# Basic Declarations and Definitions

A *declaration* introduces names and assigns them types. It can
form part of a class definition or of a
refinement in a compound type.

A *definition* introduces names that denote terms or types. It can
form part of an object or class definition or it can be local to a
block. Both declarations and definitions produce *bindings* that
associate type names with type definitions or bounds, and that
associate term names with types.

The scope of a name introduced by a declaration or definition is the whole statement sequence containing the binding. However, there is a restriction on forward references in blocks: In a statement sequence $s_1 \ldots s_n$ making up a block, if a simple name in $s_i$ refers to an entity defined by $s_j$ where $j \geq i$, then for all $s_k$ between and including $s_i$ and $s_j$,

- $s_k$ cannot be a variable definition.
- If $s_k$ is a value definition, it must be lazy.

## Value Declarations and Definitions

A value declaration `val $x$: $T$`

introduces $x$ as a name of a value of
type $T$.

A value definition `val $x$: $T$ = $e$`

defines $x$ as a
name of the value that results from the evaluation of $e$.
If the value definition is not recursive, the type
$T$ may be omitted, in which case the packed type of
expression $e$ is assumed. If a type $T$ is given, then $e$ is expected to
conform to it.

Evaluation of the value definition implies evaluation of its
right-hand side $e$, unless it has the modifier `lazy`

. The
effect of the value definition is to bind $x$ to the value of $e$
converted to type $T$. A `lazy`

value definition evaluates
its right hand side $e$ the first time the value is accessed.

A *constant value definition* is of the form

where `e`

is a constant expression.
The `final`

modifier must be
present and no type annotation may be given. References to the
constant value `x`

are themselves treated as constant expressions; in the
generated code they are replaced by the definition's right-hand side `e`

.

Value definitions can alternatively have a pattern
as left-hand side. If $p$ is some pattern other
than a simple name or a name followed by a colon and a type, then the
value definition `val $p$ = $e$`

is expanded as follows:

- If the pattern $p$ has bound variables $x_1 , \ldots , x_n$, where $n > 1$:

Here, $\$ x$ is a fresh name.

- If $p$ has a unique bound variable $x$:

- If $p$ has no bound variables:

###### Example

The following are examples of value definitions

The last two definitions have the following expansions.

The name of any declared or defined value may not end in `_=`

.

A value declaration `val $x_1 , \ldots , x_n$: $T$`

is a shorthand for the
sequence of value declarations `val $x_1$: $T$; ...; val $x_n$: $T$`

.
A value definition `val $p_1 , \ldots , p_n$ = $e$`

is a shorthand for the
sequence of value definitions `val $p_1$ = $e$; ...; val $p_n$ = $e$`

.
A value definition `val $p_1 , \ldots , p_n: T$ = $e$`

is a shorthand for the
sequence of value definitions `val $p_1: T$ = $e$; ...; val $p_n: T$ = $e$`

.

## Variable Declarations and Definitions

A variable declaration `var $x$: $T$`

is equivalent to the declarations
of both a *getter function* $x$ *and* a *setter function* `$x$_=`

:

An implementation of a class may *define* a declared variable
using a variable definition, or by defining the corresponding setter and getter methods.

A variable definition `var $x$: $T$ = $e$`

introduces a
mutable variable with type $T$ and initial value as given by the
expression $e$. The type $T$ can be omitted, in which case the type of
$e$ is assumed. If $T$ is given, then $e$ is expected to
conform to it.

Variable definitions can alternatively have a pattern
as left-hand side. A variable definition
`var $p$ = $e$`

where $p$ is a pattern other
than a simple name or a name followed by a colon and a type is expanded in the same way
as a value definition
`val $p$ = $e$`

, except that
the free names in $p$ are introduced as mutable variables, not values.

The name of any declared or defined variable may not end in `_=`

.

A variable definition `var $x$: $T$ = _`

can appear only as a member of a template.
It introduces a mutable field with type $T$ and a default initial value.
The default value depends on the type $T$ as follows:

default | type $T$ |
---|---|

`0` |
`Int` or one of its subrange types |

`0L` |
`Long` |

`0.0f` |
`Float` |

`0.0d` |
`Double` |

`false` |
`Boolean` |

`()` |
`Unit` |

`null` |
all other types |

When they occur as members of a template, both forms of variable
definition also introduce a getter function $x$ which returns the
value currently assigned to the variable, as well as a setter function
`$x$_=`

which changes the value currently assigned to the variable.
The functions have the same signatures as for a variable declaration.
The template then has these getter and setter functions as
members, whereas the original variable cannot be accessed directly as
a template member.

###### Example

The following example shows how *properties* can be
simulated in Scala. It defines a class `TimeOfDayVar`

of time
values with updatable integer fields representing hours, minutes, and
seconds. Its implementation contains tests that allow only legal
values to be assigned to these fields. The user code, on the other
hand, accesses these fields just like normal variables.

A variable declaration `var $x_1 , \ldots , x_n$: $T$`

is a shorthand for the
sequence of variable declarations `var $x_1$: $T$; ...; var $x_n$: $T$`

.
A variable definition `var $x_1 , \ldots , x_n$ = $e$`

is a shorthand for the
sequence of variable definitions `var $x_1$ = $e$; ...; var $x_n$ = $e$`

.
A variable definition `var $x_1 , \ldots , x_n: T$ = $e$`

is a shorthand for
the sequence of variable definitions
`var $x_1: T$ = $e$; ...; var $x_n: T$ = $e$`

.

## Type Declarations and Type Aliases

A *type declaration* `type $t$[$\mathit{tps}\,$] >: $L$ <: $U$`

declares
$t$ to be an abstract type with lower bound type $L$ and upper bound
type $U$. If the type parameter clause `[$\mathit{tps}\,$]`

is omitted, $t$ abstracts over a first-order type, otherwise $t$ stands for a type constructor that accepts type arguments as described by the type parameter clause.

If a type declaration appears as a member declaration of a
type, implementations of the type may implement $t$ with any type $T$
for which $L <: T <: U$. It is a compile-time error if
$L$ does not conform to $U$. Either or both bounds may be omitted.
If the lower bound $L$ is absent, the bottom type
`scala.Nothing`

is assumed. If the upper bound $U$ is absent,
the top type `scala.Any`

is assumed.

A type constructor declaration imposes additional restrictions on the concrete types for which $t$ may stand. Besides the bounds $L$ and $U$, the type parameter clause may impose higher-order bounds and variances, as governed by the conformance of type constructors.

The scope of a type parameter extends over the bounds `>: $L$ <: $U$`

and the type parameter clause $\mathit{tps}$ itself. A
higher-order type parameter clause (of an abstract type constructor
$tc$) has the same kind of scope, restricted to the declaration of the
type parameter $tc$.

To illustrate nested scoping, these declarations are all equivalent: `type t[m[x] <: Bound[x], Bound[x]]`

, `type t[m[x] <: Bound[x], Bound[y]]`

and `type t[m[x] <: Bound[x], Bound[_]]`

, as the scope of, e.g., the type parameter of $m$ is limited to the declaration of $m$. In all of them, $t$ is an abstract type member that abstracts over two type constructors: $m$ stands for a type constructor that takes one type parameter and that must be a subtype of $Bound$, $t$'s second type constructor parameter. `t[MutableList, Iterable]`

is a valid use of $t$.

A *type alias* `type $t$ = $T$`

defines $t$ to be an alias
name for the type $T$. The left hand side of a type alias may
have a type parameter clause, e.g. `type $t$[$\mathit{tps}\,$] = $T$`

. The scope
of a type parameter extends over the right hand side $T$ and the
type parameter clause $\mathit{tps}$ itself.

The scope rules for definitions
and type parameters
make it possible that a type name appears in its
own bound or in its right-hand side. However, it is a static error if
a type alias refers recursively to the defined type constructor itself.
That is, the type $T$ in a type alias `type $t$[$\mathit{tps}\,$] = $T$`

may not
refer directly or indirectly to the name $t$. It is also an error if
an abstract type is directly or indirectly its own upper or lower bound.

###### Example

The following are legal type declarations and definitions:

The following are illegal:

If a type alias `type $t$[$\mathit{tps}\,$] = $S$`

refers to a class type
$S$, the name $t$ can also be used as a constructor for
objects of type $S$.

###### Example

Suppose we make `Pair`

an alias of the parameterized class `Tuple2`

,
as follows:

As a consequence, for any two types $S$ and $T$, the type
`Pair[$S$, $T\,$]`

is equivalent to the type `Tuple2[$S$, $T\,$]`

.
`Pair`

can also be used as a constructor instead of `Tuple2`

, as in:

## Type Parameters

Type parameters appear in type definitions, class definitions, and
function definitions. In this section we consider only type parameter
definitions with lower bounds `>: $L$`

and upper bounds
`<: $U$`

whereas a discussion of context bounds
`: $U$`

and view bounds `<% $U$`

is deferred to here.

The most general form of a first-order type parameter is
`$@a_1 \ldots @a_n$ $\pm$ $t$ >: $L$ <: $U$`

.
Here, $L$, and $U$ are lower and upper bounds that
constrain possible type arguments for the parameter. It is a
compile-time error if $L$ does not conform to $U$. $\pm$ is a *variance*, i.e. an optional prefix of either `+`

, or
`-`

. One or more annotations may precede the type parameter.

The names of all type parameters must be pairwise different in their enclosing type parameter clause. The scope of a type parameter includes in each case the whole type parameter clause. Therefore it is possible that a type parameter appears as part of its own bounds or the bounds of other type parameters in the same clause. However, a type parameter may not be bounded directly or indirectly by itself.

A type constructor parameter adds a nested type parameter clause to the type parameter. The most general form of a type constructor parameter is `$@a_1\ldots@a_n$ $\pm$ $t[\mathit{tps}\,]$ >: $L$ <: $U$`

.

The above scoping restrictions are generalized to the case of nested type parameter clauses, which declare higher-order type parameters. Higher-order type parameters (the type parameters of a type parameter $t$) are only visible in their immediately surrounding parameter clause (possibly including clauses at a deeper nesting level) and in the bounds of $t$. Therefore, their names must only be pairwise different from the names of other visible parameters. Since the names of higher-order type parameters are thus often irrelevant, they may be denoted with a `‘_’`

, which is nowhere visible.

###### Example

Here are some well-formed type parameter clauses:

The following type parameter clauses are illegal:

## Variance Annotations

Variance annotations indicate how instances of parameterized types vary with respect to subtyping. A ‘+’ variance indicates a covariant dependency, a ‘-’ variance indicates a contravariant dependency, and a missing variance indication indicates an invariant dependency.

A variance annotation constrains the way the annotated type variable
may appear in the type or class which binds the type parameter. In a
type definition `type $T$[$\mathit{tps}\,$] = $S$`

, or a type
declaration `type $T$[$\mathit{tps}\,$] >: $L$ <: $U$`

type parameters labeled
‘+’ must only appear in covariant position whereas
type parameters labeled ‘-’ must only appear in contravariant
position. Analogously, for a class definition
`class $C$[$\mathit{tps}\,$]($\mathit{ps}\,$) extends $T$ { $x$: $S$ => ...}`

,
type parameters labeled
‘+’ must only appear in covariant position in the
self type $S$ and the template $T$, whereas type
parameters labeled ‘-’ must only appear in contravariant
position.

The variance position of a type parameter in a type or template is defined as follows. Let the opposite of covariance be contravariance, and the opposite of invariance be itself. The top-level of the type or template is always in covariant position. The variance position changes at the following constructs.

- The variance position of a method parameter is the opposite of the variance position of the enclosing parameter clause.
- The variance position of a type parameter is the opposite of the variance position of the enclosing type parameter clause.
- The variance position of the lower bound of a type declaration or type parameter is the opposite of the variance position of the type declaration or parameter.
- The type of a mutable variable is always in invariant position.
- The right-hand side of a type alias is always in invariant position.
- The prefix $S$ of a type selection
`$S$#$T$`

is always in invariant position. - For a type argument $T$ of a type
`$S$[$\ldots T \ldots$ ]`

: If the corresponding type parameter is invariant, then $T$ is in invariant position. If the corresponding type parameter is contravariant, the variance position of $T$ is the opposite of the variance position of the enclosing type`$S$[$\ldots T \ldots$ ]`

.

References to the type parameters in object-private or object-protected values, types, variables, or methods of the class are not checked for their variance position. In these members the type parameter may appear anywhere without restricting its legal variance annotations.

###### Example

The following variance annotation is legal.

With this variance annotation, type instances of $P$ subtype covariantly with respect to their arguments. For instance,

If the members of $P$ are mutable variables, the same variance annotation becomes illegal.

If the mutable variables are object-private, the class definition becomes legal again:

###### Example

The following variance annotation is illegal, since $a$ appears
in contravariant position in the parameter of `append`

:

The problem can be avoided by generalizing the type of `append`

by means of a lower bound:

###### Example

With that annotation, we have that
`OutputChannel[AnyRef]`

conforms to `OutputChannel[String]`

.
That is, a
channel on which one can write any object can substitute for a channel
on which one can write only strings.

## Function Declarations and Definitions

A *function declaration* has the form `def $f\,\mathit{psig}$: $T$`

, where
$f$ is the function's name, $\mathit{psig}$ is its parameter
signature and $T$ is its result type. A *function definition*
`def $f\,\mathit{psig}$: $T$ = $e$`

also includes a *function body* $e$,
i.e. an expression which defines the function's result. A parameter
signature consists of an optional type parameter clause `[$\mathit{tps}\,$]`

,
followed by zero or more value parameter clauses
`($\mathit{ps}_1$)$\ldots$($\mathit{ps}_n$)`

. Such a declaration or definition
introduces a value with a (possibly polymorphic) method type whose
parameter types and result type are as given.

The type of the function body is expected to conform to the function's declared result type, if one is given. If the function definition is not recursive, the result type may be omitted, in which case it is determined from the packed type of the function body.

A *type parameter clause* $\mathit{tps}$ consists of one or more
type declarations, which introduce type
parameters, possibly with bounds. The scope of a type parameter includes
the whole signature, including any of the type parameter bounds as
well as the function body, if it is present.

A *value parameter clause* $\mathit{ps}$ consists of zero or more formal
parameter bindings such as `$x$: $T$`

or `$x: T = e$`

, which bind value
parameters and associate them with their types.

### Default Arguments

Each value parameter declaration may optionally define a default argument. The default argument expression $e$ is type-checked with an expected type $T'$ obtained by replacing all occurrences of the function's type parameters in $T$ by the undefined type.

For every parameter $p_{i,j}$ with a default argument a method named
`$f\$$default$\$$n`

is generated which computes the default argument
expression. Here, $n$ denotes the parameter's position in the method
declaration. These methods are parametrized by the type parameter clause
`[$\mathit{tps}\,$]`

and all value parameter clauses
`($\mathit{ps}_1$)$\ldots$($\mathit{ps}_{i-1}$)`

preceding $p_{i,j}$.
The `$f\$$default$\$$n`

methods are inaccessible for user programs.

###### Example

In the method

the default expression `0`

is type-checked with an undefined expected
type. When applying `compare()`

, the default value `0`

is inserted
and `T`

is instantiated to `Int`

. The methods computing the default
arguments have the form:

The scope of a formal value parameter name $x$ comprises all subsequent parameter clauses, as well as the method return type and the function body, if they are given. Both type parameter names and value parameter names must be pairwise distinct.

A default value which depends on earlier parameters uses the actual arguments if they are provided, not the default arguments.

If an implicit argument is not found by implicit search, it may be supplied using a default argument.

### By-Name Parameters

The type of a value parameter may be prefixed by `=>`

, e.g.
`$x$: => $T$`

. The type of such a parameter is then the
parameterless method type `=> $T$`

. This indicates that the
corresponding argument is not evaluated at the point of function
application, but instead is evaluated at each use within the
function. That is, the argument is evaluated using *call-by-name*.

The by-name modifier is disallowed for parameters of classes that
carry a `val`

or `var`

prefix, including parameters of case
classes for which a `val`

prefix is implicitly generated.

###### Example

The declaration

indicates that both parameters of `whileLoop`

are evaluated using
call-by-name.

### Repeated Parameters

The last value parameter of a parameter section may be suffixed by
`'*'`

, e.g. `(..., $x$:$T$*)`

. The type of such a
*repeated* parameter inside the method is then the sequence type
`scala.Seq[$T$]`

. Methods with repeated parameters
`$T$*`

take a variable number of arguments of type $T$.
That is, if a method $m$ with type
`($p_1:T_1 , \ldots , p_n:T_n, p_s:S$*)$U$`

is applied to arguments
$(e_1 , \ldots , e_k)$ where $k \geq n$, then $m$ is taken in that application
to have type $(p_1:T_1 , \ldots , p_n:T_n, p_s:S , \ldots , p_{s'}S)U$, with
$k - n$ occurrences of type
$S$ where any parameter names beyond $p_s$ are fresh. The only exception to
this rule is if the last argument is
marked to be a *sequence argument* via a `_*`

type
annotation. If $m$ above is applied to arguments
`($e_1 , \ldots , e_n, e'$: _*)`

, then the type of $m$ in
that application is taken to be
`($p_1:T_1, \ldots , p_n:T_n,p_{s}:$scala.Seq[$S$])`

.

It is not allowed to define any default arguments in a parameter section with a repeated parameter.

###### Example

The following method definition computes the sum of the squares of a variable number of integer arguments.

The following applications of this method yield `0`

, `1`

,
`6`

, in that order.

Furthermore, assume the definition:

The following application of method `sum`

is ill-formed:

By contrast, the following application is well formed and yields again
the result `6`

:

### Procedures

Special syntax exists for procedures, i.e. functions that return the
`Unit`

value `()`

.
A *procedure declaration* is a function declaration where the result type
is omitted. The result type is then implicitly completed to the
`Unit`

type. E.g., `def $f$($\mathit{ps}$)`

is equivalent to
`def $f$($\mathit{ps}$): Unit`

.

A *procedure definition* is a function definition where the result type
and the equals sign are omitted; its defining expression must be a block.
E.g., `def $f$($\mathit{ps}$) {$\mathit{stats}$}`

is equivalent to
`def $f$($\mathit{ps}$): Unit = {$\mathit{stats}$}`

.

###### Example

Here is a declaration and a definition of a procedure named `write`

:

The code above is implicitly completed to the following code:

### Method Return Type Inference

A class member definition $m$ that overrides some other function $m'$ in a base class of $C$ may leave out the return type, even if it is recursive. In this case, the return type $R'$ of the overridden function $m'$, seen as a member of $C$, is taken as the return type of $m$ for each recursive invocation of $m$. That way, a type $R$ for the right-hand side of $m$ can be determined, which is then taken as the return type of $m$. Note that $R$ may be different from $R'$, as long as $R$ conforms to $R'$.

###### Example

Assume the following definitions:

Here, it is OK to leave out the result type of `factorial`

in `C`

, even though the method is recursive.

## Import Clauses

An import clause has the form `import $p$.$I$`

where $p$ is a
stable identifier and $I$ is an import expression.
The import expression determines a set of names of importable members of $p$
which are made available without qualification. A member $m$ of $p$ is
*importable* if it is accessible.
The most general form of an import expression is a list of *import selectors*

for $n \geq 0$, where the final wildcard `‘_’`

may be absent. It
makes available each importable member `$p$.$x_i$`

under the unqualified name
$y_i$. I.e. every import selector `$x_i$ => $y_i$`

renames
`$p$.$x_i$`

to
$y_i$. If a final wildcard is present, all importable members $z$ of
$p$ other than `$x_1 , \ldots , x_n,y_1 , \ldots , y_n$`

are also made available
under their own unqualified names.

Import selectors work in the same way for type and term members. For
instance, an import clause `import $p$.{$x$ => $y\,$}`

renames the term
name `$p$.$x$`

to the term name $y$ and the type name `$p$.$x$`

to the type name $y$. At least one of these two names must
reference an importable member of $p$.

If the target in an import selector is a wildcard, the import selector
hides access to the source member. For instance, the import selector
`$x$ => _`

“renames” $x$ to the wildcard symbol (which is
unaccessible as a name in user programs), and thereby effectively
prevents unqualified access to $x$. This is useful if there is a
final wildcard in the same import selector list, which imports all
members not mentioned in previous import selectors.

The scope of a binding introduced by an import-clause starts immediately after the import clause and extends to the end of the enclosing block, template, package clause, or compilation unit, whichever comes first.

Several shorthands exist. An import selector may be just a simple name
$x$. In this case, $x$ is imported without renaming, so the
import selector is equivalent to `$x$ => $x$`

. Furthermore, it is
possible to replace the whole import selector list by a single
identifier or wildcard. The import clause `import $p$.$x$`

is
equivalent to `import $p$.{$x\,$}`

, i.e. it makes available without
qualification the member $x$ of $p$. The import clause
`import $p$._`

is equivalent to
`import $p$.{_}`

,
i.e. it makes available without qualification all members of $p$
(this is analogous to `import $p$.*`

in Java).

An import clause with multiple import expressions
`import $p_1$.$I_1 , \ldots , p_n$.$I_n$`

is interpreted as a
sequence of import clauses
`import $p_1$.$I_1$; $\ldots$; import $p_n$.$I_n$`

.

###### Example

Consider the object definition:

Then the block

is equivalent to the block