Expressions =========== ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} Expr ::= (Bindings | id | `_') `=>' Expr | Expr1 Expr1 ::= `if' `(' Expr `)' {nl} Expr [[semi] else Expr] | `while' `(' Expr `)' {nl} Expr | `try' `{' Block `}' [`catch' `{' CaseClauses `}'] [`finally' Expr] | `do' Expr [semi] `while' `(' Expr ')' | `for' (`(' Enumerators `)' | `{' Enumerators `}') {nl} [`yield'] Expr | `throw' Expr | `return' [Expr] | [SimpleExpr `.'] id `=' Expr | SimpleExpr1 ArgumentExprs `=' Expr | PostfixExpr | PostfixExpr Ascription | PostfixExpr `match' `{' CaseClauses `}' PostfixExpr ::= InfixExpr [id [nl]] InfixExpr ::= PrefixExpr | InfixExpr id [nl] InfixExpr PrefixExpr ::= [`-' | `+' | `~' | `!'] SimpleExpr SimpleExpr ::= `new' (ClassTemplate | TemplateBody) | BlockExpr | SimpleExpr1 [`_'] SimpleExpr1 ::= Literal | Path | `_' | `(' [Exprs] `)' | SimpleExpr `.' id s | SimpleExpr TypeArgs | SimpleExpr1 ArgumentExprs | XmlExpr Exprs ::= Expr {`,' Expr} BlockExpr ::= `{' CaseClauses `}' | `{' Block `}' Block ::= {BlockStat semi} [ResultExpr] ResultExpr ::= Expr1 | (Bindings | ([`implicit'] id | `_') `:' CompoundType) `=>' Block Ascription ::= `:' InfixType | `:' Annotation {Annotation} | `:' `_' `*' ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Expressions are composed of operators and operands. Expression forms are discussed subsequently in decreasing order of precedence. Expression Typing ----------------- The typing of expressions is often relative to some _expected type_ (which might be undefined). When we write ``expression $e$ is expected to conform to type $T$'', we mean: (1) the expected type of $e$ is $T$, and (2) the type of expression $e$ must conform to $T$. The following skolemization rule is applied universally for every expression: If the type of an expression would be an existential type $T$, then the type of the expression is assumed instead to be a [skolemization](#existential-types) of $T$. Skolemization is reversed by type packing. Assume an expression $e$ of type $T$ and let $t_1[\mathit{tps}_1] >: L_1 <: U_1 , \ldots , t_n[\mathit{tps}_n] >: L_n <: U_n$ be all the type variables created by skolemization of some part of $e$ which are free in $T$. Then the _packed type_ of $e$ is ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} $T$ forSome { type $t_1[\mathit{tps}_1] >: L_1 <: U_1$; $\ldots$; type $t_n[\mathit{tps}_n] >: L_n <: U_n$ }. ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Literals -------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} SimpleExpr ::= Literal ~~~~~~~~~~~~~~~~~~~~~~~~~~~ Typing of literals is as described [here](#literals); their evaluation is immediate. The _Null_ Value -------------------- The `null` value is of type `scala.Null`, and is thus compatible with every reference type. It denotes a reference value which refers to a special “`null`” object. This object implements methods in class `scala.AnyRef` as follows: - `eq($x\,$)` and `==($x\,$)` return `true` iff the argument $x$ is also the ``null'' object. - `ne($x\,$)` and `!=($x\,$)` return true iff the argument x is not also the ``null'' object. - `isInstanceOf[$T\,$]` always returns `false`. - `asInstanceOf[$T\,$]` returns the ``null'' object itself if $T$ conforms to `scala.AnyRef`, and throws a `NullPointerException` otherwise. A reference to any other member of the ``null'' object causes a `NullPointerException` to be thrown. Designators ----------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} SimpleExpr ::= Path | SimpleExpr `.' id ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ A designator refers to a named term. It can be a _simple name_ or a _selection_. A simple name $x$ refers to a value as specified [here](#identifiers-names-and-scopes). If $x$ is bound by a definition or declaration in an enclosing class or object $C$, it is taken to be equivalent to the selection `$C$.this.$x$` where $C$ is taken to refer to the class containing $x$ even if the type name $C$ is [shadowed](#identifiers-names-and-scopes) at the occurrence of $x$. If $r$ is a [stable identifier](#paths) of type $T$, the selection $r.x$ refers statically to a term member $m$ of $r$ that is identified in $T$ by the name $x$. For other expressions $e$, $e.x$ is typed as if it was `{ val $y$ = $e$; $y$.$x$ }`, for some fresh name $y$. The expected type of a designator's prefix is always undefined. The type of a designator is the type $T$ of the entity it refers to, with the following exception: The type of a [path](#paths) $p$ which occurs in a context where a [stable type](#singleton-types) is required is the singleton type `$p$.type`. The contexts where a stable type is required are those that satisfy one of the following conditions: #. The path $p$ occurs as the prefix of a selection and it does not designate a constant, or #. The expected type $\mathit{pt}$ is a stable type, or #. The expected type $\mathit{pt}$ is an abstract type with a stable type as lower bound, and the type $T$ of the entity referred to by $p$ does not conform to $\mathit{pt}$, or #. The path $p$ designates a module. The selection $e.x$ is evaluated by first evaluating the qualifier expression $e$, which yields an object $r$, say. The selection's result is then the member of $r$ that is either defined by $m$ or defined by a definition overriding $m$. If that member has a type which conforms to `scala.NotNull`, the member's value must be initialized to a value different from `null`, otherwise a `scala.UnitializedError` is thrown. This and Super -------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} SimpleExpr ::= [id `.'] `this' | [id '.'] `super' [ClassQualifier] `.' id ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The expression `this` can appear in the statement part of a template or compound type. It stands for the object being defined by the innermost template or compound type enclosing the reference. If this is a compound type, the type of `this` is that compound type. If it is a template of a class or object definition with simple name $C$, the type of this is the same as the type of `$C$.this`. The expression `$C$.this` is legal in the statement part of an enclosing class or object definition with simple name $C$. It stands for the object being defined by the innermost such definition. If the expression's expected type is a stable type, or `$C$.this` occurs as the prefix of a selection, its type is `$C$.this.type`, otherwise it is the self type of class $C$. A reference `super.$m$` refers statically to a method or type $m$ in the least proper supertype of the innermost template containing the reference. It evaluates to the member $m'$ in the actual supertype of that template which is equal to $m$ or which overrides $m$. The statically referenced member $m$ must be a type or a method. If it is a method, it must be concrete, or the template containing the reference must have a member $m'$ which overrides $m$ and which is labeled `abstract override`. A reference `$C$.super.$m$` refers statically to a method or type $m$ in the least proper supertype of the innermost enclosing class or object definition named $C$ which encloses the reference. It evaluates to the member $m'$ in the actual supertype of that class or object which is equal to $m$ or which overrides $m$. The statically referenced member $m$ must be a type or a method. If the statically referenced member $m$ is a method, it must be concrete, or the innermost enclosing class or object definition named $C$ must have a member $m'$ which overrides $m$ and which is labeled `abstract override`. The `super` prefix may be followed by a trait qualifier `[$T\,$]`, as in `$C$.super[$T\,$].$x$`. This is called a _static super reference_. In this case, the reference is to the type or method of $x$ in the parent trait of $C$ whose simple name is $T$. That member must be uniquely defined. If it is a method, it must be concrete. (@super) Consider the following class definitions ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} class Root { def x = "Root" } class A extends Root { override def x = "A" ; def superA = super.x } trait B extends Root { override def x = "B" ; def superB = super.x } class C extends Root with B { override def x = "C" ; def superC = super.x } class D extends A with B { override def x = "D" ; def superD = super.x } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The linearization of class `C` is `{C, B, Root}` and the linearization of class `D` is `{D, B, A, Root}`. Then we have: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} (new A).superA == "Root", (new C).superB = "Root", (new C).superC = "B", (new D).superA == "Root", (new D).superB = "A", (new D).superD = "B", ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Note that the `superB` function returns different results depending on whether `B` is mixed in with class `Root` or `A`. Function Applications --------------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} SimpleExpr ::= SimpleExpr1 ArgumentExprs ArgumentExprs ::= `(' [Exprs] `)' | `(' [Exprs `,'] PostfixExpr `:' `_' `*' ')' | [nl] BlockExpr Exprs ::= Expr {`,' Expr} ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ An application `$f$($e_1 , \ldots , e_m$)` applies the function $f$ to the argument expressions $e_1 , \ldots , e_m$. If $f$ has a method type `($p_1$:$T_1 , \ldots , p_n$:$T_n$)$U$`, the type of each argument expression $e_i$ is typed with the corresponding parameter type $T_i$ as expected type. Let $S_i$ be type type of argument $e_i$ $(i = 1 , \ldots , m)$. If $f$ is a polymorphic method, [local type inference](#local-type-inference) is used to determine type arguments for $f$. If $f$ has some value type, the application is taken to be equivalent to `$f$.apply($e_1 , \ldots , e_m$)`, i.e.\ the application of an `apply` method defined by $f$. The function $f$ must be _applicable_ to its arguments $e_1 , \ldots , e_n$ of types $S_1 , \ldots , S_n$. If $f$ has a method type $(p_1:T_1 , \ldots , p_n:T_n)U$ we say that an argument expression $e_i$ is a _named_ argument if it has the form $x_i=e'_i$ and $x_i$ is one of the parameter names $p_1 , \ldots , p_n$. The function $f$ is applicable if all of the follwing conditions hold: - For every named argument $x_i=e'_i$ the type $S_i$ is compatible with the parameter type $T_j$ whose name $p_j$ matches $x_i$. - For every positional argument $e_i$ the type $S_i$ is compatible with $T_i$. - If the expected type is defined, the result type $U$ is compatible to it. If $f$ is a polymorphic method it is applicable if [local type inference](#local-type-inference) can determine type arguments so that the instantiated method is applicable. If $f$ has some value type it is applicable if it has a method member named `apply` which is applicable. Evaluation of `$f$($e_1 , \ldots , e_n$)` usually entails evaluation of $f$ and $e_1 , \ldots , e_n$ in that order. Each argument expression is converted to the type of its corresponding formal parameter. After that, the application is rewritten to the function's right hand side, with actual arguments substituted for formal parameters. The result of evaluating the rewritten right-hand side is finally converted to the function's declared result type, if one is given. The case of a formal parameter with a parameterless method type `=>$T$` is treated specially. In this case, the corresponding actual argument expression $e$ is not evaluated before the application. Instead, every use of the formal parameter on the right-hand side of the rewrite rule entails a re-evaluation of $e$. In other words, the evaluation order for `=>`-parameters is _call-by-name_ whereas the evaluation order for normal parameters is _call-by-value_. Furthermore, it is required that $e$'s [packed type](#expression-typing) conforms to the parameter type $T$. The behavior of by-name parameters is preserved if the application is transformed into a block due to named or default arguments. In this case, the local value for that parameter has the form `val $y_i$ = () => $e$` and the argument passed to the function is `$y_i$()`. The last argument in an application may be marked as a sequence argument, e.g.\ `$e$: _*`. Such an argument must correspond to a [repeated parameter](#repeated-parameters) of type `$S$*` and it must be the only argument matching this parameter (i.e.\ the number of formal parameters and actual arguments must be the same). Furthermore, the type of $e$ must conform to `scala.Seq[$T$]`, for some type $T$ which conforms to $S$. In this case, the argument list is transformed by replacing the sequence $e$ with its elements. When the application uses named arguments, the vararg parameter has to be specified exactly once. A function application usually allocates a new frame on the program's run-time stack. However, if a local function or a final method calls itself as its last action, the call is executed using the stack-frame of the caller. (@) Assume the following function which computes the sum of a variable number of arguments: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} def sum(xs: Int*) = (0 /: xs) ((x, y) => x + y) ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Then ~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} sum(1, 2, 3, 4) sum(List(1, 2, 3, 4): _*) ~~~~~~~~~~~~~~~~~~~~~~~~~~~ both yield `10` as result. On the other hand, ~~~~~~~~~~~~~~~~~~~~~~ {.scala} sum(List(1, 2, 3, 4)) ~~~~~~~~~~~~~~~~~~~~~~ would not typecheck. ### Named and Default Arguments If an application might uses named arguments $p = e$ or default arguments, the following conditions must hold. - The named arguments form a suffix of the argument list $e_1 , \ldots , e_m$, i.e.\ no positional argument follows a named one. - The names $x_i$ of all named arguments are pairwise distinct and no named argument defines a parameter which is already specified by a positional argument. - Every formal parameter $p_j:T_j$ which is not specified by either a positional or a named argument has a default argument. If the application uses named or default arguments the following transformation is applied to convert it into an application without named or default arguments. If the function $f$ has the form `$p.m$[$\mathit{targs}$]` it is transformed into the block ~~~~~~~~~~~~~~~~~~ {.scala} { val q = $p$ q.$m$[$\mathit{targs}$] } ~~~~~~~~~~~~~~~~~~ If the function $f$ is itself an application expression the transformation is applied recursively on $f$. The result of transforming $f$ is a block of the form ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} { val q = $p$ val $x_1$ = expr$_1$ $\ldots$ val $x_k$ = expr$_k$ q.$m$[$\mathit{targs}$]($\mathit{args}_1$)$, \ldots ,$($\mathit{args}_l$) } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ where every argument in $(\mathit{args}_1) , \ldots , (\mathit{args}_l)$ is a reference to one of the values $x_1 , \ldots , x_k$. To integrate the current application into the block, first a value definition using a fresh name $y_i$ is created for every argument in $e_1 , \ldots , e_m$, which is initialised to $e_i$ for positional arguments and to $e'_i$ for named arguments of the form `$x_i=e'_i$`. Then, for every parameter which is not specified by the argument list, a value definition using a fresh name $z_i$ is created, which is initialized using the method computing the [default argument](#function-declarations-and-definitions) of this parameter. Let $\mathit{args}$ be a permutation of the generated names $y_i$ and $z_i$ such such that the position of each name matches the position of its corresponding parameter in the method type `($p_1:T_1 , \ldots , p_n:T_n$)$U$`. The final result of the transformation is a block of the form ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} { val q = $p$ val $x_1$ = expr$_1$ $\ldots$ val $x_l$ = expr$_k$ val $y_1$ = $e_1$ $\ldots$ val $y_m$ = $e_m$ val $z_1$ = q.$m$\$default\$i[$\mathit{targs}$]($\mathit{args}_1$)$, \ldots ,$($\mathit{args}_l$) $\ldots$ val $z_d$ = q.$m$\$default\$j[$\mathit{targs}$]($\mathit{args}_1$)$, \ldots ,$($\mathit{args}_l$) q.$m$[$\mathit{targs}$]($\mathit{args}_1$)$, \ldots ,$($\mathit{args}_l$)($\mathit{args}$) } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Method Values ------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} SimpleExpr ::= SimpleExpr1 `_' ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The expression ~`$e$ _`~ is well-formed if $e$ is of method type or if $e$ is a call-by-name parameter. If $e$ is a method with parameters, `$e$ _`~ represents $e$ converted to a function type by [eta expansion](#eta-expansion). If $e$ is a parameterless method or call-by-name parameter of type `=>$T$`, `$e$ _`~ represents the function of type `() => $T$`, which evaluates $e$ when it is applied to the empty parameterlist `()`. (@) The method values in the left column are each equivalent to the [anonymous functions](#anonymous-functions) on their right. ------------------------------ ----------------------------------------------------- `Math.sin _`{.scala} `x => Math.sin(x)`{.scala} `Array.range _`{.scala} `(x1, x2) => Array.range(x1, x2)`{.scala} `List.map2 _`{.scala} `(x1, x2) => (x3) => List.map2(x1, x2)(x3)`{.scala} `List.map2(xs, ys)_`{.scala} `x => List.map2(xs, ys)(x)`{.scala} ------------------------------ ----------------------------------------------------- Note that a space is necessary between a method name and the trailing underscore because otherwise the underscore would be considered part of the name. Type Applications ----------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} SimpleExpr ::= SimpleExpr TypeArgs ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ A type application `$e$[$T_1 , \ldots , T_n$]` instantiates a polymorphic value $e$ of type `[$a_1$ >: $L_1$ <: $U_1, \ldots , a_n$ >: $L_n$ <: $U_n$]$S$` with argument types `$T_1 , \ldots , T_n$`. Every argument type $T_i$ must obey the corresponding bounds $L_i$ and $U_i$. That is, for each $i = 1 , \ldots , n$, we must have $\sigma L_i <: T_i <: \sigma U_i$, where $\sigma$ is the substitution $[a_1 := T_1 , \ldots , a_n := T_n]$. The type of the application is $\sigma S$. If the function part $e$ is of some value type, the type application is taken to be equivalent to `$e$.apply[$T_1 , \ldots ,$ T$_n$]`, i.e.\ the application of an `apply` method defined by $e$. Type applications can be omitted if [local type inference](#local-type-inference) can infer best type parameters for a polymorphic functions from the types of the actual function arguments and the expected result type. Tuples ------ ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} SimpleExpr ::= `(' [Exprs] `)' ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ A tuple expression `($e_1 , \ldots , e_n$)` is an alias for the class instance creation `scala.Tuple$n$($e_1 , \ldots , e_n$)`, where $n \geq 2$. The empty tuple `()` is the unique value of type `scala.Unit`. Instance Creation Expressions ----------------------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} SimpleExpr ::= `new' (ClassTemplate | TemplateBody) ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ A simple instance creation expression is of the form `new $c$` where $c$ is a [constructor invocation](#constructor-invocations). Let $T$ be the type of $c$. Then $T$ must denote a (a type instance of) a non-abstract subclass of `scala.AnyRef`. Furthermore, the _concrete self type_ of the expression must conform to the [self type](#templates) of the class denoted by $T$. The concrete self type is normally $T$, except if the expression `new $c$` appears as the right hand side of a value definition ~~~~~~~~~~~~~~~~~~~~~~~ {.scala} val $x$: $S$ = new $c$ ~~~~~~~~~~~~~~~~~~~~~~~ (where the type annotation `: $S$` may be missing). In the latter case, the concrete self type of the expression is the compound type `$T$ with $x$.type`. The expression is evaluated by creating a fresh object of type $T$ which is is initialized by evaluating $c$. The type of the expression is $T$. A general instance creation expression is of the form `new $t$` for some [class template](#templates) $t$. Such an expression is equivalent to the block ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} { class $a$ extends $t$; new $a$ } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ where $a$ is a fresh name of an _anonymous class_ which is inaccessible to user programs. There is also a shorthand form for creating values of structural types: If `{$D$}` is a class body, then `new {$D$}` is equivalent to the general instance creation expression `new AnyRef{$D$}`. (@) Consider the following structural instance creation expression: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} new { def getName() = "aaron" } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ This is a shorthand for the general instance creation expression ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} new AnyRef{ def getName() = "aaron" } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The latter is in turn a shorthand for the block ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} { class anon\$X extends AnyRef{ def getName() = "aaron" }; new anon\$X } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ where `anon\$X` is some freshly created name. Blocks ------ ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} BlockExpr ::= `{' Block `}' Block ::= {BlockStat semi} [ResultExpr] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ A block expression `{$s_1$; $\ldots$; $s_n$; $e\,$}` is constructed from a sequence of block statements $s_1 , \ldots , s_n$ and a final expression $e$. The statement sequence may not contain two definitions or declarations that bind the same name in the same namespace. The final expression can be omitted, in which case the unit value `()` is assumed. The expected type of the final expression $e$ is the expected type of the block. The expected type of all preceding statements is undefined. The type of a block `$s_1$; $\ldots$; $s_n$; $e$` is `$T$ forSome {$\,Q\,$}`, where $T$ is the type of $e$ and $Q$ contains [existential clauses](#existential-types) for every value or type name which is free in $T$ and which is defined locally in one of the statements $s_1 , \ldots , s_n$. We say the existential clause _binds_ the occurrence of the value or type name. Specifically, - A locally defined type definition `type$\;t = T$` is bound by the existential clause `type$\;t >: T <: T$`. It is an error if $t$ carries type parameters. - A locally defined value definition~ `val$\;x: T = e$` is bound by the existential clause `val$\;x: T$`. - A locally defined class definition `class$\;c$ extends$\;t$` is bound by the existential clause `type$\;c <: T$` where $T$ is the least class type or refinement type which is a proper supertype of the type $c$. It is an error if $c$ carries type parameters. - A locally defined object definition `object$\;x\;$extends$\;t$` is bound by the existential clause `val$\;x: T$` where $T$ is the least class type or refinement type which is a proper supertype of the type `$x$.type`. Evaluation of the block entails evaluation of its statement sequence, followed by an evaluation of the final expression $e$, which defines the result of the block. (@) Assuming a class `Ref[T](x: T)`, the block ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} { class C extends B {$\ldots$} ; new Ref(new C) } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ has the type `Ref[_1] forSome { type _1 <: B }`. The block ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} { class C extends B {$\ldots$} ; new C } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ simply has type `B`, because with the rules [here](#simplification-rules) the existentially quantified type `_1 forSome { type _1 <: B }` can be simplified to `B`. Prefix, Infix, and Postfix Operations ------------------------------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} PostfixExpr ::= InfixExpr [id [nl]] InfixExpr ::= PrefixExpr | InfixExpr id [nl] InfixExpr PrefixExpr ::= [`-' | `+' | `!' | `~'] SimpleExpr ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Expressions can be constructed from operands and operators. ### Prefix Operations A prefix operation $\mathit{op};e$ consists of a prefix operator $\mathit{op}$, which must be one of the identifiers ‘`+`’, ‘`-`’, ‘`!`’ or ‘`~`’. The expression $\mathit{op};e$ is equivalent to the postfix method application `e.unary_$\mathit{op}$`. Prefix operators are different from normal function applications in that their operand expression need not be atomic. For instance, the input sequence `-sin(x)` is read as `-(sin(x))`, whereas the function application `negate sin(x)` would be parsed as the application of the infix operator `sin` to the operands `negate` and `(x)`. ### Postfix Operations A postfix operator can be an arbitrary identifier. The postfix operation $e;\mathit{op}$ is interpreted as $e.\mathit{op}$. ### Infix Operations An infix operator can be an arbitrary identifier. Infix operators have precedence and associativity defined as follows: The _precedence_ of an infix operator is determined by the operator's first character. Characters are listed below in increasing order of precedence, with characters on the same line having the same precedence. ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ $\mbox{\rm\sl(all letters)}$ | ^ & < > = ! : + - * / % $\mbox{\rm\sl(all other special characters)}$ ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ That is, operators starting with a letter have lowest precedence, followed by operators starting with ``|`', etc. There's one exception to this rule, which concerns [_assignment operators_](#assignment-operators). The precedence of an assigment operator is the same as the one of simple assignment `(=)`. That is, it is lower than the precedence of any other operator. The _associativity_ of an operator is determined by the operator's last character. Operators ending in a colon ``:`' are right-associative. All other operators are left-associative. Precedence and associativity of operators determine the grouping of parts of an expression as follows. - If there are several infix operations in an expression, then operators with higher precedence bind more closely than operators with lower precedence. - If there are consecutive infix operations $e_0; \mathit{op}_1; e_1; \mathit{op}_2 \ldots \mathit{op}_n; e_n$ with operators $\mathit{op}_1 , \ldots , \mathit{op}_n$ of the same precedence, then all these operators must have the same associativity. If all operators are left-associative, the sequence is interpreted as $(\ldots(e_0;\mathit{op}_1;e_1);\mathit{op}_2\ldots);\mathit{op}_n;e_n$. Otherwise, if all operators are right-associative, the sequence is interpreted as $e_0;\mathit{op}_1;(e_1;\mathit{op}_2;(\ldots \mathit{op}_n;e_n)\ldots)$. - Postfix operators always have lower precedence than infix operators. E.g.\ $e_1;\mathit{op}_1;e_2;\mathit{op}_2$ is always equivalent to $(e_1;\mathit{op}_1;e_2);\mathit{op}_2$. The right-hand operand of a left-associative operator may consist of several arguments enclosed in parentheses, e.g. $e;\mathit{op};(e_1,\ldots,e_n)$. This expression is then interpreted as $e.\mathit{op}(e_1,\ldots,e_n)$. A left-associative binary operation $e_1;\mathit{op};e_2$ is interpreted as $e_1.\mathit{op}(e_2)$. If $\mathit{op}$ is right-associative, the same operation is interpreted as `{ val $x$=$e_1$; $e_2$.$\mathit{op}$($x\,$) }`, where $x$ is a fresh name. ### Assignment Operators An assignment operator is an operator symbol (syntax category `op` in [Identifiers](#identifiers)) that ends in an equals character “`=`”, with the exception of operators for which one of the following conditions holds: #. the operator also starts with an equals character, or #. the operator is one of `(<=)`, `(>=)`, `(!=)`. Assignment operators are treated specially in that they can be expanded to assignments if no other interpretation is valid. Let's consider an assignment operator such as `+=` in an infix operation `$l$ += $r$`, where $l$, $r$ are expressions. This operation can be re-interpreted as an operation which corresponds to the assignment ~~~~~~~~~~~~~~~~ {.scala} $l$ = $l$ + $r$ ~~~~~~~~~~~~~~~~ except that the operation's left-hand-side $l$ is evaluated only once. The re-interpretation occurs if the following two conditions are fulfilled. #. The left-hand-side $l$ does not have a member named `+=`, and also cannot be converted by an [implicit conversion](#implicit-conversions) to a value with a member named `+=`. #. The assignment `$l$ = $l$ + $r$` is type-correct. In particular this implies that $l$ refers to a variable or object that can be assigned to, and that is convertible to a value with a member named `+`. Typed Expressions ----------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} Expr1 ::= PostfixExpr `:' CompoundType ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The typed expression $e: T$ has type $T$. The type of expression $e$ is expected to conform to $T$. The result of the expression is the value of $e$ converted to type $T$. (@) Here are examples of well-typed and illegally typed expressions. ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} 1: Int // legal, of type Int 1: Long // legal, of type Long // 1: string // ***** illegal ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Annotated Expressions --------------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} Expr1 ::= PostfixExpr `:' Annotation {Annotation} ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ An annotated expression `$e$: @$a_1$ $\ldots$ @$a_n$` attaches [annotations](#user-defined-annotations) $a_1 , \ldots , a_n$ to the expression $e$. Assignments ----------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} Expr1 ::= [SimpleExpr `.'] id `=' Expr | SimpleExpr1 ArgumentExprs `=' Expr ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The interpretation of an assignment to a simple variable `$x$ = $e$` depends on the definition of $x$. If $x$ denotes a mutable variable, then the assignment changes the current value of $x$ to be the result of evaluating the expression $e$. The type of $e$ is expected to conform to the type of $x$. If $x$ is a parameterless function defined in some template, and the same template contains a setter function `$x$_=` as member, then the assignment `$x$ = $e$` is interpreted as the invocation `$x$_=($e\,$)` of that setter function. Analogously, an assignment `$f.x$ = $e$` to a parameterless function $x$ is interpreted as the invocation `$f.x$_=($e\,$)`. An assignment `$f$($\mathit{args}\,$) = $e$` with a function application to the left of the ‘`=`’ operator is interpreted as `$f.$update($\mathit{args}$, $e\,$)`, i.e.\ the invocation of an `update` function defined by $f$. (@) Here are some assignment expressions and their equivalent expansions. -------------------------- --------------------- `x.f = e`{.scala} x.f_=(e) `x.f() = e`{.scala} x.f.update(e) `x.f(i) = e`{.scala} x.f.update(i, e) `x.f(i, j) = e`{.scala} x.f.update(i, j, e) -------------------------- --------------------- (@imp-mat-mul) Here is the usual imperative code for matrix multiplication. ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} def matmul(xss: Array[Array[Double]], yss: Array[Array[Double]]) = { val zss: Array[Array[Double]] = new Array(xss.length, yss(0).length) var i = 0 while (i < xss.length) { var j = 0 while (j < yss(0).length) { var acc = 0.0 var k = 0 while (k < yss.length) { acc = acc + xss(i)(k) * yss(k)(j) k += 1 } zss(i)(j) = acc j += 1 } i += 1 } zss } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Desugaring the array accesses and assignments yields the following expanded version: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} def matmul(xss: Array[Array[Double]], yss: Array[Array[Double]]) = { val zss: Array[Array[Double]] = new Array(xss.length, yss.apply(0).length) var i = 0 while (i < xss.length) { var j = 0 while (j < yss.apply(0).length) { var acc = 0.0 var k = 0 while (k < yss.length) { acc = acc + xss.apply(i).apply(k) * yss.apply(k).apply(j) k += 1 } zss.apply(i).update(j, acc) j += 1 } i += 1 } zss } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Conditional Expressions ----------------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} Expr1 ::= `if' `(' Expr `)' {nl} Expr [[semi] `else' Expr] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The conditional expression `if ($e_1$) $e_2$ else $e_3$` chooses one of the values of $e_2$ and $e_3$, depending on the value of $e_1$. The condition $e_1$ is expected to conform to type `Boolean`. The then-part $e_2$ and the else-part $e_3$ are both expected to conform to the expected type of the conditional expression. The type of the conditional expression is the [weak least upper bound](#weak-conformance) of the types of $e_2$ and $e_3$. A semicolon preceding the `else` symbol of a conditional expression is ignored. The conditional expression is evaluated by evaluating first $e_1$. If this evaluates to `true`, the result of evaluating $e_2$ is returned, otherwise the result of evaluating $e_3$ is returned. A short form of the conditional expression eliminates the else-part. The conditional expression `if ($e_1$) $e_2$` is evaluated as if it was `if ($e_1$) $e_2$ else ()`. While Loop Expressions ---------------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} Expr1 ::= `while' `(' Expr ')' {nl} Expr ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The while loop expression `while ($e_1$) $e_2$` is typed and evaluated as if it was an application of `whileLoop ($e_1$) ($e_2$)` where the hypothetical function `whileLoop` is defined as follows. ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} def whileLoop(cond: => Boolean)(body: => Unit): Unit = if (cond) { body ; whileLoop(cond)(body) } else {} ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Do Loop Expressions ------------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} Expr1 ::= `do' Expr [semi] `while' `(' Expr ')' ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The do loop expression `do $e_1$ while ($e_2$)` is typed and evaluated as if it was the expression `($e_1$ ; while ($e_2$) $e_1$)`. A semicolon preceding the `while` symbol of a do loop expression is ignored. For Comprehensions and For Loops -------------------------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} Expr1 ::= `for' (`(' Enumerators `)' | `{' Enumerators `}') {nl} [`yield'] Expr Enumerators ::= Generator {semi Enumerator} Enumerator ::= Generator | Guard | `val' Pattern1 `=' Expr Generator ::= Pattern1 `<-' Expr [Guard] Guard ::= `if' PostfixExpr ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ A for loop `for ($\mathit{enums}\,$) $e$` executes expression $e$ for each binding generated by the enumerators $\mathit{enums}$. A for comprehension `for ($\mathit{enums}\,$) yield $e$` evaluates expression $e$ for each binding generated by the enumerators $\mathit{enums}$ and collects the results. An enumerator sequence always starts with a generator; this can be followed by further generators, value definitions, or guards. A _generator_ `$p$ <- $e$` produces bindings from an expression $e$ which is matched in some way against pattern $p$. A _value definition_ `$p$ = $e$` binds the value name $p$ (or several names in a pattern $p$) to the result of evaluating the expression $e$. A _guard_ `if $e$` contains a boolean expression which restricts enumerated bindings. The precise meaning of generators and guards is defined by translation to invocations of four methods: `map`, `withFilter`, `flatMap`, and `foreach`. These methods can be implemented in different ways for different carrier types. The translation scheme is as follows. In a first step, every generator `$p$ <- $e$`, where $p$ is not [irrefutable](#patterns) for the type of $e$ is replaced by ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} $p$ <- $e$.withFilter { case $p$ => true; case _ => false } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Then, the following rules are applied repeatedly until all comprehensions have been eliminated. - A for comprehension `for ($p$ <- $e\,$) yield $e'$` is translated to `$e$.map { case $p$ => $e'$ }`. - A for loop `for ($p$ <- $e\,$) $e'$` is translated to `$e$.foreach { case $p$ => $e'$ }`. - A for comprehension ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} for ($p$ <- $e$; $p'$ <- $e'; \ldots$) yield $e''$ ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ where `$\ldots$` is a (possibly empty) sequence of generators, definitions, or guards, is translated to ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} $e$.flatMap { case $p$ => for ($p'$ <- $e'; \ldots$) yield $e''$ } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ - A for loop ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} for ($p$ <- $e$; $p'$ <- $e'; \ldots$) $e''$ ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ where `$\ldots$` is a (possibly empty) sequence of generators, definitions, or guards, is translated to ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} $e$.foreach { case $p$ => for ($p'$ <- $e'; \ldots$) $e''$ } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ - A generator `$p$ <- $e$` followed by a guard `if $g$` is translated to a single generator `$p$ <- $e$.withFilter(($x_1 , \ldots , x_n$) => $g\,$)` where $x_1 , \ldots , x_n$ are the free variables of $p$. - A generator `$p$ <- $e$` followed by a value definition `$p'$ = $e'$` is translated to the following generator of pairs of values, where $x$ and $x'$ are fresh names: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} ($p$, $p'$) <- for ($x @ p$ <- $e$) yield { val $x' @ p'$ = $e'$; ($x$, $x'$) } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ (@) The following code produces all pairs of numbers between $1$ and $n-1$ whose sums are prime. ~~~~~~~~~~~~~~~~~~~~~~~ {.scala} for { i <- 1 until n j <- 1 until i if isPrime(i+j) } yield (i, j) ~~~~~~~~~~~~~~~~~~~~~~~ The for comprehension is translated to: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} (1 until n) .flatMap { case i => (1 until i) .withFilter { j => isPrime(i+j) } .map { case j => (i, j) } } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ (@) For comprehensions can be used to express vector and matrix algorithms concisely. For instance, here is a function to compute the transpose of a given matrix: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} def transpose[A](xss: Array[Array[A]]) = { for (i <- Array.range(0, xss(0).length)) yield for (xs <- xss) yield xs(i) } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Here is a function to compute the scalar product of two vectors: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} def scalprod(xs: Array[Double], ys: Array[Double]) = { var acc = 0.0 for ((x, y) <- xs zip ys) acc = acc + x * y acc } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Finally, here is a function to compute the product of two matrices. Compare with the imperative version of \ref{ex:imp-mat-mul}. ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} def matmul(xss: Array[Array[Double]], yss: Array[Array[Double]]) = { val ysst = transpose(yss) for (xs <- xss) yield for (yst <- ysst) yield scalprod(xs, yst) } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The code above makes use of the fact that `map`, `flatMap`, `withFilter`, and `foreach` are defined for instances of class `scala.Array`. Return Expressions ------------------ ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} Expr1 ::= `return' [Expr] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ A return expression `return $e$` must occur inside the body of some enclosing named method or function. The innermost enclosing named method or function in a source program, $f$, must have an explicitly declared result type, and the type of $e$ must conform to it. The return expression evaluates the expression $e$ and returns its value as the result of $f$. The evaluation of any statements or expressions following the return expression is omitted. The type of a return expression is `scala.Nothing`. The expression $e$ may be omitted. The return expression `return` is type-checked and evaluated as if it was `return ()`. An `apply` method which is generated by the compiler as an expansion of an anonymous function does not count as a named function in the source program, and therefore is never the target of a return expression. Returning from a nested anonymous function is implemented by throwing and catching a `scala.runtime.NonLocalReturnException`. Any exception catches between the point of return and the enclosing methods might see the exception. A key comparison makes sure that these exceptions are only caught by the method instance which is terminated by the return. If the return expression is itself part of an anonymous function, it is possible that the enclosing instance of $f$ has already returned before the return expression is executed. In that case, the thrown `scala.runtime.NonLocalReturnException` will not be caught, and will propagate up the call stack. Throw Expressions ----------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} Expr1 ::= `throw' Expr ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ A throw expression `throw $e$` evaluates the expression $e$. The type of this expression must conform to `Throwable`. If $e$ evaluates to an exception reference, evaluation is aborted with the thrown exception. If $e$ evaluates to `null`, evaluation is instead aborted with a `NullPointerException`. If there is an active [`try` expression](#try-expressions) which handles the thrown exception, evaluation resumes with the handler; otherwise the thread executing the `throw` is aborted. The type of a throw expression is `scala.Nothing`. Try Expressions --------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} Expr1 ::= `try' `{' Block `}' [`catch' `{' CaseClauses `}'] [`finally' Expr] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ A try expression is of the form `try { $b$ } catch $h$` where the handler $h$ is a [pattern matching anonymous function](#pattern-matching-anonymous-functions) ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} { case $p_1$ => $b_1$ $\ldots$ case $p_n$ => $b_n$ } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ This expression is evaluated by evaluating the block $b$. If evaluation of $b$ does not cause an exception to be thrown, the result of $b$ is returned. Otherwise the handler $h$ is applied to the thrown exception. If the handler contains a case matching the thrown exception, the first such case is invoked. If the handler contains no case matching the thrown exception, the exception is re-thrown. Let $\mathit{pt}$ be the expected type of the try expression. The block $b$ is expected to conform to $\mathit{pt}$. The handler $h$ is expected conform to type `scala.PartialFunction[scala.Throwable, $\mathit{pt}\,$]`. The type of the try expression is the [weak least upper bound](#weak-conformance) of the type of $b$ and the result type of $h$. A try expression `try { $b$ } finally $e$` evaluates the block $b$. If evaluation of $b$ does not cause an exception to be thrown, the expression $e$ is evaluated. If an exception is thrown during evaluation of $e$, the evaluation of the try expression is aborted with the thrown exception. If no exception is thrown during evaluation of $e$, the result of $b$ is returned as the result of the try expression. If an exception is thrown during evaluation of $b$, the finally block $e$ is also evaluated. If another exception $e$ is thrown during evaluation of $e$, evaluation of the try expression is aborted with the thrown exception. If no exception is thrown during evaluation of $e$, the original exception thrown in $b$ is re-thrown once evaluation of $e$ has completed. The block $b$ is expected to conform to the expected type of the try expression. The finally expression $e$ is expected to conform to type `Unit`. A try expression `try { $b$ } catch $e_1$ finally $e_2$` is a shorthand for `try { try { $b$ } catch $e_1$ } finally $e_2$`. Anonymous Functions ------------------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} Expr ::= (Bindings | [`implicit'] id | `_') `=>' Expr ResultExpr ::= (Bindings | ([`implicit'] id | `_') `:' CompoundType) `=>' Block Bindings ::= `(' Binding {`,' Binding} `)' Binding ::= (id | `_') [`:' Type] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The anonymous function `($x_1$: $T_1 , \ldots , x_n$: $T_n$) => e` maps parameters $x_i$ of types $T_i$ to a result given by expression $e$. The scope of each formal parameter $x_i$ is $e$. Formal parameters must have pairwise distinct names. If the expected type of the anonymous function is of the form `scala.Function$n$[$S_1 , \ldots , S_n$, $R\,$]`, the expected type of $e$ is $R$ and the type $T_i$ of any of the parameters $x_i$ can be omitted, in which case`$T_i$ = $S_i$` is assumed. If the expected type of the anonymous function is some other type, all formal parameter types must be explicitly given, and the expected type of $e$ is undefined. The type of the anonymous function is`scala.Function$n$[$S_1 , \ldots , S_n$, $T\,$]`, where $T$ is the [packed type](#expression-typing) of $e$. $T$ must be equivalent to a type which does not refer to any of the formal parameters $x_i$. The anonymous function is evaluated as the instance creation expression ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} new scala.Function$n$[$T_1 , \ldots , T_n$, $T$] { def apply($x_1$: $T_1 , \ldots , x_n$: $T_n$): $T$ = $e$ } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ In the case of a single untyped formal parameter, `($x\,$) => $e$` can be abbreviated to `$x$ => $e$`. If an anonymous function `($x$: $T\,$) => $e$` with a single typed parameter appears as the result expression of a block, it can be abbreviated to `$x$: $T$ => e`. A formal parameter may also be a wildcard represented by an underscore `_`. In that case, a fresh name for the parameter is chosen arbitrarily. A named parameter of an anonymous function may be optionally preceded by an `implicit` modifier. In that case the parameter is labeled [`implicit`](#implicit-parameters-and-views); however the parameter section itself does not count as an implicit parameter section in the sense defined [here](#implicit-parameters). Hence, arguments to anonymous functions always have to be given explicitly. (@) Examples of anonymous functions: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} x => x // The identity function f => g => x => f(g(x)) // Curried function composition (x: Int,y: Int) => x + y // A summation function () => { count += 1; count } // The function which takes an // empty parameter list $()$, // increments a non-local variable // `count' and returns the new value. _ => 5 // The function that ignores its argument // and always returns 5. ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ ### Placeholder Syntax for Anonymous Functions ~~~~~~~~~~~~~~~~~~~~~~ {.grammar} SimpleExpr1 ::= `_' ~~~~~~~~~~~~~~~~~~~~~~ An expression (of syntactic category `Expr`) may contain embedded underscore symbols `_` at places where identifiers are legal. Such an expression represents an anonymous function where subsequent occurrences of underscores denote successive parameters. Define an _underscore section_ to be an expression of the form `_:$T$` where $T$ is a type, or else of the form `_`, provided the underscore does not appear as the expression part of a type ascription `_:$T$`. An expression $e$ of syntactic category `Expr` _binds_ an underscore section $u$, if the following two conditions hold: (1) $e$ properly contains $u$, and (2) there is no other expression of syntactic category `Expr` which is properly contained in $e$ and which itself properly contains $u$. If an expression $e$ binds underscore sections $u_1 , \ldots , u_n$, in this order, it is equivalent to the anonymous function `($u'_1$, ... $u'_n$) => $e'$` where each $u_i'$ results from $u_i$ by replacing the underscore with a fresh identifier and $e'$ results from $e$ by replacing each underscore section $u_i$ by $u_i'$. (@) The anonymous functions in the left column use placeholder syntax. Each of these is equivalent to the anonymous function on its right. --------------------------- ------------------------------------ `_ + 1`{.scala} `x => x + 1`{.scala} `_ * _`{.scala} `(x1, x2) => x1 * x2`{.scala} `(_: Int) * 2`{.scala} `(x: Int) => (x: Int) * 2`{.scala} `if (_) x else y`{.scala} `z => if (z) x else y`{.scala} `_.map(f)`{.scala} `x => x.map(f)`{.scala} `_.map(_ + 1)`{.scala} `x => x.map(y => y + 1)`{.scala} --------------------------- ------------------------------------ Constant Expressions -------------------- Constant expressions are expressions that the Scala compiler can evaluate to a constant. The definition of ``constant expression'' depends on the platform, but they include at least the expressions of the following forms: - A literal of a value class, such as an integer - A string literal - A class constructed with [`Predef.classOf`](#the-predef-object) - An element of an enumeration from the underlying platform - A literal array, of the form `Array$(c_1 , \ldots , c_n)$`, where all of the $c_i$'s are themselves constant expressions - An identifier defined by a [constant value definition](#value-declarations-and-definitions). Statements ---------- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.grammar} BlockStat ::= Import | {Annotation} [`implicit'] Def | {Annotation} {LocalModifier} TmplDef | Expr1 | TemplateStat ::= Import | {Annotation} {Modifier} Def | {Annotation} {Modifier} Dcl | Expr | ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Statements occur as parts of blocks and templates. A statement can be an import, a definition or an expression, or it can be empty. Statements used in the template of a class definition can also be declarations. An expression that is used as a statement can have an arbitrary value type. An expression statement $e$ is evaluated by evaluating $e$ and discarding the result of the evaluation. Block statements may be definitions which bind local names in the block. The only modifier allowed in all block-local definitions is `implicit`. When prefixing a class or object definition, modifiers `abstract`, `final`, and `sealed` are also permitted. Evaluation of a statement sequence entails evaluation of the statements in the order they are written. Implicit Conversions -------------------- Implicit conversions can be applied to expressions whose type does not match their expected type, to qualifiers in selections, and to unapplied methods. The available implicit conversions are given in the next two sub-sections. We say, a type $T$ is _compatible_ to a type $U$ if $T$ conforms to $U$ after applying [eta-expansion](#eta-expansion) and [view applications](#views). ### Value Conversions The following five implicit conversions can be applied to an expression $e$ which has some value type $T$ and which is type-checked with some expected type $\mathit{pt}$. _Overloading Resolution_ \ If an expression denotes several possible members of a class, [overloading resolution](#overloading-resolution) is applied to pick a unique member. _Type Instantiation_ \ An expression $e$ of polymorphic type ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} [$a_1$ >: $L_1$ <: $U_1 , \ldots , a_n$ >: $L_n$ <: $U_n$]$T$ ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ which does not appear as the function part of a type application is converted to a type instance of $T$ by determining with [local type inference](#local-type-inference) instance types `$T_1 , \ldots , T_n$` for the type variables `$a_1 , \ldots , a_n$` and implicitly embedding $e$ in the [type application](#type-applications) `$e$[$T_1 , \ldots , T_n$]`. _Numeric Widening_ \ If $e$ has a primitive number type which [weakly conforms](#weak-conformance) to the expected type, it is widened to the expected type using one of the numeric conversion methods `toShort`, `toChar`, `toInt`, `toLong`, `toFloat`, `toDouble` defined [here](#numeric-value-types). _Numeric Literal Narrowing_ \ If the expected type is `Byte`, `Short` or `Char`, and the expression $e$ is an integer literal fitting in the range of that type, it is converted to the same literal in that type. _Value Discarding_ \ If $e$ has some value type and the expected type is `Unit`, $e$ is converted to the expected type by embedding it in the term `{ $e$; () }`. _View Application_ \ If none of the previous conversions applies, and $e$'s type does not conform to the expected type $\mathit{pt}$, it is attempted to convert $e$ to the expected type with a [view](#views).\bigskip _Dynamic Member Selection_ \ If none of the previous conversions applies, and $e$ is a prefix of a selection $e.x$, and $e$'s type conforms to class `scala.Dynamic`, then the selection is rewritten according to the rules for [dynamic member selection](#dynamic-member-selection). ### Method Conversions The following four implicit conversions can be applied to methods which are not applied to some argument list. _Evaluation_ \ A parameterless method $m$ of type `=> $T$` is always converted to type $T$ by evaluating the expression to which $m$ is bound. _Implicit Application_ \ If the method takes only implicit parameters, implicit arguments are passed following the rules [here](#implicit-parameters). _Eta Expansion_ \ Otherwise, if the method is not a constructor, and the expected type $\mathit{pt}$ is a function type $(\mathit{Ts}') \Rightarrow T'$, [eta-expansion](#eta-expansion) is performed on the expression $e$. _Empty Application_ \ Otherwise, if $e$ has method type $()T$, it is implicitly applied to the empty argument list, yielding $e()$. ### Overloading Resolution If an identifier or selection $e$ references several members of a class, the context of the reference is used to identify a unique member. The way this is done depends on whether or not $e$ is used as a function. Let $\mathscr{A}$ be the set of members referenced by $e$. Assume first that $e$ appears as a function in an application, as in `$e$($e_1 , \ldots , e_m$)`. One first determines the set of functions that is potentially applicable based on the _shape_ of the arguments. The shape of an argument expression $e$, written $\mathit{shape}(e)$, is a type that is defined as follows: - For a function expression `($p_1$: $T_1 , \ldots , p_n$: $T_n$) => $b$`: `(Any $, \ldots ,$ Any) => $\mathit{shape}(b)$`, where `Any` occurs $n$ times in the argument type. - For a named argument `$n$ = $e$`: $\mathit{shape}(e)$. - For all other expressions: `Nothing`. Let $\mathscr{B}$ be the set of alternatives in $\mathscr{A}$ that are [_applicable_](#function-applications) to expressions $(e_1 , \ldots , e_n)$ of types $(\mathit{shape}(e_1) , \ldots , \mathit{shape}(e_n))$. If there is precisely one alternative in $\mathscr{B}$, that alternative is chosen. Otherwise, let $S_1 , \ldots , S_m$ be the vector of types obtained by typing each argument with an undefined expected type. For every member $m$ in $\mathscr{B}$ one determines whether it is applicable to expressions ($e_1 , \ldots , e_m$) of types $S_1 , \ldots , S_m$. It is an error if none of the members in $\mathscr{B}$ is applicable. If there is one single applicable alternative, that alternative is chosen. Otherwise, let $\mathscr{CC}$ be the set of applicable alternatives which don't employ any default argument in the application to $e_1 , \ldots , e_m$. It is again an error if $\mathscr{CC}$ is empty. Otherwise, one chooses the _most specific_ alternative among the alternatives in $\mathscr{CC}$, according to the following definition of being ``as specific as'', and ``more specific than'': - A parameterized method $m$ of type `($p_1:T_1, \ldots , p_n:T_n$)$U$` is _as specific as_ some other member $m'$ of type $S$ if $m'$ is applicable to arguments `($p_1 , \ldots , p_n\,$)` of types $T_1 , \ldots , T_n$. - A polymorphic method of type `[$a_1$ >: $L_1$ <: $U_1 , \ldots , a_n$ >: $L_n$ <: $U_n$]$T$` is as specific as some other member of type $S$ if $T$ is as specific as $S$ under the assumption that for $i = 1 , \ldots , n$ each $a_i$ is an abstract type name bounded from below by $L_i$ and from above by $U_i$. - A member of any other type is always as specific as a parameterized method or a polymorphic method. - Given two members of types $T$ and $U$ which are neither parameterized nor polymorphic method types, the member of type $T$ is as specific as the member of type $U$ if the existential dual of $T$ conforms to the existential dual of $U$. Here, the existential dual of a polymorphic type `[$a_1$ >: $L_1$ <: $U_1 , \ldots , a_n$ >: $L_n$ <: $U_n$]$T$` is `$T$ forSome { type $a_1$ >: $L_1$ <: $U_1$ $, \ldots ,$ type $a_n$ >: $L_n$ <: $U_n$}`. The existential dual of every other type is the type itself. The _relative weight_ of an alternative $A$ over an alternative $B$ is a number from 0 to 2, defined as the sum of - 1 if $A$ is as specific as $B$, 0 otherwise, and - 1 if $A$ is defined in a class or object which is derived from the class or object defining $B$, 0 otherwise. A class or object $C$ is _derived_ from a class or object $D$ if one of the following holds: - $C$ is a subclass of $D$, or - $C$ is a companion object of a class derived from $D$, or - $D$ is a companion object of a class from which $C$ is derived. An alternative $A$ is _more specific_ than an alternative $B$ if the relative weight of $A$ over $B$ is greater than the relative weight of $B$ over $A$. It is an error if there is no alternative in $\mathscr{CC}$ which is more specific than all other alternatives in $\mathscr{CC}$. Assume next that $e$ appears as a function in a type application, as in `$e$[$\mathit{targs}\,$]`. Then all alternatives in $\mathscr{A}$ which take the same number of type parameters as there are type arguments in $\mathit{targs}$ are chosen. It is an error if no such alternative exists. If there are several such alternatives, overloading resolution is applied again to the whole expression `$e$[$\mathit{targs}\,$]`. Assume finally that $e$ does not appear as a function in either an application or a type application. If an expected type is given, let $\mathscr{B}$ be the set of those alternatives in $\mathscr{A}$ which are [compatible](#implicit-conversions) to it. Otherwise, let $\mathscr{B}$ be the same as $\mathscr{A}$. We choose in this case the most specific alternative among all alternatives in $\mathscr{B}$. It is an error if there is no alternative in $\mathscr{B}$ which is more specific than all other alternatives in $\mathscr{B}$. (@) Consider the following definitions: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} class A extends B {} def f(x: B, y: B) = $\ldots$ def f(x: A, y: B) = $\ldots$ val a: A val b: B ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Then the application `f(b, b)` refers to the first definition of $f$ whereas the application `f(a, a)` refers to the second. Assume now we add a third overloaded definition ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} def f(x: B, y: A) = $\ldots$ ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Then the application `f(a, a)` is rejected for being ambiguous, since no most specific applicable signature exists. ### Local Type Inference Local type inference infers type arguments to be passed to expressions of polymorphic type. Say $e$ is of type [$a_1$ >: $L_1$ <: $U_1 , \ldots , a_n$ >: $L_n$ <: $U_n$]$T$ and no explicit type parameters are given. Local type inference converts this expression to a type application `$e$[$T_1 , \ldots , T_n$]`. The choice of the type arguments $T_1 , \ldots , T_n$ depends on the context in which the expression appears and on the expected type $\mathit{pt}$. There are three cases. _Case 1: Selections_ \ If the expression appears as the prefix of a selection with a name $x$, then type inference is _deferred_ to the whole expression $e.x$. That is, if $e.x$ has type $S$, it is now treated as having type [$a_1$ >: $L_1$ <: $U_1 , \ldots , a_n$ >: $L_n$ <: $U_n$]$S$, and local type inference is applied in turn to infer type arguments for $a_1 , \ldots , a_n$, using the context in which $e.x$ appears. _Case 2: Values_ \ If the expression $e$ appears as a value without being applied to value arguments, the type arguments are inferred by solving a constraint system which relates the expression's type $T$ with the expected type $\mathit{pt}$. Without loss of generality we can assume that $T$ is a value type; if it is a method type we apply [eta-expansion](#eta-expansion) to convert it to a function type. Solving means finding a substitution $\sigma$ of types $T_i$ for the type parameters $a_i$ such that - None of inferred types $T_i$ is a [singleton type](#singleton-types) - All type parameter bounds are respected, i.e.\ $\sigma L_i <: \sigma a_i$ and $\sigma a_i <: \sigma U_i$ for $i = 1 , \ldots , n$. - The expression's type conforms to the expected type, i.e.\ $\sigma T <: \sigma \mathit{pt}$. It is a compile time error if no such substitution exists. If several substitutions exist, local-type inference will choose for each type variable $a_i$ a minimal or maximal type $T_i$ of the solution space. A _maximal_ type $T_i$ will be chosen if the type parameter $a_i$ appears [contravariantly](#variance-annotations) in the type $T$ of the expression. A _minimal_ type $T_i$ will be chosen in all other situations, i.e.\ if the variable appears covariantly, non-variantly or not at all in the type $T$. We call such a substitution an _optimal solution_ of the given constraint system for the type $T$. _Case 3: Methods_ \ The last case applies if the expression $e$ appears in an application $e(d_1 , \ldots , d_m)$. In that case $T$ is a method type $(p_1:R_1 , \ldots , p_m:R_m)T'$. Without loss of generality we can assume that the result type $T'$ is a value type; if it is a method type we apply [eta-expansion](#eta-expansion) to convert it to a function type. One computes first the types $S_j$ of the argument expressions $d_j$, using two alternative schemes. Each argument expression $d_j$ is typed first with the expected type $R_j$, in which the type parameters $a_1 , \ldots , a_n$ are taken as type constants. If this fails, the argument $d_j$ is typed instead with an expected type $R_j'$ which results from $R_j$ by replacing every type parameter in $a_1 , \ldots , a_n$ with {\sl undefined}. In a second step, type arguments are inferred by solving a constraint system which relates the method's type with the expected type $\mathit{pt}$ and the argument types $S_1 , \ldots , S_m$. Solving the constraint system means finding a substitution $\sigma$ of types $T_i$ for the type parameters $a_i$ such that - None of inferred types $T_i$ is a [singleton type](#singleton-types) - All type parameter bounds are respected, i.e.\ $\sigma L_i <: \sigma a_i$ and $\sigma a_i <: \sigma U_i$ for $i = 1 , \ldots , n$. - The method's result type $T'$ conforms to the expected type, i.e.\ $\sigma T' <: \sigma \mathit{pt}$. - Each argument type [weakly conforms](#weak-conformance) to the corresponding formal parameter type, i.e.\ $\sigma S_j <:_w \sigma R_j$ for $j = 1 , \ldots , m$. It is a compile time error if no such substitution exists. If several solutions exist, an optimal one for the type $T'$ is chosen. All or parts of an expected type $\mathit{pt}$ may be undefined. The rules for [conformance](#conformance) are extended to this case by adding the rule that for any type $T$ the following two statements are always true: $\mathit{undefined} <: T$ and $T <: \mathit{undefined}$ It is possible that no minimal or maximal solution for a type variable exists, in which case a compile-time error results. Because $<:$ is a pre-order, it is also possible that a solution set has several optimal solutions for a type. In that case, a Scala compiler is free to pick any one of them. (@) Consider the two methods: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} def cons[A](x: A, xs: List[A]): List[A] = x :: xs def nil[B]: List[B] = Nil ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ and the definition ~~~~~~~~~~~~~~~~~~~~~~ {.scala} val xs = cons(1, nil) ~~~~~~~~~~~~~~~~~~~~~~ The application of `cons` is typed with an undefined expected type. This application is completed by local type inference to `cons[Int](1, nil)`. Here, one uses the following reasoning to infer the type argument `Int` for the type parameter `a`: First, the argument expressions are typed. The first argument `1` has type `Int` whereas the second argument `nil` is itself polymorphic. One tries to type-check `nil` with an expected type `List[a]`. This leads to the constraint system ~~~~~~~~~~~~~~~~~~~~ {.scala} List[b?] <: List[a] ~~~~~~~~~~~~~~~~~~~~ where we have labeled `b?` with a question mark to indicate that it is a variable in the constraint system. Because class `List` is covariant, the optimal solution of this constraint is ~~~~~~~~~~~~~~~~~~ {.scala} b = scala.Nothing ~~~~~~~~~~~~~~~~~~ In a second step, one solves the following constraint system for the type parameter `a` of `cons`: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} Int <: a? List[scala.Nothing] <: List[a?] List[a?] <: $\mbox{\sl undefined}$ ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The optimal solution of this constraint system is ~~~~~~~~ {.scala} a = Int ~~~~~~~~ so `Int` is the type inferred for `a`. (@) Consider now the definition ~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} val ys = cons("abc", xs) ~~~~~~~~~~~~~~~~~~~~~~~~~ where `xs` is defined of type `List[Int]` as before. In this case local type inference proceeds as follows. First, the argument expressions are typed. The first argument `"abc"` has type `String`. The second argument `xs` is first tried to be typed with expected type `List[a]`. This fails, as `List[Int]` is not a subtype of `List[a]`. Therefore, the second strategy is tried; `xs` is now typed with expected type `List[$\mbox{\sl undefined}$]`. This succeeds and yields the argument type `List[Int]`. In a second step, one solves the following constraint system for the type parameter `a` of `cons`: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} String <: a? List[Int] <: List[a?] List[a?] <: $\mbox{\sl undefined}$ ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The optimal solution of this constraint system is ~~~~~~~~~~~~~~ {.scala} a = scala.Any ~~~~~~~~~~~~~~ so `scala.Any` is the type inferred for `a`. ### Eta Expansion _Eta-expansion_ converts an expression of method type to an equivalent expression of function type. It proceeds in two steps. First, one identifes the maximal sub-expressions of $e$; let's say these are $e_1 , \ldots , e_m$. For each of these, one creates a fresh name $x_i$. Let $e'$ be the expression resulting from replacing every maximal subexpression $e_i$ in $e$ by the corresponding fresh name $x_i$. Second, one creates a fresh name $y_i$ for every argument type $T_i$ of the method ($i = 1 , \ldots , n$). The result of eta-conversion is then: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} { val $x_1$ = $e_1$; $\ldots$ val $x_m$ = $e_m$; ($y_1: T_1 , \ldots , y_n: T_n$) => $e'$($y_1 , \ldots , y_n$) } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ ### Dynamic Member Selection The standard Scala library defines a trait `scala.Dynamic` which defines a member \@invokeDynamic@ as follows: ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} package scala trait Dynamic { def applyDynamic (name: String, args: Any*): Any ... } ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Assume a selection of the form $e.x$ where the type of $e$ conforms to `scala.Dynamic`. Further assuming the selection is not followed by any function arguments, such an expression can be rewitten under the conditions given [here](#implicit-conversions) to: ~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} $e$.applyDynamic("$x$") ~~~~~~~~~~~~~~~~~~~~~~~~ If the selection is followed by some arguments, e.g.\ $e.x(\mathit{args})$, then that expression is rewritten to ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ {.scala} $e$.applyDynamic("$x$", $\mathit{args}$) ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~