Ch5 process synchronization

OPERATING SYSTEM
Chapter 5: Process Synchronization

Chapter 5: Process Synchronization
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Background
The Critical-Section Problem
Peterson’s Solution
Synchronization Hardware
Mutex Locks
Semaphores
Classic Problems of Synchronization
Monitors
Synchronization Examples
Alternative Approaches

Objectives
• To introduce the critical-section problem, whose solutions can be used to
ensure the consistency of shared data
• To present both software and hardware solutions of the critical-section
problem

• To examine several classical process-synchronization problems
• To explore several tools that are used to solve process synchronization
problems

Background
• Processes can execute concurrently
• May be interrupted at any time, partially completing execution
• Concurrent access to shared data may result in data inconsistency
• Maintaining data consistency requires mechanisms to ensure the orderly
execution of cooperating processes

Background (Cont.)
• Illustration of the problem:
Suppose that we wanted to provide a solution to the consumer-producer
problem that fills all the buffers. We can do so by having an integer counter
that keeps track of the number of full buffers. Initially, counter is set to 0. It
is incremented by the producer after it produces a new buffer and is
decremented by the consumer after it consumes a buffer.

Producer
while (true) {
/* produce an item in next produced */
while (counter == BUFFER SIZE) ;
/* do nothing */
buffer[in] = next produced;
in = (in + 1) % BUFFER SIZE;
counter++;
}

Consumer
while (true) {
while (counter == 0)
; /* do nothing */
next consumed = buffer[out];
out = (out + 1) % BUFFER SIZE;
counter--;
/* consume the item in next consumed */
}

Race Condition
• counter++ could be implemented as
register1 = counter
register1 = register1 + 1
counter = register1
• counter-- could be implemented as
register2 = counter
register2 = register2 - 1
counter = register2

Race Condition (Cont.)
• Consider this execution interleaving with “count = 5” initially:
S0: producer execute register1 = counter
{register1 = 5}
S1: producer execute register1 = register1 + 1 {register1 = 6}
S2: consumer execute register2 = counter
{register2 = 5}
S3: consumer execute register2 = register2 – 1 {register2 = 4}
S4: producer execute counter = register1
{counter = 6 }
S5: consumer execute counter = register2
{counter = 4}

Critical Section Problem
• Consider system of n processes {p0, p1, … pn-1}
• Each process has critical section segment of code
– Process may be changing common variables, updating table, writing file, etc
– When one process in critical section, no other may be in its critical section

• Critical section problem is to design protocol to solve this
• Each process must ask permission to enter critical section in entry section,
may follow critical section with exit section, then remainder section

Critical Section
• General structure of process pi is

Solution to Critical-Section Problem
1. Mutual Exclusion - If process Pi is executing in its critical section, then no other
processes can be executing in their critical sections

2. Progress - If no process is executing in its critical section and there exist some
processes that wish to enter their critical section, then the selection of the processes
that will enter the critical section next cannot be postponed indefinitely
3. Bounded Waiting - A bound must exist on the number of times that other processes
are allowed to enter their critical sections after a process has made a request to enter
its critical section and before that request is granted

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Assume that each process executes at a nonzero speed
No assumption concerning relative speed of the n processes

Solution to Critical-Section Problem (Cont.)
• Two approaches depending on if kernel is preemptive or nonpreemptive

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Preemptive – allows preemption of process when running in kernel mode
Non-preemptive – runs until exits kernel mode, blocks, or voluntarily yields
CPU

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Essentially free of race conditions in kernel mode

Peterson’s Solution
• Good algorithmic description of solving the problem
• Two process solution
• Assume that the load and store instructions are atomic; that is,
cannot be interrupted
• The two processes share two variables:
– int turn;
– Boolean flag[2]

Peterson’s Solution (Cont.)
• The variable turn indicates whose turn it is to enter the critical section
• The flag array is used to indicate if a process is ready to enter the
critical section. flag[i] = true implies that process Pi is ready!

Synchronization Hardware
• Many systems provide hardware support for critical section code
• All solutions below based on idea of locking
– Protecting critical regions via locks

• Uniprocessors – could disable interrupts
– Currently running code would execute without preemption
– Generally too inefficient on multiprocessor systems
• Operating systems using this not broadly scalable

Synchronization Hardware (Cont.)
• Modern machines provide special atomic hardware instructions

– Atomic = non-interruptible
– Either test memory word and set value
– Or swap contents of two memory words

Mutex Locks
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Previous solutions are complicated and generally inaccessible to application
programmers
OS designers build software tools to solve critical section problem
Simplest is mutex lock
Product critical regions with it by first acquire() a lock then release() it
– Boolean variable indicating if lock is available or not

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Calls to acquire() and release() must be atomic
– Usually implemented via hardware atomic instructions

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But this solution requires busy waiting
– This lock therefore called a spinlock

Semaphore
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Synchronization tool that does not require busy waiting
Semaphore S – integer variable
Two standard operations modify S :wait() and signal()
– Originally called P() and V()
Less complicated
Can only be accessed via two indivisible (atomic) operations
wait (S) {
while (S <= 0)
; // busy wait
S--;
}
signal (S) {
S++;
}

Semaphore Usage
• Counting semaphore – integer value can range over an unrestricted
domain
• Binary semaphore – integer value can range only between 0 and 1
– Then a mutex lock

• Can implement a counting semaphore S as a binary semaphore
• Can solve various synchronization problems
• Consider P1 and P2 that require S1 to happen before S2
P1:

P2:

S1;
signal(synch);

wait(synch);
S2;

Semaphore Implementation
• Must guarantee that no two processes can execute wait() and
signal() on the same semaphore at the same time
• Thus, implementation becomes the critical section problem where the wait
and signal code are placed in the critical section
– Could now have busy waiting in critical section implementation
• But implementation code is short
• Little busy waiting if critical section rarely occupied

• Note that applications may spend lots of time in critical sections and
therefore this is not a good solution

Semaphore Implementation
with no Busy waiting
• With each semaphore there is an associated waiting queue
• Each entry in a waiting queue has two data items:
– value (of type integer)
– pointer to next record in the list

• Two operations:
– block – place the process invoking the operation on the appropriate waiting queue
– wakeup – remove one of processes in the waiting queue and place it in the ready
queue

Deadlock and Starvation
• Deadlock – two or more processes are waiting indefinitely for an event that
can be caused by only one of the waiting processes
• Starvation – indefinite blocking
– A process may never be removed from the semaphore queue in which it is
suspended

• Priority Inversion – Scheduling problem when lower-priority process holds a
lock needed by higher-priority process
– Solved via priority-inheritance protocol

Classical Problems of Synchronization
• Classical problems used to test newly-proposed
synchronization schemes
– Bounded-Buffer Problem
– Readers and Writers Problem
– Dining-Philosophers Problem

Bounded-Buffer Problem
• n buffers, each can hold one item
• Semaphore mutex initialized to the value 1
• Semaphore full initialized to the value 0
• Semaphore empty initialized to the value n

Readers-Writers Problem
• A data set is shared among a number of concurrent processes
– Readers – only read the data set; they do not perform any updates
– Writers – can both read and write

• Problem – allow multiple readers to read at the same time
– Only one single writer can access the shared data at the same time

• Several variations of how readers and writers are treated – all involve
priorities

Readers-Writers Problem (Cont.)
• Shared Data
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Data set
Semaphore rw_mutex initialized to 1
Semaphore mutex initialized to 1
Integer read_count initialized to 0

Readers-Writers Problem Variations
• First variation – no reader kept waiting unless writer has permission to use
shared object
• Second variation – once writer is ready, it performs write asap
• Both may have starvation leading to even more variations
• Problem is solved on some systems by kernel providing reader-writer locks

Dining-Philosophers Problem
• Philosophers spend their lives thinking and eating
• Don’t interact with their neighbors, occasionally try to pick up 2 chopsticks
(one at a time) to eat from bowl
– Need both to eat, then release both when done

• In the case of 5 philosophers
– Shared data
• Bowl of rice (data set)
• Semaphore chopstick [5] initialized to 1

Problems with Semaphores
• Incorrect use of semaphore operations:
– signal (mutex) …. wait (mutex)
– wait (mutex) … wait (mutex)
– Omitting of wait (mutex) or signal (mutex) (or both)

• Deadlock and starvation

Monitors
• A high-level abstraction that provides a convenient and effective
mechanism for process synchronization
• Abstract data type, internal variables only accessible by code within the
procedure
• Only one process may be active within the monitor at a time
• But not powerful enough to model some synchronization schemes
monitor monitor-name
{
// shared variable declarations
procedure P1 (…) { …. }
procedure Pn (…) {……}
Initialization code (…) { … }
}
}

Condition Variables
• condition x, y;

• Two operations on a condition variable:
– x.wait () – a process that invokes the operation is suspended until
x.signal ()
– x.signal () – resumes one of processes (if any) that invoked x.wait
()
• If no x.wait () on the variable, then it has no effect on the variable

Monitor with Condition Variables

Condition Variables Choices
• If process P invokes x.signal(), with Q in x.wait() state, what
should happen next?
– If Q is resumed, then P must wait

• Options include
– Signal and wait – P waits until Q leaves monitor or waits for another condition
– Signal and continue – Q waits until P leaves the monitor or waits for another
condition
– Both have pros and cons – language implementer can decide
– Monitors implemented in Concurrent Pascal compromise
• P executing signal immediately leaves the monitor, Q is resumed

– Implemented in other languages including Mesa, C#, Java

Resuming Processes within a Monitor
• If several processes queued on condition x, and x.signal()
executed, which should be resumed?
• FCFS frequently not adequate
• conditional-wait construct of the form x.wait(c)
– Where c is priority number
– Process with lowest number (highest priority) is scheduled next

Synchronization Examples
• Solaris

• Windows XP
• Linux
• Pthreads

Solaris Synchronization
• Implements a variety of locks to support multitasking, multithreading
(including real-time threads), and multiprocessing
• Uses adaptive mutexes for efficiency when protecting data from short code
segments
– Starts as a standard semaphore spin-lock
– If lock held, and by a thread running on another CPU, spins
– If lock held by non-run-state thread, block and sleep waiting for signal of lock
being released

Solaris Synchronization (Cont.)
• Uses condition variables
• Uses readers-writers locks when longer sections of code need access to
data
• Uses turnstiles to order the list of threads waiting to acquire either an
adaptive mutex or reader-writer lock
– Turnstiles are per-lock-holding-thread, not per-object

• Priority-inheritance per-turnstile gives the running thread the highest of the
priorities of the threads in its turnstile

Windows XP Synchronization
• Uses interrupt masks to protect access to global resources on uniprocessor
systems
• Uses spinlocks on multiprocessor systems
– Spinlocking-thread will never be preempted

• Also provides dispatcher objects user-land which may act mutexes,
semaphores, events, and timers
– Events
• An event acts much like a condition variable
– Timers notify one or more thread when time expired
– Dispatcher objects either signaled-state (object available) or non-signaled state
(thread will block)

Linux Synchronization
• Linux:
– Prior to kernel Version 2.6, disables interrupts to implement short critical sections
– Version 2.6 and later, fully preemptive

• Linux provides:
– semaphores
– spinlocks
– reader-writer versions of both

• On single-cpu system, spinlocks replaced by enabling and disabling kernel
preemption

Pthreads Synchronization
• Pthreads API is OS-independent

• It provides:
– mutex locks
– condition variables

• Non-portable extensions include:
– read-write locks
– spinlocks

Atomic Transactions
• System Model
• Log-based Recovery
• Checkpoints
• Concurrent Atomic Transactions

System Model
• Assures that operations happen as a single logical unit of work, in its
entirety, or not at all
• Related to field of database systems
• Challenge is assuring atomicity despite computer system failures
• Transaction - collection of instructions or operations that performs single
logical function
– Here we are concerned with changes to stable storage – disk
– Transaction is series of read and write operations
– Terminated by commit (transaction successful) or abort (transaction failed)
operation
– Aborted transaction must be rolled back to undo any changes it performed

Types of Storage Media
• Volatile storage – information stored here does not survive system crashes
– Example: main memory, cache

• Nonvolatile storage – Information usually survives crashes
– Example: disk and tape

• Stable storage – Information never lost
– Not actually possible, so approximated via replication or RAID to devices with
independent failure modes

Goal is to assure transaction atomicity where failures cause loss of information
on volatile storage

Log-Based Recovery
• Record to stable storage information about all modifications by a
transaction
• Most common is write-ahead logging
– Log on stable storage, each log record describes single transaction write
operation, including
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Transaction name
Data item name
Old value
New value

– <Ti starts> written to log when transaction Ti starts
– <Ti commits> written when Ti commits

• Log entry must reach stable storage before operation on data occurs

Log-Based Recovery Algorithm
• Using the log, system can handle any volatile memory errors
– Undo(Ti) restores value of all data updated by Ti
– Redo(Ti) sets values of all data in transaction Ti to new values

• Undo(Ti) and redo(Ti) must be idempotent
– Multiple executions must have the same result as one execution

• If system fails, restore state of all updated data via log
– If log contains <Ti starts> without <Ti commits>, undo(Ti)
– If log contains <Ti starts> and <Ti commits>, redo(Ti)

Checkpoints
• Log could become long, and recovery could take long
• Checkpoints shorten log and recovery time.
• Checkpoint scheme:
1. Output all log records currently in volatile storage to stable storage
2. Output all modified data from volatile to stable storage
3. Output a log record <checkpoint> to the log on stable storage

• Now recovery only includes Ti, such that Ti started executing before the
most recent checkpoint, and all transactions after Ti All other transactions
already on stable storage

Concurrent Transactions
• Must be equivalent to serial execution – serializability
• Could perform all transactions in critical section
– Inefficient, too restrictive

• Concurrency-control algorithms provide serializability

Serializability
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Consider two data items A and B
Consider Transactions T0 and T1
Execute T0, T1 atomically
Execution sequence called schedule
Atomically executed transaction order called serial schedule
For N transactions, there are N! valid serial schedules

Nonserial Schedule
• Nonserial schedule allows overlapped execute
– Resulting execution not necessarily incorrect

• Consider schedule S, operations Oi, Oj
– Conflict if access same data item, with at least one write

• If Oi, Oj consecutive and operations of different transactions & Oi and Oj
don’t conflict
– Then S’ with swapped order Oj Oi equivalent to S

• If S can become S’ via swapping nonconflicting operations
– S is conflict serializable

Schedule 2: Concurrent Serializable Schedule

Locking Protocol
• Ensure serializability by associating lock with each data item
– Follow locking protocol for access control

• Locks
– Shared – Ti has shared-mode lock (S) on item Q, Ti can read Q but not write Q
– Exclusive – Ti has exclusive-mode lock (X) on Q, Ti can read and write Q

• Require every transaction on item Q acquire appropriate lock

• If lock already held, new request may have to wait
– Similar to readers-writers algorithm

Two-phase Locking Protocol
• Generally ensures conflict serializability
• Each transaction issues lock and unlock requests in two phases
– Growing – obtaining locks
– Shrinking – releasing locks

• Does not prevent deadlock

Timestamp-based Protocols
• Select order among transactions in advance – timestamp-ordering
• Transaction Ti associated with timestamp TS(Ti) before Ti starts
– TS(Ti) < TS(Tj) if Ti entered system before Tj
– TS can be generated from system clock or as logical counter incremented at each
entry of transaction

• Timestamps determine serializability order
– If TS(Ti) < TS(Tj), system must ensure produced schedule equivalent to serial
schedule where Ti appears before Tj

Timestamp-based Protocol Implementation
• Data item Q gets two timestamps
– W-timestamp(Q) – largest timestamp of any transaction that executed write(Q)
successfully
– R-timestamp(Q) – largest timestamp of successful read(Q)
– Updated whenever read(Q) or write(Q) executed

• Timestamp-ordering protocol assures any conflicting read and write
executed in timestamp order
• Suppose Ti executes read(Q)
– If TS(Ti) < W-timestamp(Q), Ti needs to read value of Q that was already overwritten

• read operation rejected and Ti rolled back

– If TS(Ti) ≥ W-timestamp(Q)

• read executed, R-timestamp(Q) set to max(R-timestamp(Q), TS(Ti))

Timestamp-ordering Protocol
• Suppose Ti executes write(Q)
– If TS(Ti) < R-timestamp(Q), value Q produced by Ti was needed previously and Ti
assumed it would never be produced
• Write operation rejected, Ti rolled back

– If TS(Ti) < W-timestamp(Q), Ti attempting to write obsolete value of Q
• Write operation rejected and Ti rolled back

– Otherwise, write executed

• Any rolled back transaction Ti is assigned new timestamp and restarted
• Algorithm ensures conflict serializability and freedom from deadlock

Schedule Possible Under Timestamp Protocol

Ch5 process synchronization

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