Unit 06 dbms

DATABASE MANAGEMENT SYSTEMS

MALLA REDDY ENGG. COLLEGE HYD

II B. Tech CSE II Semester

UNIT-VI PPT SLIDES

Text Books: (1) DBMS by Raghu Ramakrishnan
(2) DBMS by Sudarshan and Korth

INDEX
UNIT-6 PPT SLIDES
S.NO Module as per Lecture PPT
Session planner No Slide NO
--------------------------------------------------------------------------------------------------------------------------------
1. Transaction concept & State L1 L1- 1 to L1- 7
2. Implementation of atomicity and durability L2 L2- 1 to L2- 8
3. Serializability L3 L3- 1 to L3- 8
4. Recoverability L4 L4- 1 to L4- 8
5. Implementation of isolation L5 L5- 1 to L5- 6
6. Lock based protocols L6 L6- 1 to L6 -5
7. Lock based protocols L7 L7- 1 to L7- 10
8. Timestamp based protocols L8 L8- 1 to L8- 6
9. Validation based protocol L9 L9- 1 to L9- 9

Transaction Concept
• A transaction is a unit of program execution that accesses
and possibly updates various data items.
• E.g. transaction to transfer $50 from account A to account
B:
1. read(A)
2. A := A – 50
3. write(A)
4. read(B)
5. B := B + 50
6. write(B)
• Two main issues to deal with:
– Failures of various kinds, such as hardware failures
and system crashes
– Concurrent execution of multiple transactions

Slide No.L1-1

Example of Fund Transfer
• Transaction to transfer $50 from account A to account B:
1. read(A)
2. A := A – 50
3. write(A)
4. read(B)
5. B := B + 50
6. write(B)
• Atomicity requirement
– if the transaction fails after step 3 and before step 6, money will be
“lost” leading to an inconsistent database state
• Failure could be due to software or hardware
– the system should ensure that updates of a partially executed
transaction are not reflected in the database
• Durability requirement — once the user has been notified that the
transaction has completed (i.e., the transfer of the $50 has taken
place), the updates to the database by the transaction must persist
even if there are software or hardware failures.

Slide No.L1-2

Example of Fund Transfer (Cont.)
• Transaction to transfer $50 from account A to account B:
1. read(A)
2. A := A – 50
3. write(A)
4. read(B)
5. B := B + 50
6. write(B)
• Consistency requirement in above example:
– the sum of A and B is unchanged by the execution of the
transaction
• In general, consistency requirements include
• Explicitly specified integrity constraints such as primary
keys and foreign keys
• Implicit integrity constraints
– e.g. sum of balances of all accounts, minus sum of
loan amounts must equal value of cash-in-hand
– A transaction must see a consistent database.
– During transaction execution the database may be temporarily
inconsistent.
– When the transaction completes successfully the database must be
consistent
• Erroneous transaction logic can lead to inconsistency

Slide No.L1-3

Example of Fund Transfer (Cont.)
• Isolation requirement — if between steps 3 and 6, another
transaction T2 is allowed to access the partially updated database,
it will see an inconsistent database (the sum A + B will be less than
it should be).
T1 T2
1. read(A)
2. A := A – 50
3. write(A)
read(A), read(B), print(A+B)
4. read(B)
5. B := B + 50
6. write(B
• Isolation can be ensured trivially by running transactions serially
– that is, one after the other.
• However, executing multiple transactions concurrently has
significant benefits, as we will see later.

Slide No.L1-4

ACID Properties
A transaction is a unit of program execution that accesses and possibly
updates various data items.To preserve the integrity of data the database
system must ensure:
• Atomicity. Either all operations of the transaction are properly
reflected in the database or none are.
• Consistency. Execution of a transaction in isolation preserves the
consistency of the database.
• Isolation. Although multiple transactions may execute
concurrently, each transaction must be unaware of other
concurrently executing transactions. Intermediate transaction
results must be hidden from other concurrently executed
transactions.
– That is, for every pair of transactions Ti and Tj, it appears to Ti
that either Tj, finished execution before Ti started, or Tj started
execution after Ti finished.
• Durability. After a transaction completes successfully, the changes
it has made to the database persist, even if there are system failures.

Slide No.L1-5

Transaction State
• Active – the initial state; the transaction stays in this state
while it is executing
• Partially committed – after the final statement has been
executed.
• Failed -- after the discovery that normal execution can no
longer proceed.
• Aborted – after the transaction has been rolled back and the
database restored to its state prior to the start of the
transaction. Two options after it has been aborted:
– restart the transaction
• can be done only if no internal logical error
– kill the transaction
• Committed – after successful completion.

Slide No.L1-6

Transaction State (Cont.)

Slide No.L1-7

Implementation of Atomicity and Durability
• The recovery-management component of a database system
implements the support for atomicity and durability.
• E.g. the shadow-database scheme:
– all updates are made on a shadow copy of the database
• db_pointer is made to point to the updated shadow
copy after
– the transaction reaches partial commit and
– all updated pages have been flushed to disk.

Slide No.L2-1

Implementation of Atomicity and Durability (Cont.)
• db_pointer always points to the current consistent copy of the
database.
– In case transaction fails, old consistent copy pointed to by
db_pointer can be used, and the shadow copy can be deleted.
• The shadow-database scheme:
– Assumes that only one transaction is active at a time.
– Assumes disks do not fail
– Useful for text editors, but
• extremely inefficient for large databases (why?)
– Variant called shadow paging reduces copying of
data, but is still not practical for large databases
– Does not handle concurrent transactions
• Will study better schemes in Chapter 17.

Slide No.L2-2

Concurrent Executions
• Multiple transactions are allowed to run concurrently in the
system. Advantages are:
– increased processor and disk utilization, leading to
better transaction throughput
• E.g. one transaction can be using the CPU while
another is reading from or writing to the disk
– reduced average response time for transactions: short
transactions need not wait behind long ones.
• Concurrency control schemes – mechanisms to achieve
isolation
– that is, to control the interaction among the concurrent
transactions in order to prevent them from destroying the
consistency of the database
• Will study in Chapter 16, after studying notion
of correctness of concurrent executions.

Slide No.L2-3

Schedules
• Schedule – a sequences of instructions that specify the
chronological order in which instructions of concurrent
transactions are executed
– a schedule for a set of transactions must consist of all
instructions of those transactions
– must preserve the order in which the instructions
appear in each individual transaction.
• A transaction that successfully completes its execution
will have a commit instructions as the last statement
– by default transaction assumed to execute commit
instruction as its last step
• A transaction that fails to successfully complete its
execution will have an abort instruction as the last
statement

Slide No.L2-4

Schedule 1
• Let T1 transfer $50 from A to B, and T2 transfer 10% of the
balance from A to B.
• A serial schedule in which T1 is followed by T2 :

Slide No.L2-5

Schedule 2
• A serial schedule where T2 is followed by T1

Slide No.L2-6

Schedule 3
• Let T1 and T2 be the transactions defined previously. The
following schedule is not a serial schedule, but it is
equivalent to Schedule 1.

In Schedules 1, 2 and 3, the sum A + B is preserved.
Slide No.L2-7

Schedule 4
• The following concurrent schedule does not preserve
the value of (A + B ).

Slide No.L2-8

Serializability
• Basic Assumption – Each transaction preserves database
consistency.
• Thus serial execution of a set of transactions preserves database
consistency.
• A (possibly concurrent) schedule is serializable if it is equivalent to a
serial schedule. Different forms of schedule equivalence give rise to
the notions of:
1. conflict serializability
2. view serializability
• Simplified view of transactions
– We ignore operations other than read and write instructions
– We assume that transactions may perform arbitrary computations
on data in local buffers in between reads and writes.
– Our simplified schedules consist of only read and write
instructions.

Slide No.L3-1

Conflicting Instructions
• Instructions li and lj of transactions Ti and Tj respectively,
conflict if and only if there exists some item Q accessed
by both li and lj, and at least one of these instructions
wrote Q.
1. li = read(Q), lj = read(Q). li and lj don’t conflict.
2. li = read(Q), lj = write(Q). They conflict.
3. li = write(Q), lj = read(Q). They conflict
4. li = write(Q), lj = write(Q). They conflict
• Intuitively, a conflict between li and lj forces a (logical)
temporal order between them.
– If li and lj are consecutive in a schedule and they do
not conflict, their results would remain the same even
if they had been interchanged in the schedule.

Slide No.L3-2

Conflict Serializability
• If a schedule S can be transformed into a schedule S´ by a
series of swaps of non-conflicting instructions, we say that S
and S´ are conflict equivalent.
• We say that a schedule S is conflict serializable if it is
conflict equivalent to a serial schedule

Slide No.L3-3

Conflict Serializability (Cont.)
• Schedule 3 can be transformed into Schedule 6, a serial
schedule where T2 follows T1, by series of swaps of non-
conflicting instructions.
– Therefore Schedule 3 is conflict serializable.

Schedule 3 Schedule 6
Slide No.L3-4

Conflict Serializability (Cont.)
• Example of a schedule that is not conflict serializable:

• We are unable to swap instructions in the above schedule to
obtain either the serial schedule < T3, T4 >, or the serial
schedule < T4, T3 >.

Slide No.L3-5

View Serializability
• Let S and S´ be two schedules with the same set of
transactions. S and S´ are view equivalent if the following
three conditions are met, for each data item Q,
1. If in schedule S, transaction Ti reads the initial value of Q,
then in schedule S’ also transaction Ti must read the
initial value of Q.
2. If in schedule S transaction Ti executes read(Q), and that
value was produced by transaction Tj (if any), then in
schedule S’ also transaction Ti must read the value of Q
that was produced by the same write(Q) operation of
transaction Tj .
3. The transaction (if any) that performs the final write(Q)
operation in schedule S must also perform the final
write(Q) operation in schedule S’.
As can be seen, view equivalence is also based purely on reads
and writes alone.

Slide No.L3-6

View Serializability (Cont.)
• A schedule S is view serializable if it is view equivalent to
a serial schedule.
• Every conflict serializable schedule is also view
serializable.
• Below is a schedule which is view-serializable but not
conflict serializable.

• What serial schedule is above equivalent to?
• Every view serializable schedule that is not conflict
serializable has blind writes.

Slide No.L3-7

Other Notions of Serializability
• The schedule below produces same outcome as the serial
schedule < T1, T5 >, yet is not conflict equivalent or view
equivalent to it.

Determining such equivalence requires analysis of operations
other than read and write.

Slide No.L3-8

Recoverable Schedules
Need to address the effect of transaction failures on concurrently
running transactions.
• Recoverable schedule — if a transaction Tj reads a data
item previously written by a transaction Ti , then the
commit operation of Ti appears before the commit
operation of Tj.
• The following schedule (Schedule 11) is not recoverable if
T9 commits immediately after the read

• If T8 should abort, T9 would have read (and possibly shown to the
user) an inconsistent database state. Hence, database must
ensure that schedules are recoverable.

Slide No.L4-1

Cascading Rollbacks
• Cascading rollback – a single transaction failure leads to
a series of transaction rollbacks. Consider the following
schedule where none of the transactions has yet
committed (so the schedule is recoverable)

If T10 fails, T11 and T12 must also be rolled back.
• Can lead to the undoing of a significant amount of work

Slide No.L4-2

Cascadeless Schedules
• Cascadeless schedules — cascading rollbacks cannot occur;
for each pair of transactions Ti and Tj such that Tj reads a
data item previously written by Ti, the commit operation of Ti
appears before the read operation of Tj.
• Every cascadeless schedule is also recoverable
• It is desirable to restrict the schedules to those that are
cascadeless

Slide No.L4-3

Concurrency Control
• A database must provide a mechanism that will ensure
that all possible schedules are
– either conflict or view serializable, and
– are recoverable and preferably cascadeless
• A policy in which only one transaction can execute at a
time generates serial schedules, but provides a poor degree
of concurrency
– Are serial schedules recoverable/cascadeless?
• Testing a schedule for serializability after it has executed is
a little too late!
• Goal – to develop concurrency control protocols that will
assure serializability.

Slide No.L4-4

Concurrency Control vs. Serializability Tests
• Concurrency-control protocols allow concurrent schedules,
but ensure that the schedules are conflict/view serializable,
and are recoverable and cascadeless .
• Concurrency control protocols generally do not examine the
precedence graph as it is being created
– Instead a protocol imposes a discipline that avoids
nonseralizable schedules.
– We study such protocols in Chapter 16.
• Different concurrency control protocols provide different
tradeoffs between the amount of concurrency they allow and
the amount of overhead that they incur.
• Tests for serializability help us understand why a
concurrency control protocol is correct.

Slide No.L4-5

Weak Levels of Consistency
• Some applications are willing to live with weak levels of
consistency, allowing schedules that are not serializable
– E.g. a read-only transaction that wants to get an
approximate total balance of all accounts
– E.g. database statistics computed for query optimization
can be approximate (why?)
– Such transactions need not be serializable with respect to
other transactions
• Tradeoff accuracy for performance

Slide No.L4-6

Levels of Consistency in SQL-92
• Serializable — default
• Repeatable read — only committed records to be read,
repeated reads of same record must return same value.
However, a transaction may not be serializable – it may find
some records inserted by a transaction but not find others.
• Read committed — only committed records can be read,
but successive reads of record may return different (but
committed) values.
• Read uncommitted — even uncommitted records may be
read.

• Lower degrees of consistency useful for gathering approximate
information about the database
• Warning: some database systems do not ensure serializable schedules
by default
– E.g. Oracle and PostgreSQL by default support a level of
consistency called snapshot isolation (not part of the SQL
standard)
Slide No.L4-7

Transaction Definition in SQL
• Data manipulation language must include a construct for
specifying the set of actions that comprise a transaction.
• In SQL, a transaction begins implicitly.
• A transaction in SQL ends by:
– Commit work commits current transaction and begins a
new one.
– Rollback work causes current transaction to abort.
• In almost all database systems, by default, every SQL
statement also commits implicitly if it executes successfully
– Implicit commit can be turned off by a database directive
• E.g. in JDBC, connection.setAutoCommit(false);

Slide No.L4-8

Implementation of Isolation
• Schedules must be conflict or view serializable, and
recoverable, for the sake of database consistency, and
preferably cascadeless.
• A policy in which only one transaction can execute at a
time generates serial schedules, but provides a poor
degree of concurrency.
• Concurrency-control schemes tradeoff between the
amount of concurrency they allow and the amount of
overhead that they incur.
• Some schemes allow only conflict-serializable schedules to
be generated, while others allow view-serializable
schedules that are not conflict-serializable.

Slide No.L5-1

Figure 15.6

Slide No.L5-2

Testing for Serializability

• Consider some schedule of a set of transactions T1,
T2, ..., Tn
• Precedence graph — a direct graph where the
vertices are the transactions (names).
• We draw an arc from Ti to Tj if the two transaction
conflict, and Ti accessed the data item on which the
conflict arose earlier.
• We may label the arc by the item that was accessed.
• Example 1 x

y

Slide No.L5-3

Example Schedule (Schedule A) + Precedence Graph

T1 T2 T3 T4 T5
read(X)
read(Y)
read(Z)
read(V)
T1 T2
read(W)
read(W)
read(Y)
write(Y)
write(Z)
read(U) T3 T4
read(Y)
write(Y)
read(Z)
write(Z) T5
read(U)
write(U)
Slide No.L5-4

Test for Conflict Serializability
• A schedule is conflict serializable if and only
if its precedence graph is acyclic.
• Cycle-detection algorithms exist which take
order n2 time, where n is the number of
vertices in the graph.
– (Better algorithms take order n + e where
e is the number of edges.)
• If precedence graph is acyclic, the
serializability order can be obtained by a
topological sorting of the graph.
– This is a linear order consistent with the
partial order of the graph.
– For example, a serializability order for
Schedule A would be
T5 → T1 → T 3 → T2 → T4
• Are there others?

Slide No.L5-5

Test for View Serializability
• The precedence graph test for conflict serializability
cannot be used directly to test for view serializability.
– Extension to test for view serializability has cost
exponential in the size of the precedence graph.
• The problem of checking if a schedule is view serializable
falls in the class of NP-complete problems.
– Thus existence of an efficient algorithm is extremely
unlikely.
• However practical algorithms that just check some
sufficient conditions for view serializability can still be
used.

Slide No.L5-6

Lock-Based Protocols
• A lock is a mechanism to control concurrent access to a
data item
• Data items can be locked in two modes :
1. exclusive (X) mode. Data item can be both read as well as

written. X-lock is requested using lock-X instruction.
2. shared (S) mode. Data item can only be read. S-lock is

requested using lock-S instruction.
• Lock requests are made to concurrency-control manager.
Transaction can proceed only after request is granted.

Slide No.L6-1

Lock-Based Protocols (Cont.)
• Lock-compatibility matrix

• A transaction may be granted a lock on an item if the requested
lock is compatible with locks already held on the item by other
transactions
• Any number of transactions can hold shared locks on an item,
– but if any transaction holds an exclusive on the item no other
transaction may hold any lock on the item.
• If a lock cannot be granted, the requesting transaction is made
to wait till all incompatible locks held by other transactions have
been released. The lock is then granted.

Slide No.L6-2

Lock-Based Protocols (Cont.)
• Example of a transaction performing locking:
T2: lock-S(A);
read (A);
unlock(A);
lock-S(B);
read (B);
unlock(B);
display(A+B)
• Locking as above is not sufficient to guarantee serializability
— if A and B get updated in-between the read of A and B,
the displayed sum would be wrong.
• A locking protocol is a set of rules followed by all
transactions while requesting and releasing locks. Locking
protocols restrict the set of possible schedules.

Slide No.L6-3

Pitfalls of Lock-Based Protocols
• Consider the partial schedule

• Neither T3 nor T4 can make progress — executing lock-S(B)
causes T4 to wait for T3 to release its lock on B, while executing
lock-X(A) causes T3 to wait for T4 to release its lock on A.
• Such a situation is called a deadlock.
– To handle a deadlock one of T3 or T4 must be rolled back
and its locks released.

Slide No.L6-4

Pitfalls of Lock-Based Protocols (Cont.)
• The potential for deadlock exists in most locking
protocols. Deadlocks are a necessary evil.
• Starvation is also possible if concurrency control
manager is badly designed. For example:
– A transaction may be waiting for an X-lock on an
item, while a sequence of other transactions request
and are granted an S-lock on the same item.
– The same transaction is repeatedly rolled back due to
deadlocks.
• Concurrency control manager can be designed to prevent
starvation.

Slide No.L6-5

The Two-Phase Locking Protocol
• This is a protocol which ensures conflict-serializable
schedules.
• Phase 1: Growing Phase
– transaction may obtain locks
– transaction may not release locks
• Phase 2: Shrinking Phase
– transaction may release locks
– transaction may not obtain locks
• The protocol assures serializability. It can be proved that the
transactions can be serialized in the order of their lock
points (i.e. the point where a transaction acquired its final
lock).

Slide No.L7-1

The Two-Phase Locking Protocol (Cont.)
• Two-phase locking does not ensure freedom from
deadlocks
• Cascading roll-back is possible under two-phase locking.
To avoid this, follow a modified protocol called strict
two-phase locking. Here a transaction must hold all its
exclusive locks till it commits/aborts.
• Rigorous two-phase locking is even stricter: here all
locks are held till commit/abort. In this protocol
transactions can be serialized in the order in which they
commit.

Slide No.L7-2

The Two-Phase Locking Protocol (Cont.)

• There can be conflict serializable schedules that cannot be
obtained if two-phase locking is used.
• However, in the absence of extra information (e.g., ordering of
access to data), two-phase locking is needed for conflict
serializability in the following sense:
Given a transaction Ti that does not follow two-phase locking,
we can find a transaction Tj that uses two-phase locking, and
a schedule for Ti and Tj that is not conflict serializable.

Slide No.L7-3

Lock Conversions
• Two-phase locking with lock conversions:
– First Phase:
– can acquire a lock-S on item
– can acquire a lock-X on item
– can convert a lock-S to a lock-X (upgrade)
– Second Phase:
– can release a lock-S
– can release a lock-X
– can convert a lock-X to a lock-S (downgrade)
• This protocol assures serializability. But still relies on the
programmer to insert the various locking instructions.

Slide No.L7-4

Automatic Acquisition of Locks
• A transaction Ti issues the standard read/write
instruction, without explicit locking calls.
• The operation read(D) is processed as:
if Ti has a lock on D
then
read(D)
else begin
if necessary wait until no other
transaction has a lock-X on D
grant Ti a lock-S on D;
read(D)
end

Slide No.L7-5

Automatic Acquisition of Locks (Cont.)
• write(D) is processed as:
if Ti has a lock-X on D
then
write(D)
else begin
if necessary wait until no other trans. has any lock
on D,
if Ti has a lock-S on D
then
upgrade lock on D to lock-X
else
grant Ti a lock-X on D
write(D)
end;
• All locks are released after commit or abort

Slide No.L7-6

Implementation of Locking
• A lock manager can be implemented as a separate
process to which transactions send lock and unlock
requests
• The lock manager replies to a lock request by sending a
lock grant messages (or a message asking the
transaction to roll back, in case of a deadlock)
• The requesting transaction waits until its request is
answered
• The lock manager maintains a data-structure called a
lock table to record granted locks and pending requests
• The lock table is usually implemented as an in-memory
hash table indexed on the name of the data item being
locked

Slide No.L7-7

Lock Table
• Black rectangles indicate
granted locks, white ones
indicate waiting requests
• Lock table also records the type
of lock granted or requested
• New request is added to the end
of the queue of requests for the
data item, and granted if it is
compatible with all earlier locks
• Unlock requests result in the
request being deleted, and later
requests are checked to see if
Granted they can now be granted
Waiting
• If transaction aborts, all waiting
or granted requests of the
transaction are deleted
– lock manager may keep a
list of locks held by each
Slide No.L7-8
transaction, to implement
this efficiently

Graph-Based Protocols
• Graph-based protocols are an alternative to two-phase
locking
• Impose a partial ordering → on the set D = {d1, d2 ,..., dh} of all
data items.
– If di → dj then any transaction accessing both di and dj
must access di before accessing dj.
– Implies that the set D may now be viewed as a directed
acyclic graph, called a database graph.
• The tree-protocol is a simple kind of graph protocol.

Slide No.L7-9

Tree Protocol

1. Only exclusive locks are allowed.
2. The first lock by Ti may be on any data item.
Subsequently, a data Q can be locked by Ti only if the
parent of Q is currently locked by Ti.
3. Data items may be unlocked at any time.
4. A data item that has been locked and unlocked by Ti
cannot subsequently be relocked by Ti
Slide No.L7-10

Timestamp-Based Protocols
• Each transaction is issued a timestamp when it enters the
system. If an old transaction Ti has time-stamp TS(Ti), a new
transaction Tj is assigned time-stamp TS(Tj) such that TS(Ti)
<TS(Tj).
• The protocol manages concurrent execution such that the
time-stamps determine the serializability order.
• In order to assure such behavior, the protocol maintains for
each data Q two timestamp values:
– W-timestamp(Q) is the largest time-stamp of any
transaction that executed write(Q) successfully.
– R-timestamp(Q) is the largest time-stamp of any
transaction that executed read(Q) successfully.

Slide No. L8-1

Timestamp-Based Protocols (Cont.)

• The timestamp ordering protocol ensures that any
conflicting read and write operations are executed in
timestamp order.
• Suppose a transaction Ti issues a read(Q)
1. If TS(Ti) ≤ W-timestamp(Q), then Ti needs to read a
value of Q that was already overwritten.
Hence, the read operation is rejected, and Ti is
rolled back.
1. If TS(Ti)≥ W-timestamp(Q), then the read operation is
executed, and R-timestamp(Q) is set to max(R-
timestamp(Q), TS(Ti)).

Slide No. L8-2

Timestamp-Based Protocols (Cont.)

• Suppose that transaction Ti issues write(Q).
1. If TS(Ti) < R-timestamp(Q), then the value of Q that Ti is
producing was needed previously, and the system
assumed that that value would never be produced.
Hence, the write operation is rejected, and Ti is rolled
back.
1. If TS(Ti) < W-timestamp(Q), then Ti is attempting to write
an obsolete value of Q.
Hence, this write operation is rejected, and Ti is rolled
back.
1. Otherwise, the write operation is executed, and W-
timestamp(Q) is set to TS(Ti).

Slide No. L8-3

Example Use of the Protocol
A partial schedule for several data items for transactions with
timestamps 1, 2, 3, 4, 5

T1 T2 T3 T4 T5
read(X)
read(Y)
read(Y)
write(Y)
write(Z)
read(Z)
read(X)
abort
read(X)
write(Z)
abort
write(Y)
write(Z)

Slide No. L8-4

Correctness of Timestamp-Ordering Protocol

• The timestamp-ordering protocol guarantees serializability
since all the arcs in the precedence graph are of the form:

transaction transaction
with smaller with larger
timestamp timestamp

Thus, there will be no cycles in the precedence graph
• Timestamp protocol ensures freedom from deadlock as no
transaction ever waits.
• But the schedule may not be cascade-free, and may not
even be recoverable.

Slide No. L8-5

Thomas’ Write Rule
• Modified version of the timestamp-ordering protocol in which
obsolete write operations may be ignored under certain
circumstances.
• When Ti attempts to write data item Q, if TS(Ti) < W-
timestamp(Q), then Ti is attempting to write an obsolete value
of {Q}.
– Rather than rolling back Ti as the timestamp ordering
protocol would have done, this {write} operation can be
ignored.
• Otherwise this protocol is the same as the timestamp
ordering protocol.
• Thomas' Write Rule allows greater potential concurrency.
– Allows some view-serializable schedules that are not
conflict-serializable.
Slide No. L8-6

Validation-Based Protocol
• Execution of transaction Ti is done in three phases.
1. Read and execution phase: Transaction Ti writes only to

temporary local variables
2. Validation phase: Transaction Ti performs a ``validation
test''
to determine if local variables can be written without
violating
serializability.
3. Write phase: If Ti is validated, the updates are applied to the
database; otherwise, Ti is rolled back.
• The three phases of concurrently executing transactions can
be interleaved, but each transaction must go through the
three phases in that order.
– Assume for simplicity that the validation and write phase
occur together, atomically and serially
• I.e., only one transaction executes validation/write at a
time.
• Also called as optimistic concurrency control since
transaction executes fully in the hope that all will go well
Slide No. L9-1
during validation

Validation-Based Protocol (Cont.)

• Each transaction Ti has 3 timestamps
– Start(Ti) : the time when Ti started its execution
– Validation(Ti): the time when Ti entered its validation
phase
– Finish(Ti) : the time when Ti finished its write phase
• Serializability order is determined by timestamp given at
validation time, to increase concurrency.
– Thus TS(Ti) is given the value of Validation(Ti).
• This protocol is useful and gives greater degree of
concurrency if probability of conflicts is low.
– because the serializability order is not pre-decided, and
– relatively few transactions will have to be rolled back.

Slide No. L9-2

Validation Test for Transaction Tj

• If for all Ti with TS (Ti) < TS (Tj) either one of the following
condition holds:
– finish(Ti) < start(Tj)
– start(Tj) < finish(Ti) < validation(Tj) and the set of data
items written by Ti does not intersect with the set of data
items read by Tj.
then validation succeeds and Tj can be committed.
Otherwise, validation fails and Tj is aborted.
• Justification: Either the first condition is satisfied, and there
is no overlapped execution, or the second condition is
satisfied and
the writes of Tj do not affect reads of Ti since they occur
after Ti has finished its reads.
the writes of Ti do not affect reads of Tj since Tj does not
read any item written by TiL9-3
Slide No. .

Schedule Produced by Validation
• Example of schedule produced using validation

T14 T15
read(B)
read(B)
B:= B-50
read(A)
A:= A+50
read(A)
(validate)
display (A+B)
(validate)
write (B)
write (A)

Slide No. L9-4

Multiple Granularity
• Allow data items to be of various sizes and define a hierarchy
of data granularities, where the small granularities are
nested within larger ones
• Can be represented graphically as a tree (but don't confuse
with tree-locking protocol)
• When a transaction locks a node in the tree explicitly, it
implicitly locks all the node's descendents in the same mode.
• Granularity of locking (level in tree where locking is done):
– fine granularity (lower in tree): high concurrency, high
locking overhead
– coarse granularity (higher in tree): low locking overhead,
low concurrency

Slide No. L9-5

Example of Granularity Hierarchy

The levels, starting from the coarsest (top) level are
– database
– area
– file
– record Slide No. L9-6

Intention Lock Modes
• In addition to S and X lock modes, there are three additional
lock modes with multiple granularity:
– intention-shared (IS): indicates explicit locking at a lower
level of the tree but only with shared locks.
– intention-exclusive (IX): indicates explicit locking at a
lower level with exclusive or shared locks
– shared and intention-exclusive (SIX): the subtree rooted
by that node is locked explicitly in shared mode and
explicit locking is being done at a lower level with
exclusive-mode locks.
• intention locks allow a higher level node to be locked in S or
X mode without having to check all descendent nodes.

Slide No. L9-7

Compatibility Matrix with
Intention Lock Modes

• The compatibility matrix for all lock modes is:

IS IX S S IX X
IS     ×

IX   × × ×

S  ×  × ×

S IX  × × × ×

X × × × × ×

Slide No. L9-8

Multiple Granularity Locking Scheme
• Transaction Ti can lock a node Q, using the following rules:
1. The lock compatibility matrix must be observed.
2. The root of the tree must be locked first, and may be
locked in any mode.
3. A node Q can be locked by Ti in S or IS mode only if the
parent of Q is currently locked by Ti in either IX or IS
mode.
4. A node Q can be locked by Ti in X, SIX, or IX mode only
if the parent of Q is currently locked by Ti in either IX or
SIX mode.
5. Ti can lock a node only if it has not previously unlocked
any node (that is, Ti is two-phase).
6. Ti can unlock a node Q only if none of the children of Q
are currently locked by Ti.
• Observe that locks are acquired in root-to-leaf order,
whereas they are releasedNo. L9-9
Slide in leaf-to-root order.

Unit 06 dbms

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Unit 06 dbms