CS5205Extended Type System1 CS5205: Foundation in Programming Languages Lecture 7 : Extended Type...

CS5205 Extended Type System 1

CS5205: Foundation in Programming Languages

Lecture 7 : Extended Type System

• Base & Unit Types• Sequencing• Ascription• Let bindings• Pairs, Tuples, Records• Sums, variants• Recursion • References• Exceptions

Pure Simply Typed Lambda CalculusPure Simply Typed Lambda Calculus

• t ::= termsx variable x:T.t abstractiont t application

• v ::= value x :T.t abstraction value

• T ::= typesT ! T type of functions

• ::= contexts empty context

, x:T type variable binding

Call-by-Value SemanticsCall-by-Value Semantics

( x.t) v ! [x v] t (E-AppAbs)

t1 ! t‘1

t1 t2 ! t’1 t2

(E-App1)

t2 ! t‘2

v1 t2 ! v1 t’2

(E-App2)

TypingTyping

x:T 2 ` x : T

` t1 : T1 ! T2 ` t2 : T1

` t1 t2 : T2

(T-Var)

(T-App)

, x:T1 ` t2 : T2

` x:T1.t2 : T1 ! T2 (T-Abs)

Where are the Base Types?Where are the Base Types?

• T ::= typesT ! T type of functions

Extend with uninterpreted base types, e.g.• T ::= types

T ! T type of functionsA base type 1B base type 2C base type 3 :

Examples Examples

x:A. x

f:A! B. x:A. f(x)

f:A! A. x:A. f(x)

f:B! B. x:B. f(f(x))

Base types are Bool, Nat, etc.The above types are monomorphic rather than polymorphic. For polymorphic types, we need to use type variables, denoted by say , , etc.

Unit TypeUnit Type

New Syntax:t ::= … terms

unit constant unitv ::= … values

unit constant unitT ::= … types

Unit unit typeNote that Unit type has only one possible value.

New Evaluation Rules: None

New Typing Rules : ` unit : Unit T-Unit

Sequencing : Basic Idea Sequencing : Basic Idea

Syntax : e1; e2

Evaluate an expression (to achieve some side-effect, such as printing), ignore its result and then evaluate another expression.

Examples:

(print x); x+1

(printcurrenttime); compute; (printcurrenttime)

Lambda Calculus with SequencingLambda Calculus with SequencingNew Syntax• t ::= … terms

t ; t sequence

unit ; t ! t (E-SeqUnit)

t1 ; t2 ! t‘1 ; t2

t1 ! t‘1 (E-Seq)

New Evaluation Rules:

Sequencing (cont) Sequencing (cont)

` t1 : Unit1 ` t2 : T2

` t1 ; t2 : T2

(T-Seq)

New Typing Rule:

Sequencing (Second Version)Sequencing (Second Version)

• Treat t1;t2 as an abbreviation for (x:Unit. t2) t1 .

• Then the evaluation and typing rules for abstraction and application will take care of sequencing!

• Such shortcuts are called derived forms (or syntactic sugar) and are heavily used in programming language definition.

Equivalence of Sequencing ExtensionEquivalence of Sequencing Extension

Let E be the simply typed lambda calculus with the Unit type and the sequencing construct.

Let I be the simply-typed lambda calculus with Unit only.

Let e 2 E ! I be the elaboration function that translates from E To I.

Then, we have for each term t:

• t !E t’ iff e(t) !I e(t’)

`E t:T iff `I e(t):T

Ascription : Motivation Ascription : Motivation

Sometimes, we want to say explicitly that a term has a certain type.

Reasons:• as comments for online documentation• for debugging type errors• control printing of types (together with type

syntax)• casting (as in Java)• resolve ambiguity (of more complex typing)

Ascription : Syntax Ascription : Syntax

New Syntax• t ::= … terms

t as T ascription

Example:

(f (g (h x y z))) as Bool

Ascription (cont)Ascription (cont)New Evaluation Rules:

New Typing Rules:

v as T ! v (E-Ascribe1)

t ! t‘

t as T ! t‘ as T (E- Ascribe2)

` t : T ` t as T : T

(T-Ascribe)

Let Bindings : Motivation Let Bindings : Motivation

• Let expression allow us to give a name to the result of an expression for later use and reuse.

• Examples:

let pi=<long computation> in …pi..pi..pi….

let square = x:Nat. x*x in ….(square 2)..(square 4)…

Lambda Calculus with Let BindingLambda Calculus with Let BindingNew Syntax

t ::= … termslet x=t in t let binding

let x=v in t2! [x v] t2

t1 ! t‘1

let x=t1 in t2 ! let x=t‘1 in t2

(E- Let2)

New Evaluation Rules

(E-Let1)

Let Binding (cont) Let Binding (cont)

` t1 : T1 , x:T1 ` t2 : T2

` let x=t1 in t2 : T2

(T-Let)

New Typing Rule:

Let Bindings as Derived FormLet Bindings as Derived Form

We can consider let expressions as derived form:

In untyped setting:let x=t1 in t2 abbreviates to ( x. t2) t1

In a typed setting:let x=t1 in t2 abbreviates to ( x:?. t2) t1

How to get type declaration for the formal parameter? Answer : Type inference (see next lecture).

Pairs : Motivation Pairs : Motivation

Pairs provide the simplest kind of data structures.

Examples: (9, 81)

x : Nat. (x, x*x)

Pairs : SyntaxPairs : Syntax

• t ::= … terms(t, t) variablet.1 first projectiont.2 second projection

• v ::= … value(v, v) pair value

• T ::= … typesT £ T product type

Pairs : Evaluation RulesPairs : Evaluation Rules

(v1, v2).1 ! v1

t ! t‘

t.1 ! t‘.1(E-Proj1)

(E-PairBeta1)

(v1, v2).2 ! v2 (E-PairBeta2)

t ! t‘

t.2 ! t‘.2(E-Proj2)

Pairs : Evaluation Rules (cont)Pairs : Evaluation Rules (cont)

t ! t‘

(v, t) ! (v, t‘)(E-Pair2)

t1 ! t‘1

(t1, t2) ! (t‘1, t2)

(E-Pair1)

Pairs : Typing RulesPairs : Typing Rules

` t1 : T1 ` t2 : T2

` (t1,t2) : T1 £ T2 (T-Pair)

` t : T1 £ T2

` t.1 : T1 (T-Proj1)

` t : T1 £ T2

` t.2 : T2 (T-Proj2)

Tuples Tuples

Tuples are a straightforward generalization of pairs, where n terms are combined in a tuple expression.

Example:(1, true, unit) : (Nat, Bool, Unit)(1,(true, 0)) : (Nat, (Bool, Nat))( ) : ( )

Note that n may be 0. Then the only value is ( ) with type ( ). Such a type is isomorphic to Unit.

Records Records

Sometimes it is better to have components labeled more meaningfully instead of via numbers 1..n, as in tuples

Tuples with labels are called records.

Example:{partno=5524, cost=30.27, instock =false}

has type {partno:Nat, cost:Float, instock:Bool}instead of:

(5524,30.27,false) : (Nat, Float, Bool)

Tuples as Records Tuples as Records

We can treat tuples as an abbreviation for records whose fields are numbers from 1 to n.

Example:

(1, true, unit) is abbreviation for {1=1, 2=true, 3=unit}

Pattern-Matching : Motivation Pattern-Matching : Motivation

It is often convenient to access some or all components of a record simultaneously as in:

let {partno=p, cost=c, instock=i} = accessdb key1 key2

in …p…c…i…

let construct can be extended with pattern-matching. Usually case construct also used for sum/variant type to be described next.

Pattern-Matching Pattern-Matching

New syntactic Form:

p ::= x variable pattern{li=pi}i 2 1..n record pattern

t ::= … termslet p=t in t pattern binding

Pattern-Matching (cont) Pattern-Matching (cont) Matching Rules :

match(x,v) = [x v]

for each i match(pi, vi) = i

match({li, pi}i 2 1..n, {li, vi}i 2 1..n)

= 1 ± … ± n

(M-Rcd)

(M-Var)

Evaluation Rules :

t1 ! t‘1

let p=t1 in t2! let p=t1‘ in t2

(E-Let)

let p=v in t ! match(p,v) t (E-LetV)

Sums : Motivation Sums : Motivation

Often, like to handle values of different structures with the same function. Examples:

PhysicalAddr={firstlast:String, add:String}VirtualAddr={name:String, email:String}

A sum type can then be written like this:

Addr = PhysicalAddr + VirtualAddr

Given a sum type; e.g.

K = Nat + Bool

Need to use tags inl and inr to indicate that a value is a particular member of the sum type; e.g.

inl 5 : K but not inr 5 : K nor 5 : K inr true : K

Given the address type:

Addr = PhysicalAddr + VirtualAddr

We can use case construct to analyse the particular value of sum type:

getName = a : Addr. case a ofinl x => a.firstlastinr y => y.name

Problem : Uniqueness of Types Problem : Uniqueness of Types

Given the address type:

PhysicalAddr={firstlast:String, add:String}VirtualAddr={name:String, email:String}PagingUser={name:String, pager:String}Addr = PhysicalAddr + VirtualAddrUserCoord = PagerUser + VirtualAddr

What is the type of :

inr {name : “chin”, email=“chinwn@comp.nus.edu.sg”} ?

Resolving Ambiguous Types Resolving Ambiguous Types

• Analyse the program to find out what type should be. (Type inference)

• Given a type that includes all possible sum types. (Subtyping)

• Force programmer to explicitly declare the type using ascription.

Sums : SyntaxSums : Syntax

• t ::= … termsinl t as T tagging (left)inr t as T tagging (right)case t of {pi => ti} pattern matching

• v ::= … valueinl v as T tagged value (left)inr v as T tagged value (right)

• T ::= … typesT + T sum type

Sums : Evaluation RulesSums : Evaluation Rules

case (inl v as T) of {inl x1 ) t1 | inr x2 ) t2}

! [x1 v] t1

(E-CaseInl)

case (inr v as T) of {inl x1 ) t1 | inr x2 ) t2}

! [x2 v] t2

(E-CaseInr)

t ! t‘

case t of {p ) t} ! case t‘ of {p ) t} (E-Case)

Sums : Evaluation RulesSums : Evaluation Rules

t ! t‘

inl t as T ! inl t‘ as T (E-Inl)

t ! t‘

inr t as T ! inl t‘ as T (E-Inr)

Sums : Typing RulesSums : Typing Rules

` t1 : T1

` inl t1 as T1 £ T2 : T1 £ T2 (T-Inl)

` t2 : T2

` inr t2 as T1 £ T2 : T1 £ T2 (T-Inr)

Variants : Labeled Sums Variants : Labeled Sums

Instead of inl and inr, we may like to use nicer labels, just as in records. This is what variants do.

For types, instead of: T1 + T we write: <l1:T1 + l2:T2>

For terms, instead of: inl r as T1 + T we write: <l1= t> as <l1:T1 + l2:T2>

Variants : Example Variants : Example

An example using variant type:Addr =<physical : PhysicalAddr, virtual : VirtualAddr>

A variant value:a = <physical=pa> as Addr

Function over variant value:getName = a : Addr.

case a of<physical=x> ) x.firstlast<virtual=y> ) y.name

Application : Options Application : Options

Option type denotes a maybe value:OptionNat =<none : Unit, some : Nat>Table = Nat ! OptionNat

Operations on Table:emptyTable = n : Nat. <none=unit> as OptionNat

extendTable = t:Table. m:Nat v: Nat. n:Nat.

if equal n m then <some=v> as OptionNatelse t n

Application : Enumeration Application : Enumeration

Enumeration are variants that only make use of their labels. Each possible value is unit.

Weekday =<monday:Unit, tuesday:Unit, wednesday:Unit, thursday:Unit,

friday:Unit >

Application : Single-Field Variants Application : Single-Field Variants

Labels can be convenient to add more information on how the value is used.

Example, currency denominations (or units):

DollarAmount =<dollars:Float>

EuroAmount =<euros:Float>

Recursion : Motivation Recursion : Motivation

Recall the fix-point combinator:

fix = f. ( x. f ( y. x x y)) ( x. f ( y. x x y))

Unfortunately, it is not valid (well-typed) in simply typed lambda calculus.

Solution : provide this as a language primitive that is hardwired.

Recursion : Syntax & Evaluation RulesRecursion : Syntax & Evaluation Rules

New Syntax• t ::= … terms

fix t fixed point operator

New Evaluation

t ! t‘

fix t ! fix t’ (E-Fix)

fix ( x : T. t) ! [x fix ( x : T. t)] t (E-FixBeta)

Recursion : Typing RulesRecursion : Typing Rules

` t : T ! T

` fix t : T (T-Fix)

Can you guess the inherent type of fix?

References : Motivation References : Motivation

Many languages do not syntactically distinguish between references (pointers) from their values.

In C, we write: x = x+1

For typing, it is useful to make this distinction explicit; since operationally pointers and values are different.

Introduce the syntax (ref t), which returns a reference to the result of evaluating t. The type of ref t is Ref T, if T is the type of t.

Remember that we have many type constructors already:

Nat £ float {partno:Nat,cost:float}Unit+Nat <none:Unit,some:Nat>

Typing : First AttemptTyping : First Attempt ` t : T

` ref t : Ref T (T-Ref)

` t : Ref T ` ! t : T

(T-Deref)

` t1 : Ref T1 ` t2 : T1

` t1 := t2 : Unit(T-Assign)

What should be the value of a reference?

What should the assignment “do”?

How can we capture the difference of evaluating a dereferencing depending on the value of the reference?

How do we capture side-effects of assignment?

Answer:

Introduce locations corresponding to references.

Introduce stores that map references to values.

Extend evaluation relation to work on stores.

References : Evaluation References : Evaluation

Instead of:

t ! t’

we now write:

t | !t’ | ’

where ’ denotes the changed store.

Evaluation of Application Evaluation of Application

t1 | !t’1 | ’

t1 t2 | ! t’1 t2 | ’ (E-Appl1)

( x : T.t) v | ! [x v] t | (E-AppAbs)

t2 | !t’2 | ’

v t2 | ! v t’2 | ’ (E-Appl2)

ValuesValues

The result of evaluating a ref expression is a location

• v ::= value x:T.t abstraction valueunit unit valuel store location

TermsTerms

Below is the syntax for terms. • t ::= terms

x variable x:T.t abstraction valueunit constant unitt t applicationref t reference creation! t dereferencet := t assignmentl store location

Evaluation of Deferencing Evaluation of Deferencing

t | !t’ | ’! t | ! ! t’ | ’

(E-Appl1)

lv ! l | ! v |

(E-DeRefLoc)

Evaluation of Assignment Evaluation of Assignment

t1 | !t’1 | ’

t1 := t2 | ! t’1 := t2 | ’ (E-Assign1)

(E-Assign)

t2 | !t’2 | ’

l := t2 | ! l := t’2 | ’ (E-Assign2)

l:=v2 | !unit | l v]

Evaluation of References Evaluation of References

t | !t’ | ’ref t | ! ref t’ | ’

(E-Ref)

l dom()

ref v | ! l | , (l v)(E-RefV)

Towards a Typing for Locations Towards a Typing for Locations

` (l) : T ` l : Ref T (T-Ref)

But where does come from? How about adding store to the typing relation

| ` (l) : T | ` l : Ref T (T-Ref)

..but..

Problem Problem

No finite type derivation for l with respect to the store.

(l1 x:Nat. (! l2)x, l2 x:Nat. (! l1)x)

Idea Idea

Instead of adding stores as argument to the typingrelation, we add store typings, which are mappingsfrom locations to types. Example for a store typing:

= (l1 Nat ! Nat, l2 Nat! Nat, l3 Unit)

Typing : Final Typing : Final

| ` t : T | ` ref t : Ref T

(T-Ref)

( l ) = T | ` l : Ref T

(T-Loc)

Typing : Final Typing : Final

| ` t : Ref T | ` ! t : T

(T-Deref)

| ` t1 : Ref T1 | ` t2 : T1

| ` t1 := t2 : Unit(T-Assign)

Safety = Preservation + Progress Safety = Preservation + Progress

Preservation, first attempt:

If | ` t : T and t | !t’ | ’, then | ` t’ : T

But the store and store typing may not agree!

| (l Unit ! Unit) ` ! l : Unit ! Unit ! l | (l unit) !unit | (l unit)

| ` unit : Unit for any ,

Well-Typed Stores Well-Typed Stores

A store is well-typed with respect to a typing context and a store typing , written as:

if dom()=dom() and | ` ( l ) : ( l ) for every l 2 dom()

Preservation : Second Attempt Preservation : Second Attempt

If: | ` t : T and t | !t’ | ’and | `

then | ` t’ : T

But : rule (E-RefV) extends the store:

l dom()

ref v | ! l | , ( l v)(E-RefV)

Preservation : Final VersionPreservation : Final Version

If: | ` t : T and t | !t’ | ’and | `

then for some ’ , | ’ ` t’ : T and | ’ ` ’

Exceptions : MotivationExceptions : Motivation

During execution, situations may occur that requires drastic measures such as resetting the state of program or even aborting the program.

• division by zero• arithmetic overflow• array index out of bounds• …

ErrorsErrors

We can denote error explicitly: t ::= … terms

error run-time error

(E-AppErr1)error t ! error

(E-AppErr2)v error ! error

` error : T (T-Error)

for any T

Evaluation Rules:

Typing Rule:

Examples Examples

( x:Nat. 0) error

(fix ( x:Nat. x)) error

( x:Bool. x) error

( x:Bool. x) (error true)

Error Handling : Motivation Error Handling : Motivation

In implementation, the evaluation of error will forcethe runtime stack to be cleared so that programreleases its computational resources.

Idea of error handling : Install a marker on the stack. During clearing of stack frames, the markers can be checked and when the right one is found, execution can resume normally.

Error HandlingError Handling

display try (..complicated calculation..) with “cannot compute the result”

Error HandlingError HandlingProvide a try-with (similar to try-catch of Java) mechanism. t ::= … terms

try t with t trap errors

(E-TryV)try v with t ! v

(E-TryError)try error with t ! t

t1!t’1

try t1 with t2! try t’1 with t2

(E-Try)

Typing Error HandlingTyping Error Handling

New Typing Rule:

` t1 : T ` t2 : T ` try t1 with t2 : T

(E-Try)

Exception Carrying Values: Motivation Exception Carrying Values: Motivation

Typically, we would like to know what kind of exceptional situation has occurred in order to take appropriate action.

Idea : Instead of errors, raise exception value that can be examined after trapping. This technique is called exception handling.

Exception Carrying ValueException Carrying ValueNew syntax:

t ::= … termsraise t raise exceptiontry t with t handle exception

Exception Carrying ValuesException Carrying ValuesNew Evaluation Rules:

t!t’raise t ! raise t’ (E-Raise)

(E-AppRaise1)(raise v) t ! raise v

v1 (raise v2) ! (raise v2) (E-AppRaise1)

(raise (raise v)) ! (raise v) (E-AppRaiseRaise)

Exception Carrying ValuesException Carrying Values

(E-TryV)try v with t ! v

(E-TryError)try (raise v) with t ! t v

t1!t’1

try t1 with t2! try t’1 with t2

(E-Try)

Exception Carrying ValuesException Carrying ValuesNew Typing Rules:

` t1 : T ` Texn ! T ` try t1 with t2 : T

(E-Try)

` t : Texn ` raise t : T

(E-Raise)

What Values can serve as Exceptions? What Values can serve as Exceptions?

• Texn is Nat as in return codes for Unix system calls.

• Texn is String for convenience in printing out messages.

• Texn is a certain fixed variant type, such as:

< divideByZero : Unit, overflow : Unit,

fileNotFound : String >

• …

O’Caml Exceptions O’Caml Exceptions

• Exceptions are a special extensible variant type.

• Syntax (exception l of T) does variant extension.

• Syntax (raise l(t)) is short for:raise (<l=t>) as Texn

• Syntax of try is sugar for try and case.

CS5205Extended Type System1 CS5205: Foundation in Programming Languages Lecture 7 : Extended Type...

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