Compiler Design - Computer Science and Engineering

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CS432F/CSL 728:

Compiler Design July 2004

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S. Arun-Kumar [email protected] Department of Computer Science and Engineering I. I. T. Delhi, Hauz Khas, New Delhi 110 016.

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Contents

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• first The Bigpicture last

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• first Lexical Analysis last

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• first Syntax Analysis last – – – –

Context-free Grammars last first Ambiguity last first Shift-Reduce Parsing last first Parse Trees last first

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• first Semantic Analysis last Full Screen

• first Synthesized Attributes last • first Inherited Attributes last

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Contents . . . Contd

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• first Abstract Syntax Trees last • first Symbol Tables last • first Intermediate Representation last – IR: Properties – Typical Instruction Set – IR: Generation

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• first Runtime Structure last Close

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The Big Picture: 1 stream of characters SCANNER stream of tokens

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PARSER parse tree

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PARSER parse tree SEMANTIC ANALYZER abstract syntax tree

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I.R. CODE GENERATOR intermediate representation

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OPTIMIZER optimized intermediate representation

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OPTIMIZER optimized intermediate representation CODE GENERATOR

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The Big Picture: 7 stream of characters SCANNER

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stream of tokens

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PARSER parse tree ERROR− HANDLER

SEMANTIC ANALYZER abstract syntax tree

SYMBOL TABLE

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MANAGER I.R. CODE GENERATOR intermediate representation

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OPTIMIZER optimized intermediate representation CODE GENERATOR

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The Big Picture: 7

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stream of characters SCANNER

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stream of tokens PARSER parse tree ERROR− HANDLER

SEMANTIC ANALYZER abstract syntax tree

SYMBOL TABLE MANAGER

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I.R. CODE GENERATOR intermediate representation OPTIMIZER optimized intermediate representation

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CODE GENERATOR target code

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Scanner Parser Semantic Analysis Symbol Table IR Optimization Contents

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• Takes a stream of characters and identifies tokens from the lexemes.

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• Eliminates comments and redundant whitepace. Go Back

• Keeps track of line numbers and column numbers and passes them as parameters to the other phases to enable error-reporting to the user.

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• Whitespace:

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• Lexeme: A sequence of non-whitespace characters de-

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A sequence of space, tab, newline, carriage-return, form-feed characters etc. limited by whitespace or special characters (e.g. operators like +, -, *).

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• Examples of lexemes. Go Back

– reserved words, keywords, identifiers etc. – Each comment is usually a single lexeme – preprocessor directives

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• Token: A sequence of characters to be treated as a single unit.

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• Examples of tokens. Reserved words (e.g. begin, end, struct, if etc.) Keywords (integer, true etc.) Operators (+, &&, ++ etc) Identifiers (variable names, procedure names, parameter names) – Literal constants (numeric, string, character constants etc.) – Punctuation marks (:, , etc.)

– – – –

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Scanning: 4 • Identification of tokens is usually done by a Deterministic Finite-state automaton (DFA).

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• The set of tokens of a language is represented by a large regular expression.

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• This regular expression is fed to a lexical-analyser generator such as Lex, Flex or ML-Lex.

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• A giant DFA is created by the Lexical analyser generator.

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a−e,g−z identifier

if

f 2

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identifier 0−9,a−z 4

h, j− a− other

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white space

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0−9

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0−9,a−z

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0−9 Go Back

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7 int any char

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12 comment

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Syntax Analysis

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Consider the following two languages over an alphabet A = {a, b}.

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R = {anbn|n < 100} P = {anbn|n > 0} • R may be finitely represented by a regular expression

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(even though the actual expression is very long).

• However, P cannot actually be represented by a regular

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• A regular expression is not powerful enough to represent languages which require parenthesis matching to arbitrary depths.

• All high level programming languages require an underlying language of expressions which require parentheses to be nested and matched to arbitrary depth.

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CF-Grammars: Definition A context-free grammar (CFG) G = hN, T , P, Si consists of

• a set N of nonterminal symbols,

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• a set T of terminal symbols or the alphabet, • a set P of productions or rewrite rules, • each production is of the form X −→ α, where – X ∈ N is a nonterminal and – α ∈ (N ∪ T )∗ is a string of terminals and nonterminals

• a start symbol S ∈ N .

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CFG: Example G = h{S}, {a, b}, P, Si, where S −→ ab and S −→ aSb are the only productions in P . Derivations look like this:

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S ⇒ ab Page 21 of 100

S ⇒ aSb ⇒ aabb S ⇒ aSb ⇒ aaSbb ⇒ aaabbb

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L(G), the language generated by G is {anbn|n > 0}. Full Screen

Actually can be proved by induction on the length and structure of derivations.

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CFG: Empty word

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G = h{S}, {a, b}, P, Si, where S −→ SS | aSb | ε generates all sequences of matching nested parentheses, including the empty word ε.

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A leftmost derivation might look like this:

S ⇒ SS ⇒ SSS ⇒ SS ⇒ aSbS ⇒ abS ⇒ abaSb . . . Page 22 of 100

A rightmost derivation might look like this:

S ⇒ SS ⇒ SSS ⇒ SS ⇒ SaSb ⇒ Sab ⇒ aSbab . . . Other derivations might look like God alone knows what!

S ⇒ SS ⇒ SSS ⇒ SS ⇒ . . .

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CFG: Derivation trees 1 Derivation sequences

• put an artificial order in which productions are fired.

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• instead look at trees of derivations in which we may think of productions as being fired in parallel.

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• There is then no highlighting in red to determine which copy of a nonterminal was used to get the next member of the sequence.

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• Of course, generation of the empty word ε must be shown explicitly in the tree.

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CFG: Derivation trees 2

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S S

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Derivation tree of abaabb

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Ambiguity: 1 Home Page

E → I | C | E+E | E∗E I → y | z C → 4

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Consider the sentence y + 4 ∗ z.

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E → I | C | E+E | E∗E I → y | z C → 4

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E → I | C | E+E | E∗E I → y | z C → 4

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Parsing: 0 Home Page

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Principle:

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Reduce whenever possible. Shift only when reduce is impossible

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Parsing: 8a Home Page

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Parsing: 14b Home Page

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Parsing: 8b r1. E

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Parsing: Summary: 1

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• All high-level languages are designed so that they may

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be parsed in this fashion with only a single token lookahead.

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• Parsers for a language can be automatically constructed by parger-generators such as Yacc, Bison, MLYacc.

• Shift-reduce conflicts if any, are automatically detected

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and reported by the parser-generator.

• Shift-reduce conflicts may be avoided by suitably

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Parsing: Summary: 2

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• Very often shift-reduce conflicts may occur because of the prefix problem. In such cases many parsergenerators resolve the conflict in favour of shifting.

• There is also a possiblility of reduce-reduce conflicts. This usually happens when there is more than one nonterminal symbol to which the contents of the stack may reduce.

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• A minor reworking of the grammar to avoid redundant non-terminal symbols will get rid of reduce-reduce conflicts. The Big Picture

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Semantic Analysis: 1 • Every Programming langauge can be used to program any computable function, assuming of course, it has – unbounded memory, and – unbounded time

• The parser of a programming language provides the framework within which the target code is to be generated.

• The parser also provides a structuring mechanism that divides the task of code generation into bits and pieces determined by the individual nonterminals and production rules.

• However, contex-free grammars are not powerful enough to represent all computable functions. Example, the language {an bn cn |n > 0}.

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Semantic Analysis: 2

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• There are context-sensitive aspects of a program that cannot be represented/enforced by a context-free grammar definition. Examples include – correspondence between formal and actual parameters – type consistency between declaration and use. – scope and visibility issues with respect to identifiers in a program.

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An Attribute Grammar

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E0 → E1−T B E0.val := sub(E1.val, T.val) E → T

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Inherited Attributes: 0 C -style declarations generating int x, y, z. D → TL L → L,I | I

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T → int | float I → x | y | z D

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Inherited Attributes: 5

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An Attribute Grammar

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int

D → TL

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int

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→ int

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→ float B T.type := float.f loat

B T.type := int.int

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B L.in := T.type

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Abstract Syntax: 0 Home Page

E → E−T | T T → T /F | F F → n | (E)

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Suppose we want to evaluate an expression (4 − 1)/2. What we actually want is a tree that looks like this:

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But what we actually get during parsing is a tree that looks like . . .

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Abstract Syntax: 1 . . . THIS!

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Abstract Syntax: 2

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We use attribute grammar rules to construct the abstract syntax tree (AST)!. But in order to do that we first require two procedures for tree construction. makeLeaf(literal) : Creates a node with label literal and returns a pointer to it. makeBinaryNode(opr, opd1, opd2) : Creates a node with label opr (with fields which point to opd1 and opd2) and returns a pointer to the newly created node. Now we may associate a synthesized attribute called ptr with each terminal and nonterminal symbol which points to the root of the subtree created for it.

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Abstract Syntax: 3

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E0 → E1−T B E0.ptr := makeBinaryN ode(−, E1.ptr, T.ptr) E → T T0 → T1/F

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B E.ptr := T.ptr B T0.ptr := makeBinaryN ode(/, T1.ptr, F.ptr) Page 115 of 100

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Symbol Table:1 • The store house of context-sensitive and run-time information about every identifier in the source program.

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• All accesses relating to an identifier require to first find the attributes of the identifier from the symbol table

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• Usually organized as a hash table – provides fast access.

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• Compiler-generated temporaries may also be stored in the symbol table

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Symbol Table:2

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Attributes stored in a symbol table for each identifier:

• type • size

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• scope/visibility information Page 117 of 100

• base address • addresses to location of auxiliary symbol tables (in case

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of records, procedures, classes)

• address of the location containing the string which actually names the identifier and its length in the string pool

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Symbol Table:3 • A symbol table exists through out the compilation and

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run-time.

• Major operations required of a symbol table: – insertion – search – deletions are purely logical (depending on scope and visibility) and not physical

• Keywords are often stored in the symbol table before

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Symbol Table:4 Accesses to the symbol table at every stage of the compilation process,

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Scanning: Insertion of new identifiers. Parsing: Access to the symbol table to ensure that an operand exists (declaration before use).

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Semantic analysis: • Determination of types of identifiers from declarations • type checking to ensure that operands are used in type-valid contexts. • Checking scope, visibility violations.

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Symbol Table:5 IR generation: . Memory allocation and relativea address calculation.

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Optimization: All memory accesses through symbol table Target code: Translation of relative addresses to absolute addresses in terms of word length, word boundary etc.

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i.e.relative to a base address that is known only at run-time

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Intermediate Representation

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Intermediate representations are important for reasons of portability.

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• (more or less) independent of specific features of the

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high-level language. Example. Java byte-code for any high-level language.

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• (more or less) independent of specific features of any particular target architecture (e.g. number of registers, memory size) – number of registers – memory size – word length

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IR Properties: 1 1. It is fairly low-level containing instructions common to all target architectures and assembly languages. How low can you stoop? . . . 2. It contains some fairly high-level instructions that are common to most high-level programming languages. How high can you rise? 3. To ensure portability

• an unbounded number of variables and memory locations • no commitment to Representational Issues

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• memory locations are also typed according to the data they may contain, • no commitment is made regarding word boundaries, and the structure of individual data items.

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IR: Representation? • No commitment to word boundaries or byte boundaries • No commitment to representation of – – – –

int vs. float, float vs. double, packed vs. unpacked, strings – where and how?.

Back to IR Properties:1

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IR: How low can you stoop?

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• most arithmetic and logical operations, load and store instructions etc.

• so as to be interpreted easily, Page 124 of 100

• the interpreter is fairly small, • execution speeds are high,

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• to have fixed length instructions (where each operand position has a specific meaning). Back to IR Properties:1

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IR: How high can you rise? • typed variables, • temporary variables instead of registers.

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• array-indexing, • random access to record fields, • parameter-passing,

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• pointers and pointer management • no limits on memory addresses Back to IR Properties:1

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A typical instruction set: 1 Three address code: A suite of instructions. Each instruction has at most 3 operands.

• an opcode representing an operation with at most 2

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operands Page 126 of 100

• two operands on which the binary operation is performed

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• a target operand, which accumulates the result of the (binary) operation.

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If an operation requires less than 3 operands then one or more of the operands is made null.

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A typical instruction set: 2 • Assignments (LOAD-STORE) • Jumps (conditional and unconditional)

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• Procedures and parameters Go Back

• Arrays and array-indexing • Pointer Referencing and Dereferencing

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A typical instruction set: 2

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• Assignments (LOAD-STORE) – x := y bop z, where bop is a binary operation – x := uop y, where uop is a unary operation – x := y, load, store, copy or register transfer

• Jumps (conditional and unconditional)

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• Assignments (LOAD-STORE) • Jumps (conditional and unconditional) – goto L – Unconditional jump, – x relop y goto L – Conditional jump, where relop is a relational operator

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• Procedures and parameters • Arrays and array-indexing

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• Pointer Referencing and Dereferencing

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– call p n, where n is the number of parameters – return y, return value from a procedures call – param x, parameter declaration

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A typical instruction set: 2 • Assignments (LOAD-STORE)

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• Jumps (conditional and unconditional) • Procedures and parameters • Arrays and array-indexing

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– x := a[i] – array indexing for r-value – a[j] := y – array indexing for l-value Note: The two opcodes are different depending on whether l-value or r-value is desired. x and y are always simple variables

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• Assignments (LOAD-STORE) • Jumps (conditional and unconditional) • Procedures and parameters

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• Arrays and array-indexing • Pointer Referencing and Dereferencing – x := ˆy – referencing: set x to point to y – x := *y – dereferencing: copy contents of location pointed to by y into x – *x := y – dereferencing: copy r-value of y into the location pointed to by x

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Pointers y

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@z

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IR: Generation

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• Can be generated by recursive traversal of the abstract syntax tree.

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• Can be generated by syntax-directed translation as follows: For every non-terminal symbol N in the grammar of the source language there exist two attributes N.place , which denotes the address of a temporary variable where the result of the execution of the generated code is stored N.code , which is the actual code segment generated.

• In addition a global counter for the instructions gener-

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ated is maintained as part of the generation process.

• It is independent of the source language but can express target machine operations without committing to too much detail.

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IR: Infrastructure

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Given an abstract syntax tree T, with T also denoting its root node. T.place address of temporary variable where result of execution of the T is stored. newtemp returns a fresh variable name and also installs it in the symbol table along with relevant information T.code the actual sequence of instructions generated for the tree T. newlabel returns a label to mark an instruction in the generated code which may be the target of a jump. emit emits an instructions (regarded as a string).

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IR: Infrastructure Colour and font coding of IR code generation

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• Green: Nodes of the Abstract Syntax Tree • Brown: Characters and strings of the Intermediate Representation

• Red: Variables and data structures of the language in which the IR code generator is written

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• blue: Names of relevant procedures used in IR code generation.

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IR: Example

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E

→ id

B

E.place := id.place; E.code := emit()

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E0

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E0.place := newtemp; E0.code := E1.code || E2.code || emit(E0.place := E1.place − E2.place)

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IR: Example Home Page

S

→ id := E

S.code

:= E.code || emit(id.place:=E.place)

S0

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→ while E do S1

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S0.begin := newlabel; S0.af ter := newlabel; S0.code := emit(S0.begin:) || E.code || emit(if E.place= 0 goto S0.af ter) || S1.code || emit(gotoS0.begin) || emit(S0.af ter:)

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IR: Example Home Page

S

→ id := E

S.code

:= E.code || emit(id.place:=E.place)

S0

B

→ while E do S1

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B Page 139 of 100

S0.begin := newlabel; S0.af ter := newlabel; S0.code := emit(S0.begin:) || E.code || emit(if E.place= 0 goto S0.af ter) || S1.code || emit(gotoS0.begin) || emit(S0.af ter:)

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IR: Generation

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While generating the intermediate representation, it is sometimes necessary to generate jumps into code that has not been generated as yet (hence the address of the label is unknown). This usually happens while processing

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• forward jumps • short-circuit evaluation of boolean expressions It is usual in such circumstances to either fill up the empty label entries in a second pass over the the code or through a process of backpatching (which is the maintenance of lists of jumps to the same instruction number), wherein the blank entries are filled in once the sequence number of the target instruction becomes known.

Page 140 of 100

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A Calling Chain Main program

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Globals Procedure P2 Locals of P2

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Procedure P21 Locals of P21 Body of P21

Call P21 Page 141 of 100

Body of P2

Call P21 Go Back

Procedure P1 Locals of P1

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Body of P1

Call P2 Main body

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Run-time Structure: 1

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Main program Title Page

Globals Procedure P2

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Locals of P2 Procedure P21 Locals of P21 Page 142 of 100

Body of P21 Body of P2

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Procedure P1 Full Screen

Locals of P1 Body of P1

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Locals of P2 Procedure P21 Locals of P21 Page 143 of 100

Body of P21 Body of P2 Procedure P1 Locals of P1 Body of P1

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Return address to Main Dynamic link to Main Locals of P1 Static link to Main Formal par of P1

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Locals of P2 Procedure P21 Locals of P21


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Return address to last of P1 Dynamic link to last P1 Locals of P2 Static link to last P1 Formal par P2

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Globals Procedure P2 Locals of P2 Procedure P21 Locals of P21


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Return address to last of P2 Dynamic link to last P2 Locals of P21 Static link last P2 Formal par P21 Return address to last of P1 Dynamic link to last P1 Locals of P2 Static link to last P1 Formal par P2

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Page 145 of 100

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Run-time Structure: 5 Main program Globals

Return address to last of P21 Dynamic link to last P21 Locals of P21 Static link to last P2 Formal par P21

Procedure P2 Locals of P2 Procedure P21 Locals of P21


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Return address to last of P2 Dynamic link to last P2 Locals of P21 Static link last P2 Formal par P21 Return address to last of P1 Dynamic link to last P1 Locals of P2 Static link to last P1 Formal par P2

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Compiler Design - Computer Science and Engineering

Compiler Design - Computer Science and Engineering

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