Types
We distinguish between firstorder types and type constructors, which take type parameters and yield types. A subset of firstorder types called value types represents sets of (firstclass) values. Value types are either concrete or abstract.
Every concrete value type can be represented as a class type, i.e. a type designator that refers to a class or a trait ^{1}, or as a compound type representing an intersection of types, possibly with a refinement that further constrains the types of its members.
Abstract value types are introduced by type parameters and abstract type bindings. Parentheses in types can be used for grouping.
Nonvalue types capture properties of identifiers that are not values. For example, a type constructor does not directly specify a type of values. However, when a type constructor is applied to the correct type arguments, it yields a firstorder type, which may be a value type.
Nonvalue types are expressed indirectly in Scala. E.g., a method type is
described by writing down a method signature, which in itself is not a real
type, although it gives rise to a corresponding method type.
Type constructors are another example, as one can write
type Swap[m[_, _], a,b] = m[b, a]
, but there is no syntax to write
the corresponding anonymous type function directly.
Paths
Paths are not types themselves, but they can be a part of named types and in that function form a central role in Scala's type system.
A path is one of the following.
 The empty path ε (which cannot be written explicitly in user programs).
 $C.$
this
, where $C$ references a class. The paththis
is taken as a shorthand for $C.$this
where $C$ is the name of the class directly enclosing the reference.  $p.x$ where $p$ is a path and $x$ is a stable member of $p$. Stable members are packages or members introduced by object definitions or by value definitions of nonvolatile types.
 $C.$
super
$.x$ or $C.$super
$[M].x$ where $C$ references a class and $x$ references a stable member of the super class or designated parent class $M$ of $C$. The prefixsuper
is taken as a shorthand for $C.$super
where $C$ is the name of the class directly enclosing the reference.
A stable identifier is a path which ends in an identifier.
Value Types
Every value in Scala has a type which is of one of the following forms.
Singleton Types
A singleton type is of the form $p.$type
, where $p$ is a
path pointing to a value expected to conform
to scala.AnyRef
. The type denotes the set of values
consisting of null
and the value denoted by $p$.
A stable type is either a singleton type or a type which is
declared to be a subtype of trait scala.Singleton
.
Type Projection
A type projection $T$#$x$ references the type member named $x$ of type $T$.
Type Designators
A type designator refers to a named value type. It can be simple or qualified. All such type designators are shorthands for type projections.
Specifically, the unqualified type name $t$ where $t$ is bound in some
class, object, or package $C$ is taken as a shorthand for
$C.$this.type#
$t$. If $t$ is
not bound in a class, object, or package, then $t$ is taken as a
shorthand for ε.type#
$t$.
A qualified type designator has the form p.t
where p
is
a path and t is a type name. Such a type designator is
equivalent to the type projection p.type#t
.
Example
Some type designators and their expansions are listed below. We assume
a local type parameter $t$, a value maintable
with a type member Node
and the standard class scala.Int
,
Designator  Expansion 

t  ε.type#t 
Int  scala.type#Int 
scala.Int  scala.type#Int 
data.maintable.Node  data.maintable.type#Node 
Parameterized Types
A parameterized type $T[ T_1 , \ldots , T_n ]$ consists of a type designator $T$ and type parameters $T_1 , \ldots , T_n$ where $n \geq 1$. $T$ must refer to a type constructor which takes $n$ type parameters $a_1 , \ldots , a_n$.
Say the type parameters have lower bounds $L_1 , \ldots , L_n$ and upper bounds $U_1, \ldots, U_n$. The parameterized type is wellformed if each actual type parameter conforms to its bounds, i.e. $\sigma L_i <: T_i <: \sigma U_i$ where $\sigma$ is the substitution $[ a_1 := T_1 , \ldots , a_n := T_n ]$.
Example Parameterized Types
Given the partial type definitions:
the following parameterized types are well formed:
Example
Given the above type definitions, the following types are illformed:
Tuple Types
A tuple type $(T_1 , \ldots , T_n)$ is an alias for the
class scala.Tuple$n$[$T_1$, … , $T_n$]
, where $n \geq 2$.
Tuple classes are case classes whose fields can be accessed using
selectors _1
, … , _n
. Their functionality is
abstracted in a corresponding Product
trait. The nary tuple
class and product trait are defined at least as follows in the
standard Scala library (they might also add other methods and
implement other traits).
Annotated Types
An annotated type $T$ $a_1, \ldots, a_n$ attaches annotations $a_1 , \ldots , a_n$ to the type $T$.
Example
The following type adds the @suspendable
annotation to the type String
:
Compound Types
A compound type $T_1$ with
… with
$T_n \{ R \}$
represents objects with members as given in the component types
$T_1 , \ldots , T_n$ and the refinement $\{ R \}$. A refinement
$\{ R \}$ contains declarations and type definitions.
If a declaration or definition overrides a declaration or definition in
one of the component types $T_1 , \ldots , T_n$, the usual rules for
overriding apply; otherwise the declaration
or definition is said to be “structural” ^{2}.
Within a method declaration in a structural refinement, the type of any value parameter may only refer to type parameters or abstract types that are contained inside the refinement. That is, it must refer either to a type parameter of the method itself, or to a type definition within the refinement. This restriction does not apply to the method's result type.
If no refinement is given, the empty refinement is implicitly added,
i.e. $T_1$ with
… with
$T_n$ is a shorthand for $T_1$ with
… with
$T_n \{\}$.
A compound type may also consist of just a refinement
$\{ R \}$ with no preceding component types. Such a type is
equivalent to AnyRef
$\{ R \}$.
Example
The following example shows how to declare and use a method which has a parameter type that contains a refinement with structural declarations.
Although Bird
and Plane
do not share any parent class other than
Object
, the parameter r of method takeoff
is defined using a
refinement with structural declarations to accept any object that declares
a value callsign
and a fly
method.
Infix Types
An infix type $T_1$ op
$T_2$ consists of an infix
operator op
which gets applied to two type operands $T_1$ and
$T_2$. The type is equivalent to the type application
op
$[T_1, T_2]$. The infix operator op
may be an
arbitrary identifier.
All type infix operators have the same precedence; parentheses have to be used for grouping. The associativity of a type operator is determined as for term operators: type operators ending in a colon ‘:’ are rightassociative; all other operators are leftassociative.
In a sequence of consecutive type infix operations $t_0 \, \mathit{op} \, t_1 \, \mathit{op_2} \, \ldots \, \mathit{op_n} \, t_n$, all operators $\mathit{op}_1 , \ldots , \mathit{op}_n$ must have the same associativity. If they are all leftassociative, the sequence is interpreted as $(\ldots (t_0 \mathit{op_1} t_1) \mathit{op_2} \ldots) \mathit{op_n} t_n$, otherwise it is interpreted as $t_0 \mathit{op_1} (t_1 \mathit{op_2} ( \ldots \mathit{op_n} t_n) \ldots)$.
Function Types
The type $(T_1 , \ldots , T_n) \Rightarrow U$ represents the set of function values that take arguments of types $T1 , \ldots , Tn$ and yield results of type $U$. In the case of exactly one argument type $T \Rightarrow U$ is a shorthand for $(T) \Rightarrow U$. An argument type of the form $\Rightarrow T$ represents a callbyname parameter of type $T$.
Function types associate to the right, e.g. $S \Rightarrow T \Rightarrow U$ is the same as $S \Rightarrow (T \Rightarrow U)$.
Function types are shorthands for class types that define apply
functions. Specifically, the $n$ary function type
$(T_1 , \ldots , T_n) \Rightarrow U$ is a shorthand for the class type
Function$_n$[T1 , … , $T_n$, U]
. Such class
types are defined in the Scala library for $n$ between 0 and 22 as follows.
Hence, function types are covariant in their result type and contravariant in their argument types.
Existential Types
An existential type has the form $T$ forSome { $Q$ }
where $Q$ is a sequence of
type declarations.
Let
$t_1[\mathit{tps}_1] >: L_1 <: U_1 , \ldots , t_n[\mathit{tps}_n] >: L_n <: U_n$
be the types declared in $Q$ (any of the
type parameter sections [ $\mathit{tps}_i$ ]
might be missing).
The scope of each type $t_i$ includes the type $T$ and the existential clause
$Q$.
The type variables $t_i$ are said to be bound in the type
$T$ forSome { $Q$ }
.
Type variables which occur in a type $T$ but which are not bound in $T$ are said
to be free in $T$.
A type instance of $T$ forSome { $Q$ }
is a type $\sigma T$ where $\sigma$ is a substitution over $t_1 , \ldots , t_n$
such that, for each $i$, $\sigma L_i <: \sigma t_i <: \sigma U_i$.
The set of values denoted by the existential type $T$ forSome {$\,Q\,$}
is the union of the set of values of all its type instances.
A skolemization of $T$ forSome { $Q$ }
is
a type instance $\sigma T$, where $\sigma$ is the substitution
$[t_1'/t_1 , \ldots , t_n'/t_n]$ and each $t_i'$ is a fresh abstract type
with lower bound $\sigma L_i$ and upper bound $\sigma U_i$.
Simplification Rules
Existential types obey the following four equivalences:
 Multiple forclauses in an existential type can be merged. E.g.,
$T$ forSome { $Q$ } forSome { $Q'$ }
is equivalent to$T$ forSome { $Q$ ; $Q'$}
.  Unused quantifications can be dropped. E.g.,
$T$ forSome { $Q$ ; $Q'$}
where none of the types defined in $Q'$ are referred to by $T$ or $Q$, is equivalent to$T$ forSome {$ Q $}
.  An empty quantification can be dropped. E.g.,
$T$ forSome { }
is equivalent to $T$.  An existential type
$T$ forSome { $Q$ }
where $Q$ contains a clausetype $t[\mathit{tps}] >: L <: U$
is equivalent to the type$T'$ forSome { $Q$ }
where $T'$ results from $T$ by replacing every covariant occurrence of $t$ in $T$ by $U$ and by replacing every contravariant occurrence of $t$ in $T$ by $L$.
Existential Quantification over Values
As a syntactic convenience, the bindings clause
in an existential type may also contain
value declarations val $x$: $T$
.
An existential type $T$ forSome { $Q$; val $x$: $S\,$;$\,Q'$ }
is treated as a shorthand for the type
$T'$ forSome { $Q$; type $t$ <: $S$ with Singleton; $Q'$ }
, where $t$ is a
fresh type name and $T'$ results from $T$ by replacing every occurrence of
$x$.type
with $t$.
Placeholder Syntax for Existential Types
Scala supports a placeholder syntax for existential types.
A wildcard type is of the form _$\;$>:$\,L\,$<:$\,U$
. Both bound
clauses may be omitted. If a lower bound clause >:$\,L$
is missing,
>:$\,$scala.Nothing
is assumed. If an upper bound clause <:$\,U$
is missing,
<:$\,$scala.Any
is assumed. A wildcard type is a shorthand for an
existentially quantified type variable, where the existential quantification is
implicit.
A wildcard type must appear as type argument of a parameterized type.
Let $T = p.c[\mathit{targs},T,\mathit{targs}']$ be a parameterized type where
$\mathit{targs}, \mathit{targs}'$ may be empty and
$T$ is a wildcard type _$\;$>:$\,L\,$<:$\,U$
. Then $T$ is equivalent to the
existential
type
where $t$ is some fresh type variable. Wildcard types may also appear as parts of infix types , function types, or tuple types. Their expansion is then the expansion in the equivalent parameterized type.
Example
Assume the class definitions
Here are some examples of existential types:
The last two types in this list are equivalent. An alternative formulation of the first type above using wildcard syntax is:
Example
The type List[List[_]]
is equivalent to the existential type
Example
Assume a covariant type
The type
is equivalent (by simplification rule 4 above) to
which is in turn equivalent (by simplification rules 2 and 3 above) to
List[java.lang.Number]
.
NonValue Types
The types explained in the following do not denote sets of values, nor do they appear explicitly in programs. They are introduced in this report as the internal types of defined identifiers.
Method Types
A method type is denoted internally as $(\mathit{Ps})U$, where $(\mathit{Ps})$ is a sequence of parameter names and types $(p_1:T_1 , \ldots , p_n:T_n)$ for some $n \geq 0$ and $U$ is a (value or method) type. This type represents named methods that take arguments named $p_1 , \ldots , p_n$ of types $T_1 , \ldots , T_n$ and that return a result of type $U$.
Method types associate to the right: $(\mathit{Ps}_1)(\mathit{Ps}_2)U$ is treated as $(\mathit{Ps}_1)((\mathit{Ps}_2)U)$.
A special case are types of methods without any parameters. They are
written here => T
. Parameterless methods name expressions
that are reevaluated each time the parameterless method name is
referenced.
Method types do not exist as types of values. If a method name is used as a value, its type is implicitly converted to a corresponding function type.
Example
The declarations
produce the typings
Polymorphic Method Types
A polymorphic method type is denoted internally as [$\mathit{tps}\,$]$T$
where
[$\mathit{tps}\,$]
is a type parameter section
[$a_1$ >: $L_1$ <: $U_1 , \ldots , a_n$ >: $L_n$ <: $U_n$]
for some $n \geq 0$ and $T$ is a
(value or method) type. This type represents named methods that
take type arguments $S_1 , \ldots , S_n$
which
conform to the lower bounds
$L_1 , \ldots , L_n$
and the upper bounds
$U_1 , \ldots , U_n$
and that yield results of type $T$.
Example
The declarations
produce the typings
Type Constructors
A type constructor is represented internally much like a polymorphic method type.
[$\pm$ $a_1$ >: $L_1$ <: $U_1 , \ldots , \pm a_n$ >: $L_n$ <: $U_n$] $T$
represents a type that is expected by a
type constructor parameter or an
abstract type constructor binding with
the corresponding type parameter clause.
Example
Consider this fragment of the Iterable[+X]
class:
Conceptually, the type constructor Iterable
is a name for the
anonymous type [+X] Iterable[X]
, which may be passed to the
newType
type constructor parameter in flatMap
.
Base Types and Member Definitions
Types of class members depend on the way the members are referenced. Central here are three notions, namely:
 the notion of the set of base types of a type $T$,
 the notion of a type $T$ in some class $C$ seen from some prefix type $S$,
 the notion of the set of member bindings of some type $T$.
These notions are defined mutually recursively as follows.
The set of base types of a type is a set of class types, given as follows.
 The base types of a class type $C$ with parents $T_1 , \ldots , T_n$ are
$C$ itself, as well as the base types of the compound type
$T_1$ with … with $T_n$ { $R$ }
.  The base types of an aliased type are the base types of its alias.
 The base types of an abstract type are the base types of its upper bound.
 The base types of a parameterized type
$C$[$T_1 , \ldots , T_n$]
are the base types of type $C$, where every occurrence of a type parameter $a_i$ of $C$ has been replaced by the corresponding parameter type $T_i$.  The base types of a singleton type
$p$.type
are the base types of the type of $p$.  The base types of a compound type
$T_1$ with $\ldots$ with $T_n$ { $R$ }
are the reduced union of the base classes of all $T_i$'s. This means: Let the multiset $\mathscr{S}$ be the multisetunion of the base types of all $T_i$'s. If $\mathscr{S}$ contains several type instances of the same class, say$S^i$#$C$[$T^i_1 , \ldots , T^i_n$]
$(i \in I)$, then all those instances are replaced by one of them which conforms to all others. It is an error if no such instance exists. It follows that the reduced union, if it exists, produces a set of class types, where different types are instances of different classes.  The base types of a type selection
$S$#$T$
are determined as follows. If $T$ is an alias or abstract type, the previous clauses apply. Otherwise, $T$ must be a (possibly parameterized) class type, which is defined in some class $B$. Then the base types of$S$#$T$
are the base types of $T$ in $B$ seen from the prefix type $S$.  The base types of an existential type
$T$ forSome { $Q$ }
are all types$S$ forSome { $Q$ }
where $S$ is a base type of $T$.
 The base types of a class type $C$ with parents $T_1 , \ldots , T_n$ are
$C$ itself, as well as the base types of the compound type
The notion of a type $T$ in class $C$ seen from some prefix type $S$ makes sense only if the prefix type $S$ has a type instance of class $C$ as a base type, say
$S'$#$C$[$T_1 , \ldots , T_n$]
. Then we define as follows. If
$S$ = $\epsilon$.type
, then $T$ in $C$ seen from $S$ is $T$ itself.  Otherwise, if $S$ is an existential type
$S'$ forSome { $Q$ }
, and $T$ in $C$ seen from $S'$ is $T'$, then $T$ in $C$ seen from $S$ is$T'$ forSome {$\,Q\,$}
.  Otherwise, if $T$ is the $i$'th type parameter of some class $D$, then
 If $S$ has a base type
$D$[$U_1 , \ldots , U_n$]
, for some type parameters[$U_1 , \ldots , U_n$]
, then $T$ in $C$ seen from $S$ is $U_i$.  Otherwise, if $C$ is defined in a class $C'$, then $T$ in $C$ seen from $S$ is the same as $T$ in $C'$ seen from $S'$.
 Otherwise, if $C$ is not defined in another class, then $T$ in $C$ seen from $S$ is $T$ itself.
 If $S$ has a base type
 Otherwise, if $T$ is the singleton type
$D$.this.type
for some class $D$ then If $D$ is a subclass of $C$ and $S$ has a type instance of class $D$ among its base types, then $T$ in $C$ seen from $S$ is $S$.
 Otherwise, if $C$ is defined in a class $C'$, then $T$ in $C$ seen from $S$ is the same as $T$ in $C'$ seen from $S'$.
 Otherwise, if $C$ is not defined in another class, then $T$ in $C$ seen from $S$ is $T$ itself.
 If $T$ is some other type, then the described mapping is performed to all its type components.
If $T$ is a possibly parameterized class type, where $T$'s class is defined in some other class $D$, and $S$ is some prefix type, then we use "$T$ seen from $S$" as a shorthand for "$T$ in $D$ seen from $S$".
 If
The member bindings of a type $T$ are
 all bindings $d$ such that there exists a type instance of some class $C$ among the base types of $T$ and there exists a definition or declaration $d'$ in $C$ such that $d$ results from $d'$ by replacing every type $T'$ in $d'$ by $T'$ in $C$ seen from $T$, and
 all bindings of the type's refinement, if it has one.
The definition of a type projection S#T
is the member
binding $d_T$ of the type T
in S
. In that case, we also say
that S#T
is defined by $d_T$.
Relations between types
We define the following relations between types.
Name  Symbolically  Interpretation 

Equivalence  $T \equiv U$  $T$ and $U$ are interchangeable in all contexts. 
Conformance  $T <: U$  Type $T$ conforms to ("is a subtype of") type $U$. 
Weak Conformance  $T <:_w U$  Augments conformance for primitive numeric types. 
Compatibility  Type $T$ conforms to type $U$ after conversions. 
Equivalence
Equivalence $(\equiv)$ between types is the smallest congruence ^{3} such that the following holds:
 If $t$ is defined by a type alias
type $t$ = $T$
, then $t$ is equivalent to $T$.  If a path $p$ has a singleton type
$q$.type
, then$p$.type $\equiv q$.type
.  If $O$ is defined by an object definition, and $p$ is a path consisting only of package or object selectors and ending in $O$, then
$O$.this.type $\equiv p$.type
.  Two compound types are equivalent if the sequences of their component are pairwise equivalent, and occur in the same order, and their refinements are equivalent. Two refinements are equivalent if they bind the same names and the modifiers, types and bounds of every declared entity are equivalent in both refinements.
 Two method types are equivalent if:
 neither are implicit, or they both are ^{4};
 they have equivalent result types;
 they have the same number of parameters; and
 corresponding parameters have equivalent types. Note that the names of parameters do not matter for method type equivalence.
 Two polymorphic method types are equivalent if they have the same number of type parameters, and, after renaming one set of type parameters by another, the result types as well as lower and upper bounds of corresponding type parameters are equivalent.
 Two existential types are equivalent if they have the same number of quantifiers, and, after renaming one list of type quantifiers by another, the quantified types as well as lower and upper bounds of corresponding quantifiers are equivalent.
 Two type constructors are equivalent if they have the same number of type parameters, and, after renaming one list of type parameters by another, the result types as well as variances, lower and upper bounds of corresponding type parameters are equivalent.
Conformance
The conformance relation $(<:)$ is the smallest transitive relation that satisfies the following conditions.
 Conformance includes equivalence. If $T \equiv U$ then $T <: U$.
 For every value type $T$,
scala.Nothing <: $T$ <: scala.Any
.  For every type constructor $T$ (with any number of type parameters),
scala.Nothing <: $T$ <: scala.Any
.  For every class type $T$ such that
$T$ <: scala.AnyRef
one hasscala.Null <: $T$
.  A type variable or abstract type $t$ conforms to its upper bound and its lower bound conforms to $t$.
 A class type or parameterized type conforms to any of its basetypes.
 A singleton type
$p$.type
conforms to the type of the path $p$.  A singleton type
$p$.type
conforms to the typescala.Singleton
.  A type projection
$T$#$t$
conforms to$U$#$t$
if $T$ conforms to $U$.  A parameterized type
$T$[$T_1$ , … , $T_n$]
conforms to$T$[$U_1$ , … , $U_n$]
if the following three conditions hold for $i \in { 1 , \ldots , n }$: If the $i$'th type parameter of $T$ is declared covariant, then $T_i <: U_i$.
 If the $i$'th type parameter of $T$ is declared contravariant, then $U_i <: T_i$.
 If the $i$'th type parameter of $T$ is declared neither covariant nor contravariant, then $U_i \equiv T_i$.
 A compound type
$T_1$ with $\ldots$ with $T_n$ {$R\,$}
conforms to each of its component types $T_i$.  If $T <: U_i$ for $i \in { 1 , \ldots , n }$ and for every
binding $d$ of a type or value $x$ in $R$ there exists a member
binding of $x$ in $T$ which subsumes $d$, then $T$ conforms to the
compound type
$U_1$ with $\ldots$ with $U_n$ {$R\,$}
.  The existential type
$T$ forSome {$\,Q\,$}
conforms to $U$ if its skolemization conforms to $U$.  The type $T$ conforms to the existential type
$U$ forSome {$\,Q\,$}
if $T$ conforms to one of the type instances of$U$ forSome {$\,Q\,$}
.  If $T_i \equiv T_i'$ for $i \in { 1 , \ldots , n}$ and $U$ conforms to $U'$ then the method type $(p_1:T_1 , \ldots , p_n:T_n) U$ conforms to $(p_1':T_1' , \ldots , p_n':T_n') U'$.
 The polymorphic type $[a_1 >: L_1 <: U_1 , \ldots , a_n >: L_n <: U_n] T$ conforms to the polymorphic type $[a_1 >: L_1' <: U_1' , \ldots , a_n >: L_n' <: U_n'] T'$ if, assuming $L_1' <: a_1 <: U_1' , \ldots , L_n' <: a_n <: U_n'$ one has $T <: T'$ and $L_i <: L_i'$ and $U_i' <: U_i$ for $i \in { 1 , \ldots , n }$.
 Type constructors $T$ and $T'$ follow a similar discipline. We characterize
$T$ and $T'$ by their type parameter clauses
$[a_1 , \ldots , a_n]$ and
$[a_1' , \ldots , a_n']$, where an $a_i$ or $a_i'$ may include a variance
annotation, a higherorder type parameter clause, and bounds. Then, $T$
conforms to $T'$ if any list $[t_1 , \ldots , t_n]$  with declared
variances, bounds and higherorder type parameter clauses  of valid type
arguments for $T'$ is also a valid list of type arguments for $T$ and
$T[t_1 , \ldots , t_n] <: T'[t_1 , \ldots , t_n]$. Note that this entails
that:
 The bounds on $a_i$ must be weaker than the corresponding bounds declared for $a'_i$.
 The variance of $a_i$ must match the variance of $a'_i$, where covariance matches covariance, contravariance matches contravariance and any variance matches invariance.
 Recursively, these restrictions apply to the corresponding higherorder type parameter clauses of $a_i$ and $a'_i$.
A declaration or definition in some compound type of class type $C$ subsumes another declaration of the same name in some compound type or class type $C'$, if one of the following holds.
 A value declaration or definition that defines a name $x$ with type $T$ subsumes a value or method declaration that defines $x$ with type $T'$, provided $T <: T'$.
 A method declaration or definition that defines a name $x$ with type $T$ subsumes a method declaration that defines $x$ with type $T'$, provided $T <: T'$.
 A type alias
type $t$[$T_1$ , … , $T_n$] = $T$
subsumes a type aliastype $t$[$T_1$ , … , $T_n$] = $T'$
if $T \equiv T'$.  A type declaration
type $t$[$T_1$ , … , $T_n$] >: $L$ <: $U$
subsumes a type declarationtype $t$[$T_1$ , … , $T_n$] >: $L'$ <: $U'$
if $L' <: L$ and $U <: U'$.  A type or class definition that binds a type name $t$ subsumes an abstract
type declaration
type t[$T_1$ , … , $T_n$] >: L <: U
if $L <: t <: U$.
Least upper bounds and greatest lower bounds
The $(<:)$ relation forms preorder between types, i.e. it is transitive and reflexive. This allows us to define least upper bounds and greatest lower bounds of a set of types in terms of that order. The least upper bound or greatest lower bound of a set of types does not always exist. For instance, consider the class definitions:
Then the types A[Any], A[A[Any]], A[A[A[Any]]], ...
form
a descending sequence of upper bounds for B
and C
. The
least upper bound would be the infinite limit of that sequence, which
does not exist as a Scala type. Since cases like this are in general
impossible to detect, a Scala compiler is free to reject a term
which has a type specified as a least upper or greatest lower bound,
and that bound would be more complex than some compilerset
limit ^{5}.
The least upper bound or greatest lower bound might also not be
unique. For instance A with B
and B with A
are both
greatest lower bounds of A
and B
. If there are several
least upper bounds or greatest lower bounds, the Scala compiler is
free to pick any one of them.
Weak Conformance
In some situations Scala uses a more general conformance relation. A type $S$ weakly conforms to a type $T$, written $S <:_w T$, if $S <: T$ or both $S$ and $T$ are primitive number types and $S$ precedes $T$ in the following ordering.
A weak least upper bound is a least upper bound with respect to weak conformance.
Compatibility
A type $T$ is compatible to a type $U$ if $T$ (or its corresponding function type) weakly conforms to $U$ after applying etaexpansion. If $T$ is a method type, it's converted to the corresponding function type. If the types do not weakly conform, the following alternatives are checked in order:
 view application: there's an implicit view from $T$ to $U$;
 dropping byname modifiers: if $U$ is of the shape
$=> U'$
(and $T$ is not),$T <:_w U'$
;  SAM conversion: if $T$ corresponds to a function type, and $U$ declares a single abstract method whose type corresponds to the function type $U'$,
$T <:_w U'$
.
Examples
Function compatibility via SAM conversion
Given the definitions
The application foo((x: Int) => x.toString)
resolves to the first overload,
as it's more specific:
Int => String
is compatible toToString
 when expecting a value of typeToString
, you may pass a function literal fromInt
toString
, as it will be SAMconverted to said function;ToString
is not compatible toInt => String
 when expecting a function fromInt
toString
, you may not pass aToString
.
Volatile Types
Type volatility approximates the possibility that a type parameter or abstract type instance of a type does not have any nonnull values. A value member of a volatile type cannot appear in a path.
A type is volatile if it falls into one of four categories:
A compound type $T_1$ with … with $T_n$ {$R\,$}
is volatile if one of the following two conditions hold.
 One of $T_2 , \ldots , T_n$ is a type parameter or abstract type, or
 $T_1$ is an abstract type and and either the refinement $R$ or a type $T_j$ for $j > 1$ contributes an abstract member to the compound type, or
 one of $T_1 , \ldots , T_n$ is a singleton type.
Here, a type $S$ contributes an abstract member to a type $T$ if $S$ contains an abstract member that is also a member of $T$. A refinement $R$ contributes an abstract member to a type $T$ if $R$ contains an abstract declaration which is also a member of $T$.
A type designator is volatile if it is an alias of a volatile type, or if it designates a type parameter or abstract type that has a volatile type as its upper bound.
A singleton type $p$.type
is volatile, if the underlying
type of path $p$ is volatile.
An existential type $T$ forSome {$\,Q\,$}
is volatile if
$T$ is volatile.
Type Erasure
A type is called generic if it contains type arguments or type variables. Type erasure is a mapping from (possibly generic) types to nongeneric types. We write $T$ for the erasure of type $T$. The erasure mapping is defined as follows.
 The erasure of an alias type is the erasure of its righthand side.
 The erasure of an abstract type is the erasure of its upper bound.
 The erasure of the parameterized type
scala.Array$[T_1]$
isscala.Array$[T_1]$
.  The erasure of every other parameterized type $T[T_1 , \ldots , T_n]$ is $T$.
 The erasure of a singleton type
$p$.type
is the erasure of the type of $p$.  The erasure of a type projection
$T$#$x$
is$T$#$x$
.  The erasure of a compound type
$T_1$ with $\ldots$ with $T_n$ {$R\,$}
is the erasure of the intersection dominator of $T_1 , \ldots , T_n$.  The erasure of an existential type
$T$ forSome {$\,Q\,$}
is $T$.
The intersection dominator of a list of types $T_1 , \ldots , T_n$ is computed as follows. Let $T_{i_1} , \ldots , T_{i_m}$ be the subsequence of types $T_i$ which are not supertypes of some other type $T_j$. If this subsequence contains a type designator $T_c$ that refers to a class which is not a trait, the intersection dominator is $T_c$. Otherwise, the intersection dominator is the first element of the subsequence, $T_{i_1}$.

We assume that objects and packages also implicitly define a class (of the same name as the object or package, but inaccessible to user programs). ↩

A reference to a structurally defined member (method call or access to a value or variable) may generate binary code that is significantly slower than an equivalent code to a nonstructural member. ↩

A congruence is an equivalence relation which is closed under formation of contexts. ↩

A method type is implicit if the parameter section that defines it starts with the
implicit
keyword. ↩ 
The current Scala compiler limits the nesting level of parameterization in such bounds to be at most two deeper than the maximum nesting level of the operand types ↩