the while loop, set the flag to 1 for
itself, and proceed into the critical section.
Unfortunately, the code has two problems: one of correctness, and another of performance. The correctness problem is simple to see once you
get used to thinking about concurrent programming. Imagine the code
interleaving in Table 28.1 (assume flag=0 to begin).
Thread 1
call lock()
while (flag == 1)
interrupt: switch to Thread 2

Thread 2

call lock()
while (flag == 1)
flag = 1;
interrupt: switch to Thread 1
flag = 1; // set flag to 1 (too!)

Table 28.1: Trace: No Mutual Exclusion

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T IP : T HINK A BOUT C ONCURRENCY A S M ALICIOUS S CHEDULER
What we also get from this example is a sense of the approach we
need to take when trying to understand concurrent execution. What you
are really trying to do is to pretend you are a malicious scheduler, one
that interrupts threads at the most inopportune of times in order to foil
their feeble attempts at building synchronization primitives. Although
the exact sequence of interrupts may be improbable, it is possible, and that
is all we need to show to demonstrate that a particular approach does not
work.
As you can see from this interleaving, with timely (untimely?) interrupts, we can easily produce a case where both threads set their flags to 1
and both threads are thus able to enter the critical section. This is bad! We
have obviously failed to provide the most basic requirement: providing
mutual exclusion.
The performance problem, which we will address more later on, is the
fact that the way a thread waits to acquire a lock that is already held:
it endlessly checks the value of flag, a technique known as spin-waiting.
Spin-waiting wastes time waiting for another thread to release a lock. The
waste is exceptionally high on a uniprocessor, where the thread that the
waiter is waiting for cannot even run (at least, until a context switch occurs)! Thus, as we move forward and develop more sophisticated solutions, we should also consider ways to avoid this kind of waste.

28.7

Building A Working Spin Lock
While the idea behind the example above is a good one, it is not possible to implement without some support from the hardware. Fortunately,
some systems provide an instruction to support the creation of simple
locks based on this concept. This more powerful instruction has different names – on SPARC, it is the load/store unsigned byte instruction
(ldstub), whereas on x86, it is the atomic exchange instruction (xchg)
– but basically does the same thing across platforms, and is usually generally referred to as test-and-set. We define what the test-and-set instruction does with the following C code snippet:

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int TestAndSet(int *ptr, int new) {
int old = *ptr; // fetch old value at ptr
// store ’new’ into ptr
*ptr = new;
return old;
// return the old value
}

What the test-and-set instruction does is as follows. It returns the old
value pointed to by the ptr, and simultaneously updates said value to
new. The key, of course, is that this sequence of operations is performed
atomically. The reason it is called “test and set” is that it enables you
to “test” the old value (which is what is returned) while simultaneously

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typedef struct __lock_t {
int flag;
} lock_t;

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void init(lock_t *lock) {
// 0 indicates that lock is available, 1 that it is held
lock->flag = 0;
}

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void lock(lock_t *lock) {
while (TestAndSet(&lock->flag, 1) == 1)
; // spin-wait (do nothing)
}

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void unlock(lock_t *lock) {
lock->flag = 0;
}

Figure 28.2: A Simple Spin Lock Using Test-and-set
“setting” the memory location to a new value; as it turns out, this slightly
more powerful instruction is enough to build a simple spin lock, as we
now examine in Figure 28.2.
Let’s make sure we understand why this works. Imagine first the case
where a thread calls lock() and no other thread currently holds the lock;
thus, flag should be 0. When the thread then calls TestAndSet(flag,
1), the routine will return the old value of flag, which is 0; thus, the calling thread, which is testing the value of flag, will not get caught spinning
in the while loop and will acquire the lock. The thread will also atomically set the value to 1, thus indicating that the lock is now held. When
the thread is finished with its critical section, it calls unlock() to set the
flag back to zero.
The second case we can imagine arises when one thread already has
the lock held (i.e., flag is 1). In this case, this thread will call lock() and
then call TestAndSet(flag, 1) as well. This time, TestAndSet()
will return the old value at flag, which is 1 (because the lock is held),
while simultaneously setting it to 1 again. As long as the lock is held by
another thread, TestAndSet() will repeatedly return 1, and thus this
thread will spin and spin until the lock is finally released. When the flag is
finally set to 0 by some other thread, this thread will call TestAndSet()
again, which will now return 0 while atomically setting the value to 1 and
thus acquire the lock and enter the critical section.
By making both the test (of the old lock value) and set (of the new
value) a single atomic operation, we ensure that only one thread acquires
the lock. And that’s how to build a working mutual exclusion primitive!
You may also now understand why this type of lock is usually referred
to as a spin lock. It is the simplest type of lock to build, and simply spins,
using CPU cycles, until the lock becomes available. To work correctly
on a single processor, it requires a preemptive scheduler (i.e., one that
will interrupt a thread via a timer, in order to run a different thread, from
time to time). Without preemption, spin locks don’t make much sense on
a single CPU, as a thread spinning on a CPU will never relinquish it.

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28.8

9

Evaluating Spin Locks
Given our basic spin lock, we can now evaluate how effective it is
along our previously described axes. The most important aspect of a lock
is correctness: does it provide mutual exclusion? The answer here is obviously yes: the spin lock only allows a single thread to enter the critical
section at a time. Thus, we have a correct lock.
The next axis is fairness. How fair is a spin lock to a waiting thread?
Can you guarantee that a waiting thread will ever enter the critical section? The answer here, unfortunately, is bad news: spin locks don’t provide any fairness guarantees. Indeed, a thread spinning may spin forever,
under contention. Spin locks are not fair and may lead to starvation.
The final axis is performance. What are the costs of using a spin lock?
To analyze this more carefully, we suggest thinking about a few different
cases. In the first, imagine threads competing for the lock on a single
processor; in the second, consider the threads as spread out across many
processors.
For spin locks, in the single CPU case, performance overheads can
be quite painful; imagine the case where the thread holding the lock is
pre-empted within a critical section. The scheduler might then run every
other thread (imagine there are N − 1 others), each of which tries to acquire the lock. In this case, each of those threads will spin for the duration
of a time slice before giving up the CPU, a waste of CPU cycles.
However, on multiple CPUs, spin locks work reasonably well (if the
number of threads roughly equals the number of CPUs). The thinking
goes as follows: imagine Thread A on CPU 1 and Thread B on CPU 2,
both contending for a lock. If Thread A (CPU 1) grabs the lock, and then
Thread B tries to, B will spin (on CPU 2). However, presumably the critical section is short, and thus soon the lock becomes available, and is acquired by Thread B. Spinning to wait for a lock held on another processor
doesn’t waste many cycles in this case, and thus can be quite effective.

28.9

Compare-And-Swap
Another hardware primitive that some systems provide is known as
the compare-and-swap instruction (as it is called on SPARC, for example), or compare-and-exchange (as it called on x86). The C pseudocode
for this single instruction is found in Figure 28.3.
The basic idea is for compare-and-swap to test whether the value at the

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int CompareAndSwap(int *ptr, int expected, int new) {
int actual = *ptr;
if (actual == expected)
*ptr = new;
return actual;
}

Figure 28.3: Compare-and-swap

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address specified by ptr is equal to expected; if so, update the memory
location pointed to by ptr with the new value. If not, do nothing. In
either case, return the actual value at that memory location, thus allowing
the code calling compare-and-swap to know whether it succeeded or not.
With the compare-and-swap instruction, we can build a lock in a manner quite similar to that with test-and-set. For example, we could just
replace the lock() routine above with the following:
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void lock(lock_t *lock) {
while (CompareAndSwap(&lock->flag, 0, 1) == 1)
; // spin
}

The rest of the code is the same as the test-and-set example above.
This code works quite similarly; it simply checks if the flag is 0 and if
so, atomically swaps in a 1 thus acquiring the lock. Threads that try to
acquire the lock while it is held will get stuck spinning until the lock is
finally released.
If you want to see how to really make a C-callable x86-version of
compare-and-swap, this code sequence might be useful (from [S05]):
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char CompareAndSwap(int *ptr, int old, int new) {
unsigned char ret;

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// Note that sete sets a ’byte’ not the word
__asm__ __volatile__ (
" lock\n"
" cmpxchgl %2,%1\n"
" sete %0\n"
: "=q" (ret), "=m" (*ptr)
: "r" (new), "m" (*ptr), "a" (old)
: "memory");
return ret;

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}

Finally, as you may have sensed, compare-and-swap is a more powerful instruction than test-and-set. We will make some use of this power in
the future when we briefly delve into wait-free synchronization [H91].
However, if we just build a simple spin lock with it, its behavior is identical to the spin lock we analyzed above.

28.10

Load-Linked and Store-Conditional
Some platforms provide a pair of instructions that work in concert to
help build critical sections. On the MIPS architecture [H93], for example,
the load-linked and store-conditional instructions can be used in tandem
to build locks and other concurrent structures. The C pseudocode for
these instructions is as found in Figure 28.4. Alpha, PowerPC, and ARM
provide similar instructions [W09].

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int LoadLinked(int *ptr) {
return *ptr;
}

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int StoreConditional(int *ptr, int value) {
if (no one has updated *ptr since the LoadLinked to this address) {
*ptr = value;
return 1; // success!
} else {
return 0; // failed to update
}
}

Figure 28.4: Load-linked And Store-conditional
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void lock(lock_t *lock) {
while (1) {
while (LoadLinked(&lock->flag) == 1)
; // spin until it’s zero
if (StoreConditional(&lock->flag, 1) == 1)
return; // if set-it-to-1 was a success: all done
// otherwise: try it all over again
}
}

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void unlock(lock_t *lock) {
lock->flag = 0;
}

Figure 28.5: Using LL/SC To Build A Lock
The load-linked operates much like a typical load instruction, and simply fetches a value from memory and places it in a register. The key difference comes with the store-conditional, which only succeeds (and updates
the value stored at the address just load-linked from) if no intermittent
store to the address has taken place. In the case of success, the storeconditional returns 1 and updates the value at ptr to value; if it fails,
the value at ptr is not updated and 0 is returned.
As a challenge to yourself, try thinking about how to build a lock using
load-linked and store-conditional. Then, when you are finished, look at
the code below which provides one simple solution. Do it! The solution
is in Figure 28.5.
The lock() code is the only interesting piece. First, a thread spins
waiting for the flag to be set to 0 (and thus indicate the lock is not held).
Once so, the thread tries to acquire the lock via the store-conditional; if it
succeeds, the thread has atomically changed the flag’s value to 1 and thus
can proceed into the critical section.
Note how failure of the store-conditional might arise. One thread calls
lock() and executes the load-linked, returning 0 as the lock is not held.
Before it can attempt the store-conditional, it is interrupted and another
thread enters the lock code, also executing the load-linked instruction,
and also getting a 0 and continuing. At this point, two threads have
each executed the load-linked and each are about to attempt the storeconditional. The key feature of these instructions is that only one of these
threads will succeed in updating the flag to 1 and thus acquire the lock;

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T IP : L ESS C ODE I S B ETTER C ODE (L AUER ’ S L AW )
Programmers tend to brag about how much code they wrote to do something. Doing so is fundamentally broken. What one should brag about,
rather, is how little code one wrote to accomplish a given task. Short,
concise code is always preferred; it is likely easier to understand and has
fewer bugs. As Hugh Lauer said, when discussing the construction of
the Pilot operating system: “If the same people had twice as much time,
they could produce as good of a system in half the code.” [L81] We’ll call
this Lauer’s Law, and it is well worth remembering. So next time you’re
bragging about how much code you wrote to finish the assignment, think
again, or better yet, go back, rewrite, and make the code as clear and concise as possible.
the second thread to attempt the store-conditional will fail (because the
other thread updated the value of flag between its load-linked and storeconditional) and thus have to try to acquire the lock again.
In class a few years ago, undergraduate student David Capel suggested a more concise form of the above, for those of you who enjoy
short-circuiting boolean conditionals. See if you can figure out why it
is equivalent. It certainly is shorter!
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28.11

void lock(lock_t *lock) {
while (LoadLinked(&lock->flag)||!StoreConditional(&lock->flag, 1))
; // spin
}

Fetch-And-Add
One final hardware primitive is the fetch-and-add instruction, which
atomically increments a value while returning the old value at a particular address. The C pseudocode for the fetch-and-add instruction looks
like this:

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int FetchAndAdd(int *ptr) {
int old = *ptr;
*ptr = old + 1;
return old;
}

In this example, we’ll use fetch-and-add to build a more interesting
ticket lock, as introduced by Mellor-Crummey and Scott [MS91]. The
lock and unlock code looks like what you see in Figure 28.6.
Instead of a single value, this solution uses a ticket and turn variable in
combination to build a lock. The basic operation is pretty simple: when
a thread wishes to acquire a lock, it first does an atomic fetch-and-add
on the ticket value; that value is now considered this thread’s “turn”
(myturn). The globally shared lock->turn is then used to determine
which thread’s turn it is; when (myturn == turn) for a given thread,

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typedef struct __lock_t {
int ticket;
int turn;
} lock_t;

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void lock_init(lock_t *lock) {
lock->ticket = 0;
lock->turn
= 0;
}

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void lock(lock_t *lock) {
int myturn = FetchAndAdd(&lock->ticket);
while (lock->turn != myturn)
; // spin
}

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void unlock(lock_t *lock) {
FetchAndAdd(&lock->turn);
}

Figure 28.6: Ticket Locks
it is that thread’s turn to enter the critical section. Unlock is accomplished
simply by incrementing the turn such that the next waiting thread (if
there is one) can now enter the critical section.
Note one important difference with this solution versus our previous
attempts: it ensures progress for all threads. Once a thread is assigned its
ticket value, it will be scheduled at some point in the future (once those in
front of it have passed through the critical section and released the lock).
In our previous attempts, no such guarantee existed; a thread spinning
on test-and-set (for example) could spin forever even as other threads
acquire and release the lock.

28.12

Summary: So Much Spinning
Our simple hardware-based locks are simple (only a few lines of code)
and they work (you could even prove that if you’d like to, by writing
some code), which are two excellent properties of any system or code.
However, in some cases, these solutions can be quite inefficient. Imagine
you are running two threads on a single processor. Now imagine that
one thread (thread 0) is in a critical section and thus has a lock held, and
unfortunately gets interrupted. The second thread (thread 1) now tries to
acquire the lock, but finds that it is held. Thus, it begins to spin. And spin.
Then it spins some more. And finally, a timer interrupt goes off, thread
0 is run again, which releases the lock, and finally (the next time it runs,
say), thread 1 won’t have to spin so much and will be able to acquire the
lock. Thus, any time a thread gets caught spinning in a situation like this,
it wastes an entire time slice doing nothing but checking a value that isn’t
going to change! The problem gets worse with N threads contending
for a lock; N − 1 time slices may be wasted in a similar manner, simply
spinning and waiting for a single thread to release the lock. And thus,
our next problem:

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How can we develop a lock that doesn’t needlessly waste time spinning on the CPU?
Hardware support alone cannot solve the problem. We’ll need OS support too! Let’s now figure out just how that might work.

28.13

A Simple Approach: Just Yield, Baby
Hardware support got us pretty far: working locks, and even (as with
the case of the ticket lock) fairness in lock acquisition. However, we still
have a problem: what to do when a context switch occurs in a critical
section, and threads start to spin endlessly, waiting for the interrupt (lockholding) thread to be run again?
Our first try is a simple and friendly approach: when you are going to
spin, instead give up the CPU to another thread. Or, as Al Davis might
say, “just yield, baby!” [D91]. Figure 28.7 presents the approach.
In this approach, we assume an operating system primitive yield()
which a thread can call when it wants to give up the CPU and let another thread run. Because a thread can be in one of three states (running,
ready, or blocked), you can think of this as an OS system call that moves
the caller from the running state to the ready state, and thus promotes
another thread to running.
Think about the example with two threads on one CPU; in this case,
our yield-based approach works quite well. If a thread happens to call
lock() and find a lock held, it will simply yield the CPU, and thus the
other thread will run and finish its critical section. In this simple case, the
yielding approach works well.
Let us now consider the case where there are many threads (say 100)
contending for a lock repeatedly. In this case, if one thread acquires
the lock and is preempted before releasing it, the other 99 will each call

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void init() {
flag = 0;
}

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6
7
8

void lock() {
while (TestAndSet(&flag, 1) == 1)
yield(); // give up the CPU
}

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void unlock() {
flag = 0;
}

Figure 28.7: Lock With Test-and-set And Yield

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lock(), find the lock held, and yield the CPU. Assuming some kind
of round-robin scheduler, each of the 99 will execute this run-and-yield
pattern before the thread holding the lock gets to run again. While better
than our spinning approach (which would waste 99 time slices spinning),
this approach is still costly; the cost of a context switch can be substantial,
and there is thus plenty of waste.
Worse, we have not tackled the starvation problem at all. A thread
may get caught in an endless yield loop while other threads repeatedly
enter and exit the critical section. We clearly will need an approach that
addresses this problem directly.

28.14

Using Queues: Sleeping Instead Of Spinning
The real problem with our previous approaches is that they leave too
much to chance. The scheduler determines which thread runs next; if
the scheduler makes a bad choice, a thread runs that must either spin
waiting for the lock (our first approach), or yield the CPU immediately
(our second approach). Either way, there is potential for waste and no
prevention of starvation.
Thus, we must explicitly exert some control over who gets to acquire
the lock next after the current holder releases it. To do this, we will need a
little more OS support, as well as a queue to keep track of which threads
are waiting to enter the lock.
For simplicity, we will use the support provided by Solaris, in terms of
two calls: park() to put a calling thread to sleep, and unpark(threadID)
to wake a particular thread as designated by threadID. These two routines can be used in tandem to build a lock that puts a caller to sleep if it
tries to acquire a held lock and wakes it when the lock is free. Let’s look at
the code in Figure 28.8 to understand one possible use of such primitives.
We do a couple of interesting things in this example. First, we combine
the old test-and-set idea with an explicit queue of lock waiters to make a
more efficient lock. Second, we use a queue to help control who gets the
lock next and thus avoid starvation.
You might notice how the guard is used, basically as a spin-lock around
the flag and queue manipulations the lock is using. This approach thus
doesn’t avoid spin-waiting entirely; a thread might be interrupted while
acquiring or releasing the lock, and thus cause other threads to spin-wait
for this one to run again. However, the time spent spinning is quite limited (just a few instructions inside the lock and unlock code, instead of the
user-defined critical section), and thus this approach may be reasonable.
Second, you might notice that in lock(), when a thread can not acquire the lock (it is already held), we are careful to add ourselves to a
queue (by calling the gettid() call to get the thread ID of the current
thread), set guard to 0, and yield the CPU. A question for the reader:
What would happen if the release of the guard lock came after the park(),
and not before? Hint: something bad.

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typedef struct __lock_t {
int flag;
int guard;
queue_t *q;
} lock_t;

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void lock_init(lock_t *m) {
m->flag = 0;
m->guard = 0;
queue_init(m->q);
}

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19
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void lock(lock_t *m) {
while (TestAndSet(&m->guard, 1) == 1)
; //acquire guard lock by spinning
if (m->flag == 0) {
m->flag = 1; // lock is acquired
m->guard = 0;
} else {
queue_add(m->q, gettid());
m->guard = 0;
park();
}
}

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void unlock(lock_t *m) {
while (TestAndSet(&m->guard, 1) == 1)
; //acquire guard lock by spinning
if (queue_empty(m->q))
m->flag = 0; // let go of lock; no one wants it
else
unpark(queue_remove(m->q)); // hold lock (for next thread!)
m->guard = 0;
}

Figure 28.8: Lock With Queues, Test-and-set, Yield, And Wakeup
You might also notice the interesting fact that the flag does not get set
back to 0 when another thread gets woken up. Why is this? Well, it is not
an error, but rather a necessity! When a thread is woken up, it will be as
if it is returning from park(); however, it does not hold the guard at that
point in the code and thus cannot even try to set the flag to 1. Thus, we
just pass the lock directly from the thread releasing the lock to the next
thread acquiring it; flag is not set to 0 in-between.
Finally, you might notice the perceived race condition in the solution,
just before the call to park(). With just the wrong timing, a thread will
be about to park, assuming that it should sleep until the lock is no longer
held. A switch at that time to another thread (say, a thread holding the
lock) could lead to trouble, for example, if that thread then released the
lock. The subsequent park by the first thread would then sleep forever
(potentially). This problem is sometimes called the wakeup/waiting race;
to avoid it, we need to do some extra work.
Solaris solves this problem by adding a third system call: setpark().
By calling this routine, a thread can indicate it is about to park. If it then
happens to be interrupted and another thread calls unpark before park is

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actually called, the subsequent park returns immediately instead of sleeping. The code modification, inside of lock(), is quite small:
1
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3

queue_add(m->q, gettid());
setpark(); // new code
m->guard = 0;

A different solution could pass the guard into the kernel. In that case,
the kernel could take precautions to atomically release the lock and dequeue the running thread.

28.15

Different OS, Different Support
We have thus far seen one type of support that an OS can provide in
order to build a more efficient lock in a thread library. Other OS’s provide
similar support; the details vary.
For example, Linux provides something called a futex which is similar to the Solaris interface but provides a bit more in-kernel functionality.
Specifically, each futex has associated with it a specific physical memory location; associated with each such memory location is an in-kernel
queue. Callers can use futex calls (described below) to sleep and wake as
need be.
Specifically, two calls are available. The call to futex wait(address,
expected) puts the calling thread to sleep, assuming the value at address
is equal to expected. If it is not equal, the call returns immediately. The
call to the routine futex wake(address) wakes one thread that is waiting on the queue. The usage of these in Linux is as found in 28.9.
This code snippet from lowlevellock.h in the nptl library (part of
the gnu libc library) [L09] is pretty interesting. Basically, it uses a single
integer to track both whether the lock is held or not (the high bit of the
integer) and the number of waiters on the lock (all the other bits). Thus,
if the lock is negative, it is held (because the high bit is set and that bit
determines the sign of the integer). The code is also interesting because it
shows how to optimize for the common case where there is no contention:
with only one thread acquiring and releasing a lock, very little work is
done (the atomic bit test-and-set to lock and an atomic add to release the
lock). See if you can puzzle through the rest of this “real-world” lock to
see how it works.

28.16

Two-Phase Locks
One final note: the Linux approach has the flavor of an old approach
that has been used on and off for years, going at least as far back to Dahm
Locks in the early 1960’s [M82], and is now referred to as a two-phase
lock. A two-phase lock realizes that spinning can be useful, particularly
if the lock is about to be released. So in the first phase, the lock spins for
a while, hoping that it can acquire the lock.

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void mutex_lock (int *mutex) {
int v;
/* Bit 31 was clear, we got the mutex (this is the fastpath)
if (atomic_bit_test_set (mutex, 31) == 0)
return;
atomic_increment (mutex);
while (1) {
if (atomic_bit_test_set (mutex, 31) == 0) {
atomic_decrement (mutex);
return;
}
/* We have to wait now. First make sure the futex value
we are monitoring is truly negative (i.e. locked). */
v = *mutex;
if (v >= 0)
continue;
futex_wait (mutex, v);
}
}

*/

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void mutex_unlock (int *mutex) {
/* Adding 0x80000000 to the counter results in 0 if and only if
there are not other interested threads */
if (atomic_add_zero (mutex, 0x80000000))
return;

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/* There are other threads waiting for this mutex,
wake one of them up. */
futex_wake (mutex);

Figure 28.9: Linux-based Futex Locks
However, if the lock is not acquired during the first spin phase, a second phase is entered, where the caller is put to sleep, and only woken up
when the lock becomes free later. The Linux lock above is a form of such
a lock, but it only spins once; a generalization of this could spin in a loop
for a fixed amount of time before using futex support to sleep.
Two-phase locks are yet another instance of a hybrid approach, where
combining two good ideas may indeed yield a better one. Of course,
whether it does depends strongly on many things, including the hardware environment, number of threads, and other workload details. As
always, making a single general-purpose lock, good for all possible use
cases, is quite a challenge.

28.17

Summary
The above approach shows how real locks are built these days: some
hardware support (in the form of a more powerful instruction) plus some
operating system support (e.g., in the form of park() and unpark()
primitives on Solaris, or futex on Linux). Of course, the details differ, and
the exact code to perform such locking is usually highly tuned. Check
out the Solaris or Linux open source code bases if you want to see more
details; they are a fascinating read [L09, S09].

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References
[D91] “Just Win, Baby: Al Davis and His Raiders”
Glenn Dickey, Harcourt 1991
There is even an undoubtedly bad book about Al Davis and his famous “just win” quote. Or, we suppose,
the book is more about Al Davis and the Raiders, and maybe not just the quote. Read the book to find
out?
[D68] “Cooperating sequential processes”
Edsger W. Dijkstra, 1968
Available: http://www.cs.utexas.edu/users/EWD/ewd01xx/EWD123.PDF
One of the early seminal papers in the area. Discusses how Dijkstra posed the original concurrency
problem, and Dekker’s solution.
[H93] “MIPS R4000 Microprocessor User’s Manual”.
Joe Heinrich, Prentice-Hall, June 1993
Available: http://cag.csail.mit.edu/raw/
documents/R4400 Uman book Ed2.pdf
[H91] “Wait-free Synchronization”
Maurice Herlihy
ACM Transactions on Programming Languages and Systems (TOPLAS)
Volume 13, Issue 1, January 1991
A landmark paper introducing a different approach to building concurrent data structures. However,
because of the complexity involved, many of these ideas have been slow to gain acceptance in deployed
systems.
[L81] “Observations on the Development of an Operating System”
Hugh Lauer
SOSP ’81
A must-read retrospective about the development of the Pilot OS, an early PC operating system. Fun
and full of insights.
[L09] “glibc 2.9 (include Linux pthreads implementation)”
Available: http://ftp.gnu.org/gnu/glibc/
In particular, take a look at the nptl subdirectory where you will find most of the pthread support in
Linux today.
[M82] “The Architecture of the Burroughs B5000
20 Years Later and Still Ahead of the Times?”
Alastair J.W. Mayer, 1982
www.ajwm.net/amayer/papers/B5000.html
From the paper: “One particularly useful instruction is the RDLK (read-lock). It is an indivisible
operation which reads from and writes into a memory location.” RDLK is thus an early test-and-set
primitive, if not the earliest. Some credit here goes to an engineer named Dave Dahm, who apparently
invented a number of these things for the Burroughs systems, including a form of spin locks (called
“Buzz Locks” as well as a two-phase lock eponymously called “Dahm Locks.”)
[MS91] “Algorithms for Scalable Synchronization on Shared-Memory Multiprocessors”
John M. Mellor-Crummey and M. L. Scott
ACM TOCS, February 1991
An excellent survey on different locking algorithms. However, no OS support is used, just fancy hardware instructions.
[P81] “Myths About the Mutual Exclusion Problem”
G.L. Peterson
Information Processing Letters. 12(3) 1981, 115–116
Peterson’s algorithm introduced here.

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[S05] “Guide to porting from Solaris to Linux on x86”
Ajay Sood, April 29, 2005
Available: http://www.ibm.com/developerworks/linux/library/l-solar/
[S09] “OpenSolaris Thread Library”
Available: http://src.opensolaris.org/source/xref/onnv/onnv-gate/
usr/src/lib/libc/port/threads/synch.c
This is also pretty interesting to look at, though who knows what will happen to it now that Oracle owns
Sun. Thanks to Mike Swift for the pointer to the code.
[W09] “Load-Link, Store-Conditional”
Wikipedia entry on said topic, as of October 22, 2009
http://en.wikipedia.org/wiki/Load-Link/Store-Conditional
Can you believe we referenced wikipedia? Pretty shabby. But, we found the information there first,
and it felt wrong not to cite it. Further, they even listed the instructions for the different architectures: ldl l/stl c and ldq l/stq c (Alpha), lwarx/stwcx (PowerPC), ll/sc (MIPS), and
ldrex/strex (ARM version 6 and above).
[WG00] “The SPARC Architecture Manual: Version 9”
David L. Weaver and Tom Germond, September 2000
SPARC International, San Jose, California
Available: http://www.sparc.org/standards/SPARCV9.pdf
Also see: http://developers.sun.com/solaris/articles/atomic sparc/ for some
more details on Sparc atomic operations.

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