Types

Type ::= FunType | TypeLambda | MatchType | InfixType FunType ::= FunTypeArgs ‘=>’ Type | TypeLambdaParams '=>' Type TypeLambda ::= TypeLambdaParams ‘=>>’ Type MatchType ::= InfixType ‘match’ <<< TypeCaseClauses >>> InfixType ::= RefinedType | RefinedTypeOrWildcard id [nl] RefinedTypeOrWildcard {id [nl] RefinedTypeOrWildcard} RefinedType ::= AnnotType {[nl] Refinement} AnnotType ::= SimpleType {Annotation} SimpleType ::= SimpleLiteral | SimpleType1 SimpleType1 ::= id | Singleton ‘.’ id | Singleton ‘.’ ‘type’ | ‘(’ TypesOrWildcards ‘)’ | Refinement | SimpleType1 TypeArgs | SimpleType1 ‘#’ id Singleton ::= SimpleRef | SimpleLiteral | Singleton ‘.’ id SimpleRef ::= id | [id ‘.’] ‘this’ | [id ‘.’] ‘super’ [‘[’ id ‘]’] ‘.’ id ParamType ::= [‘=>’] ParamValueType ParamValueType ::= ParamValueType [‘*’] TypeArgs ::= ‘[’ TypesOrWildcards ‘]’ Refinement ::= :<<< [RefineDef] {semi [RefineDef]} >>> FunTypeArgs ::= InfixType | ‘(’ [ FunArgTypes ] ‘)’ | FunParamClause FunArgTypes ::= FunArgType { ‘,’ FunArgType } FunArgType ::= Type | ‘=>’ Type FunParamClause ::= ‘(’ TypedFunParam {‘,’ TypedFunParam } ‘)’ TypedFunParam ::= id ‘:’ Type TypeLambdaParams ::= ‘[’ TypeLambdaParam {‘,’ TypeLambdaParam} ‘]’ TypeLambdaParam ::= {Annotation} (id | ‘_’) [TypeParamClause] TypeBounds TypeParamClause ::= ‘[’ VariantTypeParam {‘,’ VariantTypeParam} ‘]’ VariantTypeParam ::= {Annotation} [‘+’ | ‘-’] (id | ‘_’) [TypeParamClause] TypeBounds TypeCaseClauses ::= TypeCaseClause { TypeCaseClause } TypeCaseClause ::= ‘case’ (InfixType | ‘_’) ‘=>’ Type [semi] RefineDef ::= ‘val’ ValDef | ‘def’ DefDef | ‘type’ {nl} TypeDef TypeBounds ::= [‘>:’ Type] [‘<:’ Type] TypesOrWildcards ::= TypeOrWildcard {‘,’ TypeOrWildcard} TypeOrWildcard ::= Type | WildcardType RefinedTypeOrWildcard ::= RefinedType | WildcardType WildcardType ::= (‘?‘ | ‘_‘) TypeBounds 

The above grammer describes the concrete syntax of types that can be written in user code. Semantic operations on types in the Scala type system are better defined in terms of internal types, which are desugared from the concrete type syntax.

Internal Types

The following abstract grammar defines the shape of internal types. In this specification, unless otherwise noted, "types" refer to internal types. Internal types abstract away irrelevant details such as precedence and grouping, and contain shapes of types that cannot be directly expressed using the concrete syntax. They also contain simplified, decomposed shapes for complex concrete syntax types, such as refined types.

Type ::= ‘AnyKind‘ | ‘Nothing‘ | TypeLambda | DesignatorType | ParameterizedType | ThisType | SuperType | LiteralType | ByNameType | AnnotatedType | RefinedType | RecursiveType | RecursiveThis | UnionType | IntersectionType | MatchType | SkolemType TypeLambda ::= ‘[‘ TypeParams ‘]‘ ‘=>>‘ Type TypeParams ::= TypeParam {‘,‘ TypeParam} TypeParam ::= ParamVariance id TypeBounds ParamVariance ::= ε | ‘+‘ | ‘-‘ DesignatorType ::= Prefix ‘.‘ id Prefix ::= Type | PackageRef | ε PackageRef ::= id {‘.‘ id} ParameterizedType ::= Type ‘[‘ TypeArgs ‘]‘ TypeArgs ::= TypeArg {‘,‘ TypeArg} TypeArg ::= Type | WilcardTypeArg WildcardTypeArg ::= ‘?‘ TypeBounds ThisType ::= classid ‘.‘ ‘this‘ SuperType ::= classid ‘.‘ ‘super‘ ‘[‘ classid ‘]‘ LiteralType ::= SimpleLiteral ByNameType ::= ‘=>‘ Type AnnotatedType ::= Type Annotation RefinedType ::= Type ‘{‘ Refinement ‘}‘ Refinement ::= ‘type‘ id TypeAliasOrBounds | ‘def‘ id ‘:‘ TypeOrMethodic | ‘val‘ id ‘:‘ Type RecursiveType ::= ‘{‘ recid ‘=>‘ Type ‘}‘ RecursiveThis ::= recid ‘.‘ ‘this‘ UnionType ::= Type ‘｜‘ Type IntersectionType ::= Type ‘＆‘ Type MatchType ::= Type ‘match‘ ‘<:‘ Type ‘{‘ {TypeCaseClause} ‘}‘ TypeCaseClause ::= ‘case‘ TypeCasePattern ‘=>‘ Type TypeCasePattern ::= TypeCapture | TypeCaseAppliedPattern | Type TypeCaseAppliedPattern ::= Type ‘[‘ TypeCasePattern { ‘,‘ TypeCasePattern } ‘]‘ TypeCapture ::= (id | ‘_‘) TypeBounds SkolemType ::= ‘∃‘ skolemid ‘:‘ Type TypeOrMethodic ::= Type | MethodicType MethodicType ::= MethodType | PolyType MethodType ::= ‘(‘ MethodTypeParams ‘)‘ TypeOrMethodic MethodTypeParams ::= ε | MethodTypeParam {‘,‘ MethodTypeParam} MethodTypeParam ::= id ‘:‘ Type PolyType ::= ‘[‘ PolyTypeParams ‘]‘ TypeOrMethodic PolyTypeParams ::= PolyTypeParam {‘,‘ PolyTypeParam} PolyTypeParam ::= id TypeBounds TypeAliasOrBounds ::= TypeAlias | TypeBounds TypeAlias ::= ‘=‘ Type TypeBounds ::= ‘<:‘ Type ‘>:‘ Type 

Translation of Concrete Types into Internal Types

Concrete types are recursively translated, or desugared, into internal types. Most shapes of concrete types have a one-to-one translation to shapes of internal types. We elaborate hereafter on the translation of the other ones.

Infix Types

InfixType ::= CompoundType {id [nl] CompoundType} 

A concrete infix type ´T_1´ op ´T_2´ consists of an infix operator op which gets applied to two type operands ´T_1´ and ´T_2´. The type is translated to the internal type application op´[T_1, T_2]´. The infix operator op may be an arbitrary identifier.

Type operators follow the same precedence and associativity as term operators. For example, A + B * C parses as A + (B * C) and A | B & C parses as A | (B & C). Type operators ending in a colon ‘:’ are right-associative; all other operators are left-associative.

In a sequence of consecutive type infix operations ´t_0 \, \mathit{op} \, t_1 \, \mathit{op_2} \, ... \, \mathit{op_n} \, t_n´, all operators ´\mathit{op}_1, ..., \mathit{op}_n´ must have the same associativity. If they are all left-associative, the sequence is interpreted as ´(... (t_0 \mathit{op_1} t_1) \mathit{op_2} ...) \mathit{op_n} t_n´, otherwise it is interpreted as ´t_0 \mathit{op_1} (t_1 \mathit{op_2} ( ... \mathit{op_n} t_n) ...)´.

Under -source:future, if the type name is alphanumeric and the target type is not marked infix, a deprecation warning is emitted.

The type operators | and & are not really special. Nevertheless, unless shadowed, they resolve to the fundamental type aliases scala.| and scala.&, which represent union and intersection types, respectively.

Function Types

Type ::= FunTypeArgs ‘=>’ Type FunTypeArgs ::= InfixType | ‘(’ [ FunArgTypes ] ‘)’ | FunParamClause FunArgTypes ::= FunArgType { ‘,’ FunArgType } FunArgType ::= Type | ‘=>’ Type FunParamClause ::= ‘(’ TypedFunParam {‘,’ TypedFunParam } ‘)’ TypedFunParam ::= id ‘:’ Type 

The concrete function type ´(T_1, ..., T_n) \Rightarrow R´ represents the set of function values that take arguments of types ´T_1, ..., Tn´ and yield results of type ´R´. The case of exactly one argument type ´T \Rightarrow R´ is a shorthand for ´(T) \Rightarrow R´. An argument type of the form ´\Rightarrow T´ represents a call-by-name parameter of type ´T´.

Function types associate to the right, e.g. ´S \Rightarrow T \Rightarrow R´ is the same as ´S \Rightarrow (T \Rightarrow R)´.

Function types are covariant in their result type and contravariant in their argument types.

Function types translate into internal class types that define an apply method. Specifically, the ´n´-ary function type ´(T_1, ..., T_n) \Rightarrow R´ translates to the internal class type scala.Function´_n´[´T_1´, ..., ´T_n´, ´R´]. In particular ´() \Rightarrow R´ is a shorthand for class type scala.Function´_0´[´R´].

Such class types behave as if they were instances of the following trait:

trait Function´_n´[-´T_1´, ..., -´T_n´, +´R´]: def apply(´x_1´: ´T_1´, ..., ´x_n´: ´T_n´): ´R´ 

Their exact supertype and implementation can be consulted in the function classes section of the standard library page in this document.

Dependent function types are function types whose parameters are named and can referred to in result types. In the concrete type ´(x_1: T_1, ..., x_n: T_n) \Rightarrow R´, ´R´ can refer to the parameters ´x_i´, notably to form path-dependent types. It translates to the internal refined type

scala.Function´_n´[´T_1´, ..., ´T_n´, ´S´] { def apply(´x_1´: ´T_1´, ..., ´x_n´: ´T_n´): ´R´ } 

where ´S´ is the least super type of ´R´ that does not mention any of the ´x_i´.

Polymorphic function types are function types that take type arguments. Their result type must be a function type. In the concrete type ´[a_1 >: L_1 <: H_1, ..., a_n >: L_1 <: H_1] => (T_1, ..., T_m) => R´, the types ´T_j´ and ´R´ can refer to the type parameters ´a_i´. It translates to the internal refined type

scala.PolyFunction { def apply[´a_1 >: L_1 <: H_1, ..., a_n >: L_1 <: H_1´](´x_1´: ´T_1´, ..., ´x_n´: ´T_n´): ´R´ } 

Tuple Types

SimpleType1 ::= ... | ‘(’ TypesOrWildcards ‘)’ 

A tuple type ´(T_1, ..., T_n)´ where ´n \geq 2´ is sugar for the type ´T_1´ *: ... *: ´T_n´ *: scala.EmptyTuple, which is itself a series of nested infix types which are sugar for *:[´T_1´, *:[´T_2´, ... *:[´T_n´, scala.EmptyTuple]]]. The ´T_i´ can be wildcard type arguments.

Notes:

• (´T_1´) is the type ´T_1´, and not ´T_1´ *: scala.EmptyTuple (´T_1´ cannot be a wildcard type argument in that case).
• () is not a valid type (i.e. it is not desugared to scala.EmptyTuple).

Concrete Refined Types

RefinedType ::= AnnotType {[nl] Refinement} SimpleType1 ::= ... | Refinement Refinement ::= :<<< [RefineDef] {semi [RefineDef]} >>> RefineDef ::= ‘val’ ValDef | ‘def’ DefDef | ‘type’ {nl} TypeDef 

In the concrete syntax of types, refinements can contain several refined definitions. They must all be abstract. Moreover, the refined definitions can refer to each other as well as to members of the parent type, i.e., they have access to this.

In the internal types, each refinement defines exactly one refined definition, and references to this must be made explicit in a recursive type.

The conversion from the concrete syntax to the abstract syntax works as follows:

1. Create a fresh recursive this name ´\alpha´.
2. Replace every implicit or explicit reference to this in the refinement definitions by ´\alpha´.
3. Create nested refined types, one for every refined definition.
4. Unless ´\alpha´ was never actually used, wrap the result in a recursive type{ ´\alpha´ => ´...´ }.

Concrete Match Types

MatchType ::= InfixType ‘match’ <<< TypeCaseClauses >>> TypeCaseClauses ::= TypeCaseClause { TypeCaseClause } TypeCaseClause ::= ‘case’ (InfixType | ‘_’) ‘=>’ Type [semi] 

In the concrete syntax of match types, patterns are arbitrary InfixTypes, and there is no explicit notion of type capture. In the abstract syntax, however, captures are made explicit and can only appear as arguments to TypeCaseAppliedPatterns.

If the concrete pattern is _, its conversion is the internal type scala.Any. If it is a concrete InfixType, it is first converted to an internal type ´P´. If ´P´ is not a ParameterizedType, then use ´P´ as the internal pattern. Otherwise, ´P´ is recursively converted into a TypeCasePattern as follows:

1. If ´P´ is a WildcardTypeArg of the form ? >: ´L´ <: ´H´, return a TypeCapture of the form _ >: ´L´ <: ´H´.
2. If ´P´ is a direct type designator ´t´ whose name starts with a lowercase and was not written using backticks, return a TypeCapture´t´ >: ´L´ <: ´H´ where >: ´L´ <: ´H´ is the declared type definition of ´t´.
3. If ´P´ is a ParameterizedType of the form ´T´[´T_1´, ..., ´T_n´]:
   1. Recursively convert each ´T_i´ into a pattern ´P_i´.
   2. If ´P_i´ is a Type for all ´i´, return ´P´.
   3. Otherwise, return the TypeCaseAppliedPattern´T´[´P_1´, ..., ´P_n´].
4. Otherwise, return ´P´.

This conversion ensures that every TypeCaseAppliedPattern recursively contains at least one TypeCapture. Moreover, at the top level, the pattern is never a TypeCapture: all TypeCaptures are nested within a TypeCaseAppliedPattern.

The bound of the internal MatchType is always <: scala.Any by default. It can be overridden in a type member definition.

Concrete Type Lambdas

TypeLambda ::= TypeLambdaParams ‘=>>’ Type TypeLambdaParams ::= ‘[’ TypeLambdaParam {‘,’ TypeLambdaParam} ‘]’ TypeLambdaParam ::= {Annotation} (id | ‘_’) [TypeParamClause] TypeBounds TypeParamClause ::= ‘[’ VariantTypeParam {‘,’ VariantTypeParam} ‘]’ VariantTypeParam ::= {Annotation} [‘+’ | ‘-’] (id | ‘_’) [TypeParamClause] TypeBounds 

At the top level of concrete type lambda parameters, variance annotations are not allowed. However, in internal types, all type lambda parameters have explicit variance annotations.

When translating a concrete type lambda into an internal one, the variance of each type parameter is inferred from its usages in the body of the type lambda.

Definitions

From here onwards, we refer to internal types by default.

Kinds

The Scala type system is fundamentally higher-kinded. Types are either proper types, type constructors or poly-kinded types.

• Proper types are the types of terms.
• Type constructors are type-level functions from types to types.
• Poly-kinded types can take various kinds.

All types live in a single lattice with respect to a conformance relationship ´<:´. The top type is AnyKind and the bottom type is Nothing: all types conform to AnyKind, and Nothing conforms to all types. They can be referred to with the fundamental type aliases scala.AnyKind and scala.Nothing, respectively.

Types can be concrete or abstract. An abstract type ´T´ always has lower and upper bounds ´L´ and ´H´ such that ´L >: T´ and ´T <: H´. A concrete type ´T´ is considered to have itself as both lower and upper bound.

The kind of a type is indicated by its (transitive) upper bound:

• A type ´T <:´ scala.Any is a proper type.
• A type ´T <: K´ where ´K´ is a type lambda (of the form [´\pm a_1 >: L_1 <: H_1´, ..., ´\pm a_n >: L_n <: H_n´] =>> ´U´) is a type constructor.
• Other types are poly-kinded; they are neither proper types nor type constructors.

As a consequece, AnyKind itself is poly-kinded. Nothing is universally-kinded: it has all kinds at the same time, since it conforms to all types.

With this representation, it is rarely necessary to explicitly talk about the kinds of types. Usually, the kinds of types are implicit through their bounds.

Another way to look at it is that type bounds are kinds. They represent sets of types: ´>: L <: H´ denotes the set of types ´T´ such that ´L <: T´ and ´T <: H´. A set of types can be seen as a type of types, i.e., as a kind.

Conventions

Type bounds are formally always of the form ´>: L <: H´. By convention, we can omit either of both bounds in writing.

• When omitted, the lower bound ´L´ is Nothing.
• When omitted, the higher bound ´H´ is Any (notAnyKind).

These conventions correspond to the defaults in the concrete syntax.

Proper Types

Proper types are also called value types, as they represent sets of values.

Stable types are value types that contain exactly one non-null value. Stable types can be used as prefixes in named designator types. The stable types are

• designator types referencing a stable term,
• this types,
• super types,
• literal types,
• recursive this types, and
• skolem types.

Every stable type ´T´ is concrete and has an underlying type ´U´ such that ´T <: U´.

Type Constructors

To each type constructor corresponds an inferred type parameter clause which is computed as follows:

• For a type lambda, its type parameter clause (including variance annotations).
• For a polymorphic class type, the type parameter clause of the referenced class definition.
• For a non-class type designator, the inferred clause of its upper bound.

Type Definitions

A type definition ´D´ represents the right-hand-side of a type member definition or the bounds of a type parameter. It is either:

• a type alias of the form ´= U´, or
• an abstract type definition with bounds ´>: L <: H´.

All type definitions have a lower bound ´L´ and an upper bound ´H´, which are types. For type aliases, ´L = H = U´.

The type definition of a type parameter is never a type alias.

Types

Type Lambdas

TypeLambda ::= ‘[‘ TypeParams ‘]‘ ‘=>>‘ Type TypeParams ::= TypeParam {‘,‘ TypeParam} TypeParam ::= ParamVariance id TypeBounds ParamVariance ::= ε | ‘+‘ | ‘-‘ 

A type lambda of the form [´\pm a_1 >: L_1 <: H_1´, ..., ´\pm a_n >: L_n <: H_n´] =>> ´U´ is a direct representation of a type constructor with ´n´ type parameters. When applied to ´n´ type arguments that conform to the specified bounds, it produces another type ´U´. Type lambdas are always concrete types.

The scope of a type parameter extends over the result type ´U´ as well as the bounds of the type parameters themselves.

All type constructors conform to some type lambda.

The type bounds of the parameters of a type lambda are in contravariant position, while its result type is in covariant position. If some type constructor ´T <:´ [´\pm a_1 >: L_1 <: H_1´, ..., ´\pm a_n >: L_n <: H_n´] =>> ´U´, then ´T´'s ´i´th type parameter bounds contain the bounds ´>: L_i <: H_i´, and its result type conforms to ´U´.

Note: the concrete syntax of type lambdas does not allow to specify variances for type parameters. Instead, variances are inferred from the body of the lambda to be as general as possible.

Example

type Lst = [T] =>> List[T] // T is inferred to be covariant with bounds >: Nothing <: Any type Fn = [A <: Seq[?], B] =>> (A => B) // A is inferred to be contravariant, B covariant val x: Lst[Int] = List(1) // ok, Lst[Int] expands to List[Int] val f: Fn[List[Int], Int] = (x: List[Int]) => x.head // ok val g: Fn[Int, Int] = (x: Int) => x // error: Int does not conform to the bound Seq[?] def liftPair[F <: [T] =>> Any](f: F[Int]): Any = f liftPair[Lst](List(1)) // ok, Lst <: ([T] =>> Any) 

Designator Types

DesignatorType ::= Prefix ‘.‘ id Prefix ::= Type | PackageRef | ε PackageRef ::= id {‘.‘ id} 

A designator type (or designator for short) is a reference to a definition. Term designators refer to term definitions, while type designators refer to type definitions.

In the abstract syntax, the id retains whether it is a term or type. In the concrete syntax, an id refers to a type designator, while id.type refers to a term designator. In that context, term designators are often called singleton types.

Designators with an empty prefix ´\epsilon´ are called direct designators. They refer to local definitions available in the scope:

• Local type, object, val, lazy val, var or def definitions
• Term or type parameters

The ids of direct designators are protected from accidental shadowing in the abstract syntax. They retain the identity of the exact definition they refer to, rather than relying on scope-based name resolution. ¹

The ´\epsilon´ prefix cannot be written in the concrete syntax. A bare id is used instead and resolved based on scopes.

Named designators refer to member definitions of a non-empty prefix:

• Top-level definitions, including top-level classes, have a package ref prefix
• Class member definitions and refinements have a type prefix

Term Designators

A term designator ´p.x´ referring to a term definition t has an underlying type ´U´. If ´p = \epsilon´ or ´p´ is a package ref, the underlying type ´U´ is the declared type of t and ´p.x´ is a stable type if an only if t is a val or object definition. Otherwise, the underlying type ´U´ and whether ´p.x´ is a stable type are determined by memberType(´p´, ´x´).

All term designators are concrete types. If scala.Null ´<: U´, the term designator denotes the set of values consisting of null and the value denoted by ´t´, i.e., the value ´v´ for which t eq v. Otherwise, the designator denotes the singleton set only containing ´v´.

Type Designators

A type designator ´p.C´ referring to a class definition (including traits and hidden object classes) is a class type. If the class is monomorphic, the type designator is a value type denoting the set of instances of ´C´ or any of its subclasses. Otherwise it is a type constructor with the same type parameters as the class definition. All class types are concrete, non-stable types.

If a type designator ´p.T´ is not a class type, it refers to a type definition T (a type parameter or a type member definition) and has an underlying type definition. If ´p = \epsilon´ or ´p´ is a package ref, the underlying type definition is the declared type definition of T. Otherwise, it is determined by memberType(´p´, ´T´). A non-class type designator is concrete (resp. stable) if and only if its underlying type definition is an alias ´U´ and ´U´ is itself concrete (resp. stable).

Parameterized Types

ParameterizedType ::= Type ‘[‘ TypeArgs ‘]‘ TypeArgs ::= TypeArg {‘,‘ TypeArg} TypeArg ::= Type | WilcardTypeArg WildcardTypeArg ::= ‘?‘ TypeBounds 

A parameterized type ´T[T_1, ..., T_n]´ consists of a type constructor ´T´ and type arguments ´T_1, ..., T_n´ where ´n \geq 1´. The parameterized type is well-formed if

• ´T´ is a type constructor which takes ´n´ type parameters ´a_1, ..., a_n´, i.e., it must conform to a type lambda of the form ´[\pm a_1 >: L_1 <: H_1, ..., \pm a_n >: L_n <: H_n] => U´, and
• if ´T´ is an abstract type constructor, none of the type arguments is a wildcard type argument, and
• each type argument conforms to its bounds, i.e., given ´\sigma´ the substitution ´[a_1 := T_1, ..., a_n := T_n]´, for each type ´i´:
  • if ´T_i´ is a type and ´\sigma L_i <: T_i <: \sigma H_i´, or
  • ´T_i´ is a wildcard type argument ´? >: L_{Ti} <: H_{Ti}´ and ´\sigma L_i <: L_{Ti}´ and ´H_{Ti} <: \sigma H_i´.

´T[T_1, ..., T_n]´ is a parameterized class type if and only if ´T´ is a class type. All parameterized class types are value types.

In the concrete syntax of wildcard type arguments, if both bounds are omitted, the real bounds are inferred from the bounds of the corresponding type parameter in the target type constructor (which must be concrete). If only one bound is omitted, Nothing or Any is used, as usual.

Also in the concrete syntax, _ can be used instead of ? for compatibility reasons, with the same meaning. This alternative will be deprecated in the future, and is already deprecated under -source:future.

Simplification Rules

Wildcard type arguments used in covariant or contravariant positions can always be simplified to regular types.

Let ´T[T_1, ..., T_n]´ be a parameterized type for a concrete type constructor. Then, applying a wildcard type argument ´? >: L <: H´ at the ´i´'th position obeys the following equivalences:

• If the type parameter ´T_i´ is declared covariant, then ´T[..., ? >: L <: H, ...] =:= T[..., H, ...]´.
• If the type parameter ´T_i´ is declared contravariant, then ´T[..., ? >: L <: H, ...] =:= T[..., L, ...]´.

Example Parameterized Types

Given the partial type definitions:

class TreeMap[A <: Comparable[A], B] { ... } class List[+A] { ... } class I extends Comparable[I] { ... } class F[M[A], X] { ... } // M[A] desugars to M <: [A] =>> Any class S[K <: String] { ... } class G[M[Z <: I], I] { ... } // M[Z <: I] desugars to M <: [Z <: I] =>> Any 

the following parameterized types are well-formed:

TreeMap[I, String] List[I] List[List[Boolean]] F[List, Int] F[[X] =>> List[X], Int] G[S, String] List[?] // ? inferred as List[_ >: Nothing <: Any], equivalent to List[Any] List[? <: String] // equivalent to List[String] S[? <: String] F[?, Boolean] // ? inferred as ? >: Nothing <: [A] =>> Any 

and the following types are ill-formed:

TreeMap[I] // illegal: wrong number of parameters TreeMap[List[I], Int] // illegal: type parameter not within bound List[[X] => List[X]] F[Int, Boolean] // illegal: Int is not a type constructor F[TreeMap, Int] // illegal: TreeMap takes two parameters, // F expects a constructor taking one F[[X, Y] => (X, Y)] G[S, Int] // illegal: S constrains its parameter to // conform to String, // G expects type constructor with a parameter // that conforms to Int 

The following code also contains an ill-formed type:

trait H[F[A]]: // F[A] desugars to F <: [A] =>> Any, which is abstract def f: F[_] // illegal : an abstract type constructor // cannot be applied to wildcard arguments. 

This Types

ThisType ::= classid ‘.‘ ‘this‘ 

A this type´C´.this denotes the this value of class ´C´ within ´C´.

This types often appear implicitly as the prefix of designator types referring to members of ´C´. They play a particular role in the type system, since they are affected by the as seen from operation on types.

This types are stable types. The underlying type of ´C´.this is the self type of ´C´.

Super Types

SuperType ::= classid ‘.‘ ‘super‘ ‘[‘ classid ‘]‘ 

A super type´C´.super[´D´] denotes the this value of class C within C, but "widened" to only see members coming from a parent class or trait ´D´.

Super types exist for compatibility with Scala 2, which allows shadowing of inner classes. In a Scala 3-only context, a super type can always be replaced by the corresponding this type. Therefore, we omit further discussion of super types in this specification.

Literal Types

LiteralType ::= SimpleLiteral 

A literal type lit denotes the single literal value lit. Thus, the type ascription 1: 1 gives the most precise type to the literal value 1: the literal type 1.

At run time, an expression e is considered to have literal type lit if e == lit. Concretely, the result of e.isInstanceOf[lit] and e match { case _ : lit => } is determined by evaluating e == lit.

Literal types are available for all primitive types, as well as for String. However, only literal types for Int, Long, Float, Double, Boolean, Char and String can be expressed in the concrete syntax.

Literal types are stable types. Their underlying type is the primitive type containing their value.

Example

val x: 1 = 1 val y: false = false val z: false = y val int: Int = x val badX: 1 = int // error: Int is not a subtype of 1 val badY: false = true // error: true is not a subtype of false 

By-Name Types

ByNameType ::= ‘=>‘ Type 

A by-name type ´=> T´ denotes the declared type of a by-name term parameter. By-name types can only appear as the types of parameters in method types, and as type arguments in parameterized types.

Annotated Types

AnnotatedType ::= Type Annotation 

An annotated type ´T a´ attaches the annotation ´a´ to the type ´T´.

Example

The following type adds the @suspendable annotation to the type String:

String @suspendable 

Refined Types

RefinedType ::= Type ‘{‘ Refinement ‘}‘ Refinement ::= ‘type‘ id TypeAliasOrBounds | ‘def‘ id ‘:‘ TypeOrMethodic | ‘val‘ id ‘:‘ Type 

A refined type ´T { R }´ denotes the set of values that belong to ´T´ and also have a member conforming to the refinement ´R´.

The refined type ´T { R }´ is well-formed if:

• ´T´ is a proper type, and
• if ´R´ is a term (def or val) refinement, the refined type is a proper type, and
• if ´R´ overrides a member of ´T´, the usual rules for overriding apply, and
• if ´R´ is a def refinement with a polymorphic method type, then ´R´ overrides a member definition of ´T´.

As an exception to the last rule, a polymorphic method type refinement is allowed if ´T <:´ scala.PolyFunction and ´id´ is the name apply.

If the refinement ´R´ overrides no member of ´T´ and is not an occurrence of the scala.PolyFunction exception, the refinement is said to be “structural” ².

Note: since a refinement does not define a class, it is not possible to use a this type to reference term and type members of the parent type ´T´ within the refinement. When the surface syntax of refined types makes such references, a recursive type wraps the refined type, given access to members of self through a recursive-this type.

Example

Given the following class definitions:

trait T: type X <: Option[Any] def foo: Any def fooPoly[A](x: A): Any trait U extends T: override def foo: Int override def fooPoly[A](x: A): A trait V extends T type X = Some[Int] def bar: Int def barPoly[A](x: A): A 

We get the following conformance relationships:

• U <: T { def foo: Int }
• U <: T { def fooPoly[A](x: A): A }
• U <: (T { def foo: Int }) { def fooPoly[A](x: A): A } (we can chain refined types to refine multiple members)
• V <: T { type X <: Some[Any] }
• V <: T { type X >: Some[Nothing] }
• V <: T { type X = Some[Int] }
• V <: T { def bar: Any } (a structural refinement)

The following refined types are not well-formed:

• T { def barPoly[A](x: A): A } (structural refinement for a polymorphic method type)
• T { type X <: List[Any] } (does not satisfy overriding rules)
• List { def head: Int } (the parent type List is not a proper type)
• T { def foo: List } (the refined type List is not a proper type)
• T { def foo: T.this.X } (T.this is not allowed outside the body of T)

Recursive Types

RecursiveType ::= ‘{‘ recid ‘=>‘ Type ‘}‘ RecursiveThis ::= recid ‘.‘ ‘this‘ 

A recursive type of the form { ´\alpha´ => ´T´ } represents the same values as ´T´, while offering ´T´ access to its recursive this type ´\alpha´.

Recursive types cannot directly be expressed in the concrete syntax. They are created as needed when a refined type in the concrete syntax contains a refinement that needs access to the this value. Each recursive type defines a unique self-reference ´\alpha´, distinct from any other recursive type in the system.

Recursive types can be unfolded during subtyping as needed, replacing references to its ´\alpha´ by a stable reference to the other side of the conformance relationship.

Example

Given the class definitions in the refined types section, we can write the following refined type in the source syntax:

T { def foo: X } // equivalent to T { def foo: this.X } 

This type is not directly expressible as a refined type alone, as the refinement cannot access the this value. Instead, in the abstract syntax of types, it is translated to { ´\alpha´ => ´T´ { def foo: ´\alpha´.X } }.

Given the following definitions:

trait Z extends T: type X = Option[Int] def foo: Option[Int] = Some(5) val z: Z 

we can check that z ´<:´ { ´\alpha´ => ´T´ { def foo: ´\alpha´.X } }. We first unfold the recursive type, substituting ´z´ for ´\alpha´, resulting in z ´<:´ T { def foo: z.X }. Since the underlying type of ´z´ is ´Z´, we can resolve z.X to mean Option[Int], and then validate that z ´<:´ T and that z has a member def foo: Option[Int].

Union and Intersection Types

UnionType ::= Type ‘｜‘ Type IntersectionType ::= Type ‘＆‘ Type 

Syntactically, the types S | T and S & T are infix types, where the infix operators are | and &, respectively (see infix types).

However, in this specification, ´S ｜ T´ and ´S ＆ T´ refer to the underlying core concepts of union and intersection types, respectively.

• The type ´S ｜ T´ represents the set of values that are represented by either ´S´ or ´T´.
• The type ´S ＆ T´ represents the set of values that are represented by both ´S´ and ´T´.

From the conformance rules rules on union and intersection types, we can show that ´＆´ and ´｜´ are commutative and associative. Moreover, ＆ is distributive over ｜. For any type ´A´, ´B´ and ´C´, all of the following relationships hold:

• ´A ＆ B =:= B ＆ A´,
• ´A ｜ B =:= B ｜ A´,
• ´(A ＆ B) ＆ C =:= A ＆ (B ＆ C)´,
• ´(A ｜ B) ｜ C =:= A ｜ (B ｜ C)´, and
• ´A ＆ (B ｜ C) =:= (A ＆ B) ｜ (A ＆ C)´.

If ´C´ is a co- or contravariant type constructor, ´C[A] ＆ C[B]´ can be simplified using the following rules:

• If ´C´ is covariant, ´C[A] ＆ C[B] =:= C[A ＆ B]´
• If ´C´ is contravariant, ´C[A] ＆ C[B] =:= C[A | B]´

The right-to-left validity of the above two rules can be derived from the definition of covariance and contravariance and the conformance rules of union and intersection types:

• When ´C´ is covariant, we can derive ´C[A ＆ B] <: C[A] ＆ C[B]´.
• When ´C´ is contravariant, we can derive ´C[A ｜ B] <: C[A] ＆ C[B]´.

Join of a union type

In some situations, a union type might need to be widened to a non-union type. For this purpose, we define the join of a union type ´T_1 ｜ ... ｜ T_n´ as the smallest intersection type of base class instances of ´T_1, ..., T_n´. Note that union types might still appear as type arguments in the resulting type, this guarantees that the join is always finite.

For example, given

trait C[+T] trait D trait E class A extends C[A] with D class B extends C[B] with D with E 

The join of ´A ｜ B´ is ´C[A ｜ B] ＆ D´

Match Types

MatchType ::= Type ‘match‘ ‘<:‘ Type ‘{‘ {TypeCaseClause} ‘}‘ TypeCaseClause ::= ‘case‘ TypeCasePattern ‘=>‘ Type TypeCasePattern ::= TypeCapture | TypeCaseAppliedPattern | Type TypeCaseAppliedPattern ::= Type ‘[‘ TypeCasePattern { ‘,‘ TypeCasePattern } ‘]‘ TypeCapture ::= (id | ‘_‘) TypeBounds 

A match type contains a scrutinee, a list of case clauses, and an upper bound. The scrutinee and the upper bound must both be proper types. A match type can be reduced to the body of a case clause if the scrutinee matches its pattern, and if it is provably disjoint from every earlier pattern.

Legal patterns

A TypeCasePattern is a legal pattern if and only if one of the following is true:

• It is a Type, or
• It is a TypeCaseAppliedPattern of the form ´P´[´P_1´, ..., ´P_n´] where ´P´ is a "pattern-legal type constructor" and for each ´i´, either:
  • ´P_i´ is a TypeCapture, or
  • ´P_i´ is a Type, or
  • ´P_i´ is a TypeCaseAppliedPattern, the type constructor ´P´ is covariant in its ´i´th type parameter, and ´P_i´ is recursively a legal pattern.

A type ´P´ is a "pattern-legal type constructor" if one of the following is true:

• It is a class type constructor, or
• It is the scala.compiletime.ops.int.S type constructor, or
• It is an abstract type constructor, or
• It is a type lambda with a refined result type of the form [´a´ >: ´L´ <: ´H´] =>> ´B´ { type ´T´ = ´a´ } where:
  • ´B´ contains no occurrence of ´a´,
  • there exists a type member ´T´ in ´B´, and
  • the bounds >: ´L´ <: ´H´ are not any more restrictive than those of ´T´ in ´B´, i.e., ´L´ <: ´(∃ \alpha : B).T´ <: ´H´.
• It is a type lambda of the form [´\pm a_1 >: L_1 <: H_1´, ..., ´\pm a_n >: L_n <: H_n´] =>> ´U´ such that:
  • Its bounds contain all possible values of its arguments, and
  • When applied to the type arguments, it beta-reduces to a new legal MatchTypeAppliedPattern that contains exactly one instance of every type capture present in the type arguments.
• It is a concrete type designator with underlying type definition = ´U´ and ´U´ is recursively a "pattern-legal type constructor".

Examples of legal patterns

Given the following definitions:

class Inv[A] class Cov[+A] class Contra[-A] class Base { type Y } type YExtractor[t] = Base { type Y = t } type ZExtractor[t] = Base { type Z = t } type IsSeq[t <: Seq[Any]] = t 

Here are examples of legal patterns:

// TypeWithoutCapture's case Any => // also the desugaring of `case _ =>` when the _ is at the top-level case Int => case List[Int] => case Array[String] => // Class type constructors with direct captures case scala.collection.immutable.List[t] => // not Predef.List; it is a type alias case Array[t] => case Contra[t] => case Either[s, t] => case Either[s, Contra[Int]] => case h *: t => case Int *: t => // The S type constructor case S[n] => // An abstract type constructor // given a [F[_]] or `type F[_] >: L <: H` in scope case F[t] => // Nested captures in covariant position case Cov[Inv[t]] => case Cov[Cov[t]] => case Cov[Contra[t]] => case Array[h] *: t => // sugar for *:[Array[h], t] case g *: h *: EmptyTuple => // Type aliases case List[t] => // which is Predef.List, itself defined as `type List[+A] = scala.collection.immutable.List[A]` // Refinements (through a type alias) case YExtractor[t] => 

The following patterns are not legal:

// Type capture nested two levels below a non-covariant type constructor case Inv[Cov[t]] => case Inv[Inv[t]] => case Contra[Cov[t]] => // Type constructor with bounds that do not contain all possible instantiations case IsSeq[t] => // Type refinement where the refined type member is not a member of the parent case ZExtractor[t] => 

Matching

Given a scrutinee X and a match type case case P => R with type captures ts, matching proceeds in three steps:

1. Compute instantiations for the type captures ts', and check that they are specific enough.
2. If successful, check that X <:< [ts := ts']P.
3. If successful, reduce to [ts := ts']R.

The instantiations are computed by the recursive function matchPattern(X, P, variance, scrutIsWidenedAbstract). At the top level, variance = 1 and scrutIsWidenedAbstract = false.

matchPattern behaves according to what kind is P:

• If P is a TypeWithoutCapture:
  • Do nothing (always succeed).
• If P is a WildcardCaptureti = _:
  • If X is of the form _ >: L <: H, instantiate ti := H (anything between L and H would work here),
  • Otherwise, instantiate ti := X.
• If P is a TypeCaptureti:
  • If X is of the form _ >: L <: H,
    • If scrutIsWidenedAbstract is true, fail as not specific.
    • Otherwise, if variance = 1, instantiate ti := H.
    • Otherwise, if variance = -1, instantiate ti := L.
    • Otherwise, fail as not specific.
  • Otherwise, if variance = 0 or scrutIsWidenedAbstract is false, instantiate ti := X.
  • Otherwise, fail as not specific.
• If P is a MatchTypeAppliedPattern of the form T[Qs]:
  • Assert: variance = 1 (from the definition of legal patterns).
  • If T is a class type constructor of the form p.C:
    • If baseType(X, C) is not defined, fail as not matching.
    • Otherwise, it is of the form q.C[Us].
    • If p =:= q is false, fail as not matching.
    • Let innerScrutIsWidenedAbstract be true if either scrutIsWidenedAbstract or X is not a concrete type.
    • For each pair of (Ui, Qi), compute matchPattern(Ui, Qi, vi, innerScrutIsWidenedAbstract) where vi is the variance of the ith type parameter of T.
  • If T is scala.compiletime.ops.int.S:
    • If n = natValue(X) is undefined or n <= 0, fail as not matching.
    • Otherwise, compute matchPattern(n - 1, Q1, 1, scrutIsWidenedAbstract).
  • If T is an abstract type constructor:
    • If X is not of the form F[Us] or F =:= T is false, fail as not matching.
    • Otherwise, for each pair of (Ui, Qi), compute matchPattern(Ui, Qi, vi, scrutIsWidenedAbstract) where vi is the variance of the ith type parameter of T.
  • If T is a refined type of the form Base { type Y = ti }:
    • Let q be X if X is a stable type, or the skolem type ∃α:X otherwise.
    • If q does not have a type member Y, fail as not matching (that implies that X <:< Base is false, because Base must have a type member Y for the pattern to be legal).
    • If q is not a skolem type:
      • Compute matchPattern(ti, q.Y, 0, scrutIsWidenedAbstract).
    • Otherwise, if q.Y is a type alias with underlying type definition = U:
      • Let U' be the result of perform type avoidance on U to remove references to q.
      • If successful, compute matchPattern(ti, U', 0, scrutIsWidenedAbstract).
      • Otherwise, fail as not specific.
  • If T is a concrete type alias to a type lambda:
    • Let P' be the beta-reduction of P.
    • Compute matchPattern(P', X, variance, scrutIsWidenedAbstract).

Disjointness

A scrutinee ´X´ is provably disjoint from a pattern ´P´ iff it is provably disjoint from the type ´P'´ obtained by replacing every type capture in ´P´ by a wildcard type argument with the same bounds.

We note ´X ⋔ Y´ to say that ´X´ and ´Y´ are provably disjoint. Intuitively, that notion is based on the following properties of the Scala language:

• Single inheritance of classes
• Final classes cannot be extended
• Sealed traits have a known set of direct children
• Constant types with distinct values are nonintersecting
• Singleton paths to distinct enum case values are nonintersecting

However, a precise definition of provably-disjoint is complicated and requires some helpers. We start with the notion of "simple types", which are a minimal subset of Scala internal types that capture the concepts mentioned above.

A "simple type" is one of:

• Nothing
• AnyKind
• ´p.C[...X_i]´ a possibly parameterized class type, where ´p´ and ´...X_i´ are arbitrary types (not just simple types)
• ´c´ a literal type
• ´p.C.x´ where C is an enum class and x is one of its value cases
• ´X_1 ＆ X_2´ where ´X_1´ and ´X_2´ are both simple types
• ´X_1 ｜ X_2´ where ´X_1´ and ´X_2´ are both simple types
• ´[...a_i]´ =>> ´X_1´ where ´X_1´ is a simple type

We define ´⌈X⌉´ a function from a full Scala type to a simple type. Intuitively, it returns the "smallest" simple type that is a supertype of X. It is defined as follows:

• ´⌈X⌉ = X´ if ´X´ is a simple type
• ´⌈X⌉ = ⌈U⌉´ if ´X´ is a stable type but not a simple type and its underlying type is ´U´
• ´⌈X⌉ = ⌈H⌉´ if ´X´ is a non-class type designator with upper bound ´H´
• ´⌈X⌉ = ⌈η(X)⌉´ if ´X´ is a polymorphic class type designator, where ´η(X)´ is its eta-expansion
• ´⌈X⌉ = ⌈Y⌉´ if ´X´ is a match type that reduces to ´Y´
• ´⌈X⌉ = ⌈H⌉´ if ´X´ is a match type that does not reduce and ´H´ is its upper bound
• ´⌈X[...T_i]⌉ = ⌈Y⌉´ where Y is the beta-reduction of ´X[...T_i]´ if ´X´ is a type lambda
• ´⌈X[...T_i]⌉ = ⌈⌈X⌉[...T_i]⌉´ if ´X´ is neither a type lambda nor a class type designator
• ´⌈X @a⌉ = ⌈X⌉´
• ´⌈X´ { ´R´ }´⌉ = ⌈X⌉´
• ´⌈´{ ´α´ => ´X´ }´ = ⌈X⌉⌉´
• ´⌈X_1 ＆ X_2⌉ = ⌈X_1⌉ ＆ ⌈X_2⌉´
• ´⌈X_1 ｜ X_2⌉ = ⌈X_1⌉ ｜ ⌈X_2⌉´
• ´⌈[...a_i]´ =>> ´X_1⌉ = [...a_i]´ =>> ´⌈X_1⌉´

The following properties hold about ´⌈X⌉´ (we have paper proofs for those):

• ´X <: ⌈X⌉´ for all type ´X´.
• If ´S <: T´, and the subtyping derivation does not use the "lower-bound rule" of ´<:´ anywhere, then ´⌈S⌉ <: ⌈T⌉´.

The "lower-bound rule" states that ´S <: T´ if ´T = q.X´ and ´q.X´ is a non-class type designator and ´S <: L´ where ´L´ is the lower bound of the underlying type definition of ´q.X´". That rule is known to break transitivy of subtyping in Scala already.

Second, we define the relation ´⋔´ on classes (including traits, hidden classes of objects, and enum terms) as:

• ´C ⋔ D´ if ´C ∉´ baseClasses´(D)´ and ´D´ is final
• ´C ⋔ D´ if ´D ∉´ baseClasses´(C)´ and ´C´ is final
• ´C ⋔ D´ if there exists classes ´C' ∈´ baseClasses´(C)´ and ´D' ∈´ baseClasses´(D)´ such that ´C' ∉´ baseClasses´(D')´ and ´D' ∉´ baseClasses´(C')´.
• ´C ⋔ D´ if ´C´ is sealed without anonymous child and ´C_i ⋔ D´ for all direct children ´C_i´ of ´C´
• ´C ⋔ D´ if ´D´ is sealed without anonymous child and ´C ⋔ D_i´ for all direct children ´D_i´ of ´D´

We can now define ´⋔´ for types.

For arbitrary types ´X´ and ´Y´, we define ´X ⋔ Y´ as ´⌈X⌉ ⋔ ⌈Y⌉´.

Two simple types ´S´ and ´T´ are provably disjoint if there is a finite derivation tree for ´S ⋔ T´ using the following rules. Most rules go by pair, which makes the whole relation symmetric:

• Nothing is disjoint from everything (including itself):
  • Nothing ´⋔ T´
  • ´S ⋔´ Nothing
• A union type is disjoint from another type if both of its parts are disjoint from that type:
  • ´S ⋔ T_1 ｜ T_2´ if ´S ⋔ T_1´ and ´S ⋔ T_2´
  • ´S_1 ｜ S_2 ⋔ T´ if ´S_1 ⋔ T´ and ´S_2 ⋔ T´
• An intersection type is disjoint from another type if at least one of its parts is disjoint from that type:
  • ´S ⋔ T_1 ＆ T_2´ if ´S ⋔ T_1´ or ´S ⋔ T_2´
  • ´S_1 ＆ S_2 ⋔ T´ if ´S_1 ⋔ T´ or ´S_1 ⋔ T´
• A type lambda is disjoint from any other type that is not a type lambda with the same number of parameters:
  • ´[...a_i]´ =>> ´S_1 ⋔ q.D.y´
  • ´[...a_i]´ =>> ´S_1 ⋔ d´
  • ´[...a_i]´ =>> ´S_1 ⋔ q.D[...T_i]´
  • ´p.C.x ⋔ [...b_i]´ =>> ´T_1´
  • ´c ⋔ [...b_i]´ =>> ´T_1´
  • ´p.C[...S_i] ⋔ [...b_i]´ =>> ´T_1´
  • ´[a_1, ..., a_n]´ =>> ´S_1 ⋔ [b_1, ..., b_m]´ =>> ´T_1´ if ´m \neq n´
• Two type lambdas with the same number of type parameters are disjoint if their result types are disjoint:
  • ´[a_1, ..., a_n]´ =>> ´S_1 ⋔ [b_1, ..., b_n]´ =>> ´T_1´ if ´S_1 ⋔ T_1´
• An enum value case is disjoint from any other enum value case (identified by either not being in the same enum class, or having a different name):
  • ´p.C.x ⋔ q.D.y´ if ´C \neq D´ or ´x \neq y´
• Two literal types are disjoint if they are different:
  • ´c ⋔ d´ if ´c \neq d´
• An enum value case is always disjoint from a literal type:
  • ´c ⋔ q.D.y´
  • ´p.C.x ⋔ d´
• An enum value case or a constant is disjoint from a class type if it does not extend that class (because it's essentially final):
  • ´p.C.x ⋔ q.D[...T_i]´ if baseType´(p.C.x, D)´ is not defined
  • ´p.C[...S_i] ⋔ q.D.y´ if baseType´(q.D.y, C)´ is not defined
  • ´c ⋔ q.D[...T_i]´ if baseType´(c, D)´ is not defined
  • ´p.C[...S_i] ⋔ d´ if baseType´(d, C)´ is not defined
• Two class types are disjoint if the classes themselves are disjoint, or if there exists a common super type with conflicting type arguments.
  • ´p.C[...S_i] ⋔ q.D[...T_i]´ if ´C ⋔ D´
  • ´p.C[...S_i] ⋔ q.D[...T_i]´ if there exists a class ´E´ such that baseType´(p.C[...S_i], E) = a.E[...A_i]´ and baseType´(q.D[...T_i], E) = b.E[...B_i]´ and there exists a pair ´(A_i, B_i)´ such that
    • ´A_i ⋔ B_i´ and it is in covariant position and there exists a field of that type parameter in ´E´, or
    • ´A_i ⋔ B_i´ and it is in invariant position, and:
      • there exists a field of that type parameter in ´E´, or
      • ´A_i´ cannot be instantiated to Nothing, or
      • ´B_i´ cannot be instantiated to Nothing.

It is worth noting that this definition disregards prefixes entirely. ´p.C´ and ´q.C´ are never provably disjoint, even if ´p´ could be proven disjoint from ´q´. It also disregards type members.

We have a proof sketch of the following property for ´⋔´:

• If ´S <: T´ and ´T ⋔ U´, then ´S ⋔ U´.

This is a very desirable property. It means that if we make the scrutinee of a match type more precise (a subtype) through substitution, and the match type previously reduced, then the match type will still reduce to the same case.

Note: if ´⋔´ were a "true" disjointness relationship, and not a provably-disjoint relationship, that property would trivially hold based on elementary set theoretic properties. It would amount to proving that if ´S ⊆ T´ and ´T ⋂ U = ∅´, then ´S ⋂ U = ∅´.

Reduction

The result of reducing ´X´ match { case ´P_1´ => ´R_1; ...;´ case ´P_n´ => ´R_n´ } can be a type, undefined, or a compile error.

For ´n \geq 1´, it is specified as:

• If ´X´ matches ´P_1´ with type capture instantiations ´[...t_i \to t_i']´:
  • If ´X ⋔ P_1´, do not reduce.
  • Otherwise, reduce as ´[...t_i \to t_i']R_1´.
• Otherwise,
  • If ´X ⋔ P1´, the result of reducing ´X´ match { case ´P_2´ => ´R_2; ...;´ case ´P_n´ => ´R_n´ } (i.e., proceed with subsequent cases).
  • Otherwise, do not reduce.

The reduction of an "empty" match type ´X´ match { } (which cannot be written in user programs) is a compile error.

Skolem Types

SkolemType ::= ‘∃‘ skolemid ‘:‘ Type 

Skolem types cannot directly be written in the concrete syntax. Moreover, although they are proper types, they can never be inferred to be part of the types of term definitions (vals, vars and defs). They are exclusively used temporarily during subtyping derivations.

Skolem types are stable types. A skolem type of the form ´∃ \alpha : T´ represents a stable reference to unknown value of type ´T´. The identifier ´\alpha´ is chosen uniquely every time a skolem type is created. However, as a skolem type is stable, it can be substituted in several occurrences in other types. When "copied" through substitution, all the copies retain the same ´\alpha´, and are therefore equivalent.

Methodic Types

TypeOrMethodic ::= Type | MethodicType MethodicType ::= MethodType | PolyType 

Methodic types are not real types. They are not part of the type lattice.

However, they share some meta-properties with types. In particular, when contained within other types that undertake some substitution, the substitution carries to the types within methodic types. It is therefore often convenient to think about them as types themselves.

Methodic types are used as the "declared type" of def definitions that have at least one term or type parameter list.

Method Types

MethodType ::= ‘(‘ MethodTypeParams ‘)‘ TypeOrMethodic MethodTypeParams ::= ε | MethodTypeParam {‘,‘ MethodTypeParam} MethodTypeParam ::= id ‘:‘ Type 

A method type is denoted internally as ´(\mathit{Ps})U´, where ´(\mathit{Ps})´ is a sequence of parameter names and types ´(p_1:T_1, ..., p_n:T_n)´ for some ´n \geq 0´ and ´U´ is a (value or method) type. This type represents named methods that take arguments named ´p_1, ..., p_n´ of types ´T_1, ..., T_n´ and that return a result of type ´U´.

Method types associate to the right: ´(\mathit{Ps}_1)(\mathit{Ps}_2)U´ is treated as ´(\mathit{Ps}_1)((\mathit{Ps}_2)U)´.

Method types do not exist as types of values. If a method name is used as a value, its type is implicitly converted to a corresponding function type.

Example

The definitions

def a: Int def b (x: Int): Boolean def c (x: Int) (y: String, z: String): String 

produce the typings

a: Int b: (Int) Boolean c: (Int) (String, String) String 

Polymorphic Method Types

PolyType ::= ‘[‘ PolyTypeParams ‘]‘ TypeOrMethodic PolyTypeParams ::= PolyTypeParam {‘,‘ PolyTypeParam} PolyTypeParam ::= ‘id‘ TypeBounds 

A polymorphic method type, or poly type for short, is denoted internally as [´\mathit{tps}\,´]´T´ where [´\mathit{tps}\,´] is a type parameter section [´a_1´ >: ´L_1´ <: ´U_1, ..., a_n´ >: ´L_n´ <: ´U_n´] for some ´n \geq 0´ and ´T´ is a (value or method) type. This type represents named methods that take type arguments ´S_1, ..., S_n´ which conform to the lower bounds ´L_1, ..., L_n´ and the upper bounds ´U_1, ..., U_n´ and that yield results of type ´T´.

Example

The definitions

def empty[A]: List[A] def union[A <: Comparable[A]] (x: Set[A], xs: Set[A]): Set[A] 

produce the typings

empty : [A >: Nothing <: Any] List[A] union : [A >: Nothing <: Comparable[A]] (x: Set[A], xs: Set[A]) Set[A] 

Operations on Types

This section defines a few meta-functions on types and methodic types.

• baseType(´T´, ´C´): computes the smallest type ´U´ of the form ´p´.´C´[´T_1, ..., T_n´] such that ´T <: U´.
• asSeenFrom(´T´, ´C´, ´p´): rebases the type ´T´ visible inside the class ´C´ "as seen from" the prefix ´p´.
• memberType(´T´, ´id´): finds a member of a type (T.id) and computes its underlying type or bounds.

These meta-functions are mutually recursive.

Base Type

The meta-function baseType(´T´, ´C´), where ´T´ is a proper type and ´C´ is a class identifier, computes the smallest type ´U´ of the form ´p.C´ or ´p.C´[´U_1, ..., U_n´] such that ´T <: U´. If no such type exists, the function is not defined. The main purpose of baseType is to substitute prefixes and class type parameters along the inheritance chain.

We define baseType(´T´, ´C´) as follows. For brevity, we write ´p.X´[´U_1, ..., U_n´] instead of ´p.X´ with ´n = 0´.

• baseType(´T = p.C´[´T_1, ..., T_n´], ´C´) ´≜ T´
• baseType(´p.D´[´T_1, ..., T_n´], ´C´) with ´D´ ≠ ´C ≜ \sigma W´ if ´Q´ is defined where
  • ´D´ is declared as ´D[\pm a_1 >: L_1 <: H_1, ..., \pm a_n >: L_n <: H_n]´ extends ´P_1, ..., P_m´
  • ´Q =´ meet(baseType(´P_i´, ´C´) for all ´i´ such that baseType(´P_i´, ´C´) is defined)
  • ´W = Q´ if ´p = \epsilon´ or if ´p´ is a package ref; otherwise, ´W =´ asSeenFrom(´Q´, ´D´, ´p´) (in that case, ´p´ is a stable type and ´D´ must be declared inside another class ´B´)
  • ´\sigma = [a_1 := T_1, ..., a_n := T_n]´ the substitution of the declared type parameters of ´D´ by the actual type arguments
• baseType(´T_1 ＆ T_2´, ´C´) ´≜´ meet´(´baseType(´T_1´, ´C´), baseType(´T_2´, ´C´)´)´
• baseType(´T_1 ｜ T_2´, ´C´) ´≜´ join´(´baseType(´T_1´, ´C´), baseType(´T_2´, ´C´)´)´
• baseType(´T´, ´C´) ´≜´ baseType(superType(´T´), ´C´) if superType(´T´) is defined

The definition above uses the following helper functions.

superType(´T´) computes the "next upper bound" of ´T´, if it exists:

• superType(´T´) where ´T´ is a stable type is its underlying type
• superType(´p.X´) where ´p.X´ is a non-class type designator is the upper bound of its underlying type definition
• superType(´([a_1 >: L_1 <: H_1, ..., a_n >: L_n <: H_n]´ =>> ´U)[T_1, ..., T_n]´) is ´[a_1 =: T_1, ..., a_n := T_n]U´ (i.e., the beta-reduction of the type lambda redex)
• superType(´T[T_1, ..., T_n]´) is superType(´T´)´[T_1, ..., T_n]´ if superType(´T´) is defined

Note that the cases of superType do not overlap with each other nor with any baseType case other than the superType-based one. The cases of baseType therefore do not overlap with each other either. That makes baseType an algorithmic partial function.

meet(´p.C[T_1, ..., T_n]´, ´q.C[U_1, ..., U_n]´) computes an intersection of two (parameterized) class types for the same class, and join computes a union:

• if ´p =:= q´ is false, then it is not defined
• otherwise, let ´W_i´ for ´i \in 1, ..., n´ be:
  • ´T_i ＆ U_i´ for meet (resp. ´T_i ｜ U_i´ for join) if the ´i´th type parameter of ´C´ is covariant
  • ´T_i ｜ U_i´ for meet (resp. ´T_i ＆ U_i´ for join) if the ´i´th type parameter of ´C´ is contravariant
  • ´T_i´ if ´T_i =:= U_i´ and the ´i´th type parameter of ´C´ is invariant
  • not defined otherwise
• if any of the ´W_i´ are not defined, the result is not defined
• otherwise, the result is ´p.C[W_1, ..., W_n]´

We generalize meet(´T_1, ..., T_n´) for a sequence as:

• not defined for ´n = 0´
• ´T_1´ if ´n = 1´
• meet(meet(´T_1, ..., T_{n-1}´), ´T_n´) if meet(´T_1, ..., T_{n-1}´) is defined
• not defined otherwise

Examples

Given the following definitions:

trait Iterable[+A] trait List[+A] extends Iterable[A] trait Map[K, +V] extends Iterable[(K, V)] trait Foo 

we have the following baseType results:

• baseType(List[Int], List) = List[Int]
• baseType(List[Int], Iterable) = Iterable[Int]
• baseType(List[A] & Iterable[B], Iterable) = meet(Iterable[A], Iterable[B]) = Iterable[A & B]
• baseType(List[A] & Foo, Iterable) = Iterable[A] (because baseType(Foo, Iterable) is not defined)
• baseType(Int, Iterable) is not defined
• baseType(Map[Int, String], Iterable) = Iterable[(Int, String)]
• baseType(Map[Int, String] & Map[String, String], Map) is not defined (because K is invariant)

As Seen From

The meta-function asSeenFrom(´T´, ´C´, ´p´), where ´T´ is a type or methodic type visible inside the class ´C´ and ´p´ is a stable type, rebases the type ´T´ "as seen from" the prefix ´p´. Essentially, it substitutes this-types and class type parameters in ´T´ to appropriate types visible from outside. Since T is visible inside ´C´, it can contain this-types and class type parameters of ´C´ itself as well as of all its enclosing classes. This-types of enclosing classes must be mapped to appropriate subprefixes of ´p´, while class type parameters must be mapped to appropriate concrete type arguments.

asSeenFrom(´T´, ´C´, ´p´) only makes sense if ´p´ has a base type for ´C´, i.e., if baseType(´p´, ´C´) is defined.

We define asSeenFrom(´T´, ´C´, ´p´) where baseType(´p´, ´C´) = ´q.C[U_1, ..., U_n]´ as follows:

• If ´T´ is a reference to the ´i´th class type parameter of some class ´D´:
  • If baseType(´p´, ´D´) ´= r.D[W_1, ..., W_m]´ is defined, then ´W_i´
  • Otherwise, if ´q = \epsilon´ or ´q´ is a package ref, then ´T´
  • Otherwise, ´q´ is a type, ´C´ must be defined in another class ´B´ and baseType(´q´, ´B´) must be defined, then asSeenFrom(´T´, ´B´, ´q´)
• Otherwise, if ´T´ is a this-type ´D´.this:
  • If ´D´ is a subclass of ´C´ and baseType(´p´, ´D´) is defined, then ´p´ (this is always the case when ´D = C´)
  • Otherwise, if ´q = \epsilon´ or ´q´ is a package ref, then ´T´
  • Otherwise, ´q´ is a type, ´C´ must be defined in another class ´B´ and baseType(´q´, ´B´) must be defined, then asSeenFrom(´T´, ´B´, ´q´)
• Otherwise, ´T´ where each if of its type components ´T_i´ is mapped to asSeenFrom(´T_i´, ´C´, ´p´).

For convenience, we generalize asSeenFrom to type definitions ´D´.

• If ´D´ is an alias ´= U´, then asSeenFrom(´D´, ´C´, ´p´) = asSeenFrom(´U´, ´C´, ´p´).
• If ´D´ is an abstract type definition with bounds ´>: L <: H´, then asSeenFrom(´D´, ´C´, ´p´) = ´>:´ asSeenFrom(´L´, ´C´, ´p´) ´<:´ asSeenFrom(´H´, ´C´, ´p´).

Member Type

The meta-function memberType(´T´, ´id´, ´p´), where ´T´ is a proper type, ´id´ is a term or type identifier, and ´p´ is a stable type, finds a member of a type (T.id) and computes its underlying type (for a term) or type definition (for a type) as seen from the prefix ´p´. For a term, it also computes whether the term is stable. memberType is the fundamental operation that computes the underlying type or underlying type definition of a named designator type.

The result ´M´ of a memberType is one of:

• undefined,
• a term result with underlying type or methodic type ´U´ and a stable flag,
• a class result with class ´C´, or
• a type result with underlying type definition ´D´.

As short-hand, we define memberType(´T´, ´id´) to be the same as memberType(´T´, ´id´, ´T´) when ´T´ is a stable type.

We define memberType(´T´, ´id´, ´p´) as follows:

• If ´T´ is a possibly parameterized class type of the form ´q.C[T_1, ..., T_n]´ (with ´n \geq 0´):
  • Let ´m´ be the class member of ´C´ with name ´id´.
  • If ´m´ is not defined, the result is undefined.
  • If ´m´ is a class definition, the result is a class result with class ´m´.
  • If ´m´ is a term definition in class ´D´ with declared type ´U´, the result is a term result with underlying type asSeenFrom(´U´, ´D´, ´p´) and stable flag true if and only if ´m´ is stable.
  • If ´m´ is a type member definition in class ´D´, the result is a type result with underlying type definition asSeenFrom(´U´, ´D´, ´p´) where ´U´ is defined as follows:
    • If ´m´ is an opaque type alias member definition with declared definition ´>: L <: H = V´, then
      • ´U´ is ´= V´ if ´p = D.´this or if we are computing memberType in a transparent mode,
      • ´U´ is ´>: L <: H´ otherwise.
    • ´U´ is the declared type definition of ´m´ otherwise.
• If ´T´ is another monomorphic type designator of the form ´q.X´:
  • Let ´U´ be memberType(´q´, ´X´)
  • Let ´H´ be the upper bound of ´U´
  • The result is memberType(´H´, ´id´, ´p´)
• If ´T´ is another parameterized type designator of the form ´q.X[T_1, ..., T_n]´ (with ´n \geq 0´):
  • Let ´U´ be memberType(´q´, ´X´)
  • Let ´H´ be the upper bound of ´U´
  • The result is memberType(´H[T_1, ..., T_n]´, ´id´, ´p´)
• If ´T´ is a parameterized type lambda of the form ´([\pm a_1 >: L_1 <: H_1, ..., \pm a_n >: L_n <: H_n]´ =>> ´U)[T_1, ..., T_n]´:
  • The result is memberType(´[a_1 := T_1, ..., a_n := T_n] U´, ´id´, ´p´), i.e., we beta-reduce the type redex.
• If ´T´ is a refined type of the form ´T_1´ { ´R´ }:
  • Let ´M_1´ be the result of memberType(´T_1´, ´id´, ´p´).
  • If the name of the refinement ´R´ is not ´id´, let ´M_2´ be undefined.
  • Otherwise, let ´M_2´ be the type or type definition of the refinement ´R´, as well as whether it is stable.
  • The result is mergeMemberType(´M_1´, ´M_2´).
• If ´T´ is a union type of the form ´T_1 ｜ T_2´:
  • Let ´J´ be the join of ´T´.
  • The result is memberType(´J´, ´id´, ´p´).
• If ´T´ is an intersection type of the form ´T_1 ＆ T_2´:
  • Let ´M_1´ be the result of memberType(´T_1´, ´id´, ´p´).
  • Let ´M_2´ be the result of memberType(´T_2´, ´id´, ´p´).
  • The result is mergeMemberType(´M_1´, ´M_2´).
• If ´T´ is a recursive type of the form { ´\alpha´ => ´T_1´ }:
  • The result is memberType(´T_1´, ´id´, ´p ´).
• If ´T´ is a stable type:
  • Let ´U´ be the underlying type of ´T´.
  • The result is memberType(´U´, ´id´, ´p´).
• Otherwise, the result is undefined.

We define the helper function mergeMemberType(´M_1´, ´M_2´) as:

• If either ´M_1´ or ´M_2´ is undefined, the result is the other one.
• Otherwise, if either ´M_1´ or ´M_2´ is a class result, the result is that one.
• Otherwise, ´M_1´ and ´M_2´ must either both be term results or both be type results.
  • If they are term results with underlying types ´U_1´ and ´U_2´ and stable flags ´s_1´ and ´s_2´, the result is a term result whose underlying type is meet(´U_1´, ´U_2´) and whose stable flag is ´s_1 \lor s_2´.
  • If they are type results with underlying type definitions ´D_1´ and ´D_2´, the result is a type result whose underlying type definition is intersect(´D_1´, ´D_2´).

Relations between types

We define the following relations between types.

Conformance

The conformance relation ´(<:)´ is the smallest relation such that ´S <: T´ is true if any of the following conditions hold. Note that the conditions are not all mutually exclusive.

• ´S = T´ (i.e., conformance is reflexive by definition).
• ´S´ is Nothing.
• ´T´ is AnyKind.
• ´S´ is a stable type with underlying type ´S_1´ and ´S_1 <: T´.
• ´S = p.x´ and ´T = q.x´ are term designators and
  • isSubPrefix(´p´, ´q´).
• ´S = p.X[S_1, ..., S_n]´ and ´T = q.X[T_1, ..., T_n]´ are possibly parameterized type designators with ´n \geq 0´ and:
  • isSubPrefix(´p´, ´q´), and
  • it is not the case that ´p.x´ and ´q.X´ are class type designators for different classes, and
  • for each ´i \in { 1, ..., n }´:
    • the ´i´th type parameter of ´q.X´ is covariant and ´S_i <: T_i´ ³, or
    • the ´i´th type parameter of ´q.X´ is contravariant and ´T_i <: S_i´ ³, or
    • the ´i´th type parameter of ´q.X´ is invariant and:
      • ´S_i´ and ´T_i´ are types and ´S_i =:= T_i´, or
      • ´S_i´ is a type and ´T_i´ is a wildcard type argument of the form ´? >: L_2 <: H_2´ and ´L_2 <: S_i´ and ´S_i <: H_2´, or
      • ´S_i´ is a wildcard type argument of the form ´? >: L_1 <: H_1´ and ´T_i´ is a wildcard type argument of the form ´? >: L_2 <: H_2´ and ´L_2 <: L_1´ and ´H_1 <: H_2´ (i.e., the ´S_i´ "interval" is contained in the ´T_i´ "interval").
• ´T = q.C[T_1, ..., T_n]´ with ´n \geq 0´ and baseType(´S´, ´C´) is defined and baseType(´S´, ´C´) ´<: T´.
• ´S = p.X[S_1, ..., S_n]´ and ´p.X´ is non-class type designator and ´H <: T´ where ´H´ is the upper bound of the underlying type definition of ´p.X´.
• ´S = p.C´ and ´T = C´.this and ´C´ is the hidden class of an object and:
  • ´p = \epsilon´ or ´p´ is a package ref, or
  • isSubPrefix(´p´, ´D´.this) where ´D´ is the enclosing class of ´C´.
• ´S = C´.this and ´T = q.C´ and ´C´ is the hidden class of an object and:
  • either ´q = \epsilon´ or ´q´ is a package ref, or
  • isSubPrefix(´D´.this, ´q´) where ´D´ is the enclosing class of ´C´.
• ´S = S_1 ｜ S_2´ and ´S_1 <: T´ and ´S_2 <: T´.
• ´T = T_1 ｜ T_2´ and either ´S <: T_1´ or ´S <: T_2´.
• ´T = T_1 ＆ T_2´ and ´S <: T_1´ and ´S <: T_2´.
• ´S = S_1 ＆ S_2´ and either ´S_1 <: T´ or ´S_2 <: T´.
• ´S = S_1´ @a and ´S_1 <: T´.
• ´T = T_1´ @a and ´S <: T_1´ (i.e., annotations can be dropped).
• ´T = q.X´ and ´q.X´ is a non-class type designator and ´S <: L´ where ´L´ is the lower bound of the underlying type definition of ´q.X´.
• ´S = p.X´ and ´p.X´ is a non-class type designator and ´H <: T´ where ´H´ is the upper bound of the underlying type definition of ´p.X´.
• ´S = [\pm a_1 >: L_1 <: H_1, ..., \pm a_n >: L_n <: H_n]´ =>> ´S_1´ and ´T = [\pm b_1 >: M_1 <: G_1, ..., \pm b_n >: M_n <: G_n]´ =>> ´T_1´, and given ´\sigma = [b_1 := a_1, ..., b_n := a_n]´:
  • ´S_1 <: \sigma T_1´, and
  • for each ´i \in { 1, ..., n }´:
    • the variance of ´a_i´ conforms to the variance of ´b_i´ (´+´ conforms to ´+´ and ´\epsilon´, ´-´ conforms to ´-´ and ´\epsilon´, and ´\epsilon´ conforms to ´\epsilon´), and
    • ´\sigma (>: M_i <: G_i)´ is contained in ´>: L_i <: H_i´ (i.e., ´L_i <: \sigma M_i´ and ´\sigma G_i <: H_i´).
• ´S = p.X´ and ´T = [\pm b_1 >: M_1 <: G_1, ..., \pm b_n >: M_n <: G_n]´ =>> ´T_1´ and ´S´ is a type constructor with ´n´ type parameters and:
  • ´([\pm a_1 >: L_1 <: H_1, ..., \pm a_n >: L_n <: H_n]´ =>> ´S[a_1, ..., a_n]) <: T´ where the ´a_i´ are copies of the type parameters of ´S´ (i.e., we can eta-expand ´S´ to compare it to a type lambda).
• ´T = T_1´ { ´R´ } and ´S <: T_1´ and, given ´p = S´ if ´S´ is a stable type and ´p = ∃ \alpha : S´ otherwise:
  • ´R =´ type ´X >: L <: H´ and memberType(´p´, ´X´) is a class result for ´C´ and ´L <: p.C´ and ´p.C <: H´, or
  • ´R =´ type ´X >: L_2 <: H_2´ and memberType(´p´, ´X´) is a type result with bounds ´>: L_1 <: H_1´ and ´L_2 <: L_1´ and ´H_1 <: H_2´, or
  • ´R =´ val ´X: T_2´ and memberType(´p´, ´X´) is a stable term result with type ´S_2´ and ´S_2 <: T_2´, or
  • ´R =´ def ´X: T_2´ and memberType(´p´, ´X´) is a term result with type ´S_2´ and ´T_2´ is a type and ´S_2 <: T_2´, or
  • ´R =´ def ´X: T_2´ and memberType(´p´, ´X´) is a term result with methodic type ´S_2´ and ´T_2´ is a methodic type and matches(´S_2´, ´T_2´).
• ´S = S_1´ { ´R´ } and ´S_1 <: T´.
• ´S =´ { ´\alpha´ => ´S_1´ } and ´T =´ { ´\beta´ => ´T_1´ } and ´S_1 <: [\beta := \alpha]T_1´.
• ´T =´ { ´\beta´ => ´T_1´ } and ´S´ is a proper type but not a recursive type and ´p' <: [\beta := p]T_1´ where:
  • ´p´ is ´S´ if ´S´ is a stable type and ´∃ \alpha : S´ otherwise, and
  • ´p'´ is the result of replacing any top-level recursive type { ´\gamma´ => ´Z´ } in ´p´ with ´[\gamma := p]Z´ (TODO specify this better).
• ´(X´ match <: ´H´ { ... }´) <: T´ if ´X´ match ... reduces to ´S_1´ and ´S_1 <: T´
• ´S <: (X´ match <: ´H´ { ... }´)´ if ´X´ match ... reduces to ´T_1´ and ´S <: T_1´
• ´(X´ match <: ´H´ { ... }´) <: T´ if ´H <: T´
• ´(X´ match <: ´H_X´ { case ´P_1´ => ´A_1´; ...; case ´P_n´ => ´A_n´ }´) <: (Y´ match <: ´H_Y´ { case ´Q_1´ => ´B_1´; ...; ´Q_n´ => ´B_n´ }´)´ if ´X =:= Y´ and ´P_i =:= Q_i´ for each ´i´ and ´A_i <: B_i´ for each ´i´
• ´S = (´=> ´S_1)´ and ´T = (´=> ´T_1)´ and ´S_1 <: T_1´.
• ´S =´ scala.Null and:
  • ´T = q.C[T_1, ..., T_n]´ with ´n \geq 0´ and ´C´ does not derive from scala.AnyVal and ´C´ is not the hidden class of an object, or
  • ´T = q.x´ is a term designator with underlying type ´U´ and scala.Null ´<: U´, or
  • ´T = T_1´ { ´R´ } and scala.Null ´<: T_1´, or
  • ´T =´ { ´\beta´ => ´T_1´ } and scala.Null ´<: T_1´.
• ´S´ is a stable type and ´T = q.x´ is a term designator with underlying type ´T_1´ and ´T_1´ is a stable type and ´S <: T_1´.
• ´S = S_1´ { ´R´ } and ´S_1 <: T´.
• ´S =´ { ´\alpha´ => ´S_1´ } and ´S_1 <: T´.
• ´T =´ scala.Tuple´_n[T_1, ..., T_n]´ with ´1 \leq n \leq 22´, and ´S <: T_1´ *: ... *: ´T_n´ *: scala.EmptyTuple.

We define isSubPrefix(´p´, ´q´) where ´p´ and ´q´ are prefixes as:

• If both ´p´ and ´q´ are types, then ´p <: q´.
• Otherwise, ´p = q´ (for empty prefixes and package refs).

We define matches(´S´, ´T´) where ´S´ and ´T´ are types or methodic types as:

• If ´S´ and ´T´ are types, then ´S <: T´.
• If ´S´ and ´T´ are method types ´(a_1: S_1, ..., a_n: S_n)S'´ and ´(b_1: T_1, ..., b_n: T_n)T'´, then ´\sigma S_i =:= T_i´ for each ´i´ and matches(´\sigma S'´, ´T'´), where ´\sigma = [a_1 := b_1, ..., a_n := b_n]´.
• If ´S´ and ´T´ are poly types ´[a_1 >: L_{s1} <: H_{s1}, ..., a_n >: L_{sn} <: H_{sn}]S'´ and ´[b_1 >: L_{t1} <: H_{t1}, ..., b_n >: L_{tn} <: H_{tn}]T'´, then ´\sigma L_{si} =:= L_{ti}´ and ´\sigma H_{si} =:= H_{ti}´ for each ´i´ and matches(´\sigma S'´, ´T'´), where ´\sigma = [a_1 := b_1, ..., a_n := b_n]´.

Note that conformance in Scala is not transitive. Given two abstract types ´A´ and ´B´, and one abstract type ´C >: A <: B´ available on prefix ´p´, we have ´A <: p.C´ and ´C <: p.B´ but not necessarily ´A <: B´.

Least upper bounds and greatest lower bounds

The ´(<:)´ relation forms pre-order between types, i.e. it is transitive and reflexive. This allows us to define least upper bounds and greatest lower bounds of a set of types in terms of that order.

• the least upper bound of A and B is the smallest type L such that A <: L and B <: L.
• the greatest lower bound of A and B is the largest type G such that G <: A and G <: B.

By construction, for all types A and B, the least upper bound of A and B is A ｜ B, and their greatest lower bound is A ＆ B.

Equivalence

Equivalence is defined as mutual conformance.

´S =:= T´ if and only if both ´S <: T´ and ´T <: S´.

Weak Conformance

In some situations Scala uses a more general conformance relation. A type ´S´ weakly conforms to a type ´T´, written ´S <:_w T´, if ´S <: T´ or both ´S´ and ´T´ are primitive number types and ´S´ precedes ´T´ in the following ordering.

Byte ´<:_w´ Short Short ´<:_w´ Int Char ´<:_w´ Int Int ´<:_w´ Long Long ´<:_w´ Float Float ´<:_w´ Double 

A weak least upper bound is a least upper bound with respect to weak conformance.

Compatibility

A type ´T´ is compatible to a type ´U´ if ´T´ (or its corresponding function type) weakly conforms to ´U´ after applying eta-expansion. If ´T´ is a method type, it's converted to the corresponding function type. If the types do not weakly conform, the following alternatives are checked in order:

• dropping by-name modifiers: if ´U´ is of the shape ´=> U'´ (and ´T´ is not), ´T <:_w U'´;
• SAM conversion: if ´T´ corresponds to a function type, and ´U´ declares a single abstract method whose type corresponds to the function type ´U'´, ´T <:_w U'´.
• implicit conversion: there's an implicit conversion from ´T´ to ´U´ in scope;

Examples

Function compatibility via SAM conversion

Given the definitions

def foo(x: Int => String): Unit def foo(x: ToString): Unit trait ToString { def convert(x: Int): String } 

The application foo((x: Int) => x.toString)resolves to the first overload, as it's more specific:

• Int => String is compatible to ToString -- when expecting a value of type ToString, you may pass a function literal from Int to String, as it will be SAM-converted to said function;
• ToString is not compatible to Int => String -- when expecting a function from Int to String, you may not pass a ToString.

Realizability

A type ´T´ is realizable if and only if it is inhabited by non-null values. It is defined as:

• A term designator ´p.x´ with underlying type ´U´ is realizable if ´p´ is ´\epsilon´ or a package ref or a realizable type and
  • memberType(´p´, ´x´) has the stable flag, or
  • the type returned by memberType(´p´, ´x´) is realizable.
• A stable type that is not a term designator is realizable.
• Another type ´T´ is realizable if
  • ´T´ is concrete, and
  • ´T´ has good bounds.

A concrete type ´T´ has good bounds if all of the following apply:

• all its non-class type members have good bounds, i.e., their bounds ´L´ and ´H´ are such that ´L <: H´,
• all its type refinements have good bounds, and
• for all base classes ´C´ of ´T´:
  • baseType(´T´, ´C´) is defined with some result ´p.C[T_1, ..., T_n]´, and
  • for all ´i \in { 1, ..., n }´, ´T_i´ is a real type or (when it is a wildcard type argument) it has good bounds.

Note: it is possible for baseType(´T´, ´C´) not to be defined because of the meet computation, which may fail to merge prefixes and/or invariant type arguments.

Type Erasure

A type is called generic if it contains type arguments or type variables. Type erasure is a mapping from (possibly generic) types to non-generic types. We write ´|T|´ for the erasure of type ´T´. The erasure mapping is defined as follows. Internal computations are performed in a transparent mode, which has an effect on how memberType behaves for opaque type aliases.

• The erasure of AnyKind is Object.
• The erasure of a non-class type designator is the erasure of its underlying upper bound.
• The erasure of a term designator is the erasure of its underlying type.
• The erasure of the parameterized type scala.Array´[T_1]´ is scala.Array´[|T_1|]´.
• The erasure of every other parameterized type ´T[T_1, ..., T_n]´ is ´|T|´.
• The erasure of a stable type ´p´ is the erasure of the underlying type of ´p´.
• The erasure of a by-name type => ´T_1´ is scala.Function0.
• The erasure of an annotated type ´T_1 a´ is ´|T_1|´.
• The erasure of a refined type ´T_1´ { ´R´ } is ´|T_1|´.
• The erasure of a recursive type { ´\alpha´ => ´T_1´ } and the associated recursive this type ´\alpha´ is ´|T_1|´.
• The erasure of a union type ´S ｜ T´ is the erased least upper bound (elub) of the erasures of ´S´ and ´T´.
• The erasure of an intersection type ´S ＆ T´ is the eglb (erased greatest lower bound) of the erasures of ´S´ and ´T´.

The erased LUB is computed as follows:

• if both argument are arrays of objects, an array of the erased LUB of the element types
• if both arguments are arrays of same primitives, an array of this primitive
• if one argument is array of primitives and the other is array of objects, Object
• if one argument is an array, Object
• otherwise a common superclass or trait S of the argument classes, with the following two properties:
  • S is minimal: no other common superclass or trait derives from S, and
  • S is last: in the linearization of the first argument type ´|A|´ there are no minimal common superclasses or traits that come after S. The reason to pick last is that we prefer classes over traits that way, which leads to more predictable bytecode and (?) faster dynamic dispatch.

The rules for ´eglb(A, B)´ are given below in pseudocode:

eglb(scala.Array[A], JArray[B]) = scala.Array[eglb(A, B)] eglb(scala.Array[T], _) = scala.Array[T] eglb(_, scala.Array[T]) = scala.Array[T] eglb(A, B) = A if A extends B eglb(A, B) = B if B extends A eglb(A, _) = A if A is not a trait eglb(_, B) = B if B is not a trait eglb(A, _) = A // use first